2005-04-16 15:20:36 -07:00
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2011-12-07 13:06:11 -05:00
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2006-01-11 12:17:46 -08:00
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2006-12-08 02:37:56 -08:00
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2012-07-26 21:08:32 -07:00
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2005-04-16 15:20:36 -07:00
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2017-02-02 17:54:15 +01:00
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2008-03-26 22:11:34 +01:00
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2013-09-11 14:26:10 -07:00
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2011-12-07 13:06:11 -05:00
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2013-04-12 01:50:06 +01:00
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2013-05-01 17:51:54 -07:00
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2014-02-28 13:46:44 -05:00
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2014-08-08 13:08:20 -04:00
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2017-02-04 01:20:53 +01:00
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2005-11-07 17:19:07 -05:00
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2007-07-15 23:41:25 -07:00
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2005-04-16 15:20:36 -07:00
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2016-09-28 00:27:17 -05:00
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2014-02-28 13:46:44 -05:00
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2008-02-06 01:37:57 -08:00
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2014-02-27 14:40:10 -05:00
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2008-03-26 22:11:34 +01:00
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2008-03-27 13:06:23 +01:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2009-06-24 03:12:00 -04:00
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2005-04-16 15:20:36 -07:00
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2014-03-20 21:10:51 -04:00
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2014-02-28 13:46:44 -05:00
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2006-12-06 20:33:20 -08:00
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2013-09-16 21:34:53 -04:00
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2005-04-16 15:20:36 -07:00
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2006-01-16 22:14:23 -08:00
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2007-10-29 14:17:23 -06:00
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2006-01-16 22:14:23 -08:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2013-09-29 22:06:07 -04:00
|
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|
|
fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
|
|
|
|
2014-03-20 21:10:51 -04:00
|
|
|
|
2005-04-16 15:20:36 -07:00
|
|
|
|
2005-11-07 17:16:09 -05:00
|
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2014-02-28 13:46:44 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 20:38:33 -05:00
|
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|
2008-03-26 22:11:34 +01:00
|
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|
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|
|
|
|
|
|
fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
|
|
|
|
2011-11-25 00:50:41 -05:00
|
|
|
|
2009-06-24 03:12:00 -04:00
|
|
|
|
2011-11-25 00:50:41 -05:00
|
|
|
|
fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
|
|
|
|
2008-03-26 22:11:34 +01:00
|
|
|
|
|
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|
2011-11-24 20:38:33 -05:00
|
|
|
|
2008-03-26 22:11:34 +01:00
|
|
|
|
2011-11-25 00:50:41 -05:00
|
|
|
|
fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
|
|
|
|
2009-06-24 03:12:00 -04:00
|
|
|
|
|
|
|
|
|
|
|
|
|
|
fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
|
|
|
|
2008-03-26 22:11:34 +01:00
|
|
|
|
|
|
|
|
|
2008-03-27 13:06:23 +01:00
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
2011-11-24 19:54:23 -05:00
|
|
|
|
2008-03-27 13:06:23 +01:00
|
|
|
|
2009-06-24 03:12:00 -04:00
|
|
|
|
|
|
|
|
|
2008-03-27 13:06:23 +01:00
|
|
|
|
|
|
|
|
|
|
|
|
|
|
2009-06-24 03:12:00 -04:00
|
|
|
|
|
|
|
|
|
2011-11-25 00:50:41 -05:00
|
|
|
|
2009-06-24 03:12:00 -04:00
|
|
|
|
2011-11-25 00:50:41 -05:00
|
|
|
|
2009-06-24 03:12:00 -04:00
|
|
|
|
|
|
|
|
|
2008-03-27 13:06:23 +01:00
|
|
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|
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|
|
2011-11-24 19:54:23 -05:00
|
|
|
|
2008-03-27 13:06:23 +01:00
|
|
|
|
2011-11-25 00:50:41 -05:00
|
|
|
|
2009-06-24 03:12:00 -04:00
|
|
|
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|
|
|
|
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|
2011-11-25 00:50:41 -05:00
|
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|
2008-03-27 13:06:23 +01:00
|
|
|
|
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|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
|
|
|
|
|
|
|
|
|
|
|
|
|
2011-11-24 22:37:54 -05:00
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
|
|
|
|
|
|
|
|
2011-11-24 22:53:09 -05:00
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
|
|
|
|
|
|
|
|
2011-11-24 22:53:09 -05:00
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
|
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|
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2011-11-24 22:37:54 -05:00
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
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|
2011-01-14 22:30:21 -05:00
|
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|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
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|
|
|
|
2011-11-24 22:53:09 -05:00
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
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|
2011-11-24 22:53:09 -05:00
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
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2015-01-10 19:01:08 -05:00
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2011-11-24 20:38:33 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 02:35:16 -05:00
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2008-03-26 22:11:34 +01:00
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2011-11-25 02:35:16 -05:00
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2008-07-21 18:06:36 +08:00
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2015-02-13 14:36:41 -08:00
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2011-11-25 02:35:16 -05:00
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2008-07-21 18:06:36 +08:00
|
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2008-03-26 22:11:34 +01:00
|
|
|
|
|
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
|
|
|
2011-11-25 02:35:16 -05:00
|
|
|
|
|
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
|
|
|
|
|
|
|
|
2011-11-25 02:35:16 -05:00
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
|
|
|
2011-11-25 02:35:16 -05:00
|
|
|
|
|
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
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2014-03-20 21:10:51 -04:00
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2011-11-25 02:35:16 -05:00
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2013-09-22 19:37:01 -07:00
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2016-10-24 16:16:13 -05:00
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2015-01-10 19:01:08 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 02:35:16 -05:00
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2008-07-21 18:06:36 +08:00
|
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2009-04-26 20:25:54 +10:00
|
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2015-02-13 14:36:41 -08:00
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2009-04-26 20:25:54 +10:00
|
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|
2008-07-21 18:06:36 +08:00
|
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2011-11-25 02:35:16 -05:00
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2008-07-21 18:06:36 +08:00
|
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2011-11-25 02:35:16 -05:00
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2008-07-21 18:06:36 +08:00
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2005-04-16 15:20:36 -07:00
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|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
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2008-02-15 14:38:00 -08:00
|
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|
2017-07-17 08:45:34 +01:00
|
|
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|
2008-02-15 14:38:00 -08:00
|
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|
|
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|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
|
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|
|
2011-11-24 22:37:54 -05:00
|
|
|
|
2009-04-26 20:25:54 +10:00
|
|
|
|
|
|
|
|
|
2011-11-24 22:53:09 -05:00
|
|
|
|
2009-04-26 20:25:54 +10:00
|
|
|
|
2011-11-24 22:53:09 -05:00
|
|
|
|
2009-04-26 20:25:54 +10:00
|
|
|
|
|
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
|
|
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|
2011-11-24 22:37:54 -05:00
|
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|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
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2009-04-26 20:25:54 +10:00
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2011-11-24 22:53:09 -05:00
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2009-04-26 20:25:54 +10:00
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2011-11-24 22:53:09 -05:00
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|
|
2009-04-26 20:25:54 +10:00
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
|
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|
2011-11-24 22:37:54 -05:00
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
|
|
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|
2009-04-26 20:25:54 +10:00
|
|
|
|
|
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
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|
|
|
2011-11-24 22:53:09 -05:00
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
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2009-04-26 20:25:54 +10:00
|
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|
|
|
|
|
|
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|
|
|
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
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2011-11-21 12:11:31 +01:00
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2008-02-15 14:37:30 -08:00
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2012-06-12 16:20:35 +02:00
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2008-02-15 14:37:30 -08:00
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2012-06-12 16:20:35 +02:00
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2011-11-24 22:37:54 -05:00
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2008-02-15 14:37:30 -08:00
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2012-06-12 16:20:35 +02:00
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2008-02-15 14:37:30 -08:00
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2012-06-12 16:20:35 +02:00
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2008-02-15 14:37:30 -08:00
|
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|
2011-11-24 22:37:54 -05:00
|
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|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
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2009-04-26 20:25:54 +10:00
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2011-01-07 17:50:10 +11:00
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2009-04-26 20:25:54 +10:00
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2011-01-07 17:50:10 +11:00
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2009-04-26 20:25:54 +10:00
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locking/atomics: COCCINELLE/treewide: Convert trivial ACCESS_ONCE() patterns to READ_ONCE()/WRITE_ONCE()
Please do not apply this to mainline directly, instead please re-run the
coccinelle script shown below and apply its output.
For several reasons, it is desirable to use {READ,WRITE}_ONCE() in
preference to ACCESS_ONCE(), and new code is expected to use one of the
former. So far, there's been no reason to change most existing uses of
ACCESS_ONCE(), as these aren't harmful, and changing them results in
churn.
However, for some features, the read/write distinction is critical to
correct operation. To distinguish these cases, separate read/write
accessors must be used. This patch migrates (most) remaining
ACCESS_ONCE() instances to {READ,WRITE}_ONCE(), using the following
coccinelle script:
----
// Convert trivial ACCESS_ONCE() uses to equivalent READ_ONCE() and
// WRITE_ONCE()
// $ make coccicheck COCCI=/home/mark/once.cocci SPFLAGS="--include-headers" MODE=patch
virtual patch
@ depends on patch @
expression E1, E2;
@@
- ACCESS_ONCE(E1) = E2
+ WRITE_ONCE(E1, E2)
@ depends on patch @
expression E;
@@
- ACCESS_ONCE(E)
+ READ_ONCE(E)
----
Signed-off-by: Mark Rutland <mark.rutland@arm.com>
Signed-off-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: davem@davemloft.net
Cc: linux-arch@vger.kernel.org
Cc: mpe@ellerman.id.au
Cc: shuah@kernel.org
Cc: snitzer@redhat.com
Cc: thor.thayer@linux.intel.com
Cc: tj@kernel.org
Cc: viro@zeniv.linux.org.uk
Cc: will.deacon@arm.com
Link: http://lkml.kernel.org/r/1508792849-3115-19-git-send-email-paulmck@linux.vnet.ibm.com
Signed-off-by: Ingo Molnar <mingo@kernel.org>
2017-10-23 14:07:29 -07:00
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2009-04-26 20:25:54 +10:00
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2011-11-21 12:11:31 +01:00
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2011-01-07 17:50:10 +11:00
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|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
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2009-04-26 20:25:54 +10:00
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2012-06-12 16:20:35 +02:00
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[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
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2008-02-15 14:37:30 -08:00
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2009-04-26 20:25:55 +10:00
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2011-11-24 22:37:54 -05:00
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2009-04-26 20:25:55 +10:00
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2012-06-12 16:20:35 +02:00
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2009-04-26 20:25:55 +10:00
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2012-06-12 16:20:35 +02:00
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2009-04-26 20:25:55 +10:00
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2012-06-12 16:20:35 +02:00
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2009-04-26 20:25:55 +10:00
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2014-03-14 12:02:47 -04:00
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2012-06-12 16:20:35 +02:00
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2009-04-26 20:25:55 +10:00
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2012-06-12 16:20:35 +02:00
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2017-09-05 12:53:12 +02:00
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2012-06-12 16:20:35 +02:00
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2017-09-05 12:53:12 +02:00
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2012-06-12 16:20:35 +02:00
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2017-09-05 12:53:12 +02:00
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2012-06-12 16:20:35 +02:00
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2017-09-05 12:53:12 +02:00
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2017-09-18 22:53:43 +03:00
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2017-09-05 12:53:12 +02:00
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2009-04-26 20:25:55 +10:00
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2008-02-15 14:37:30 -08:00
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2012-06-12 16:20:35 +02:00
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2008-02-15 14:37:30 -08:00
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2012-06-12 16:20:35 +02:00
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2008-02-15 14:37:30 -08:00
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2012-06-12 16:20:35 +02:00
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2008-02-15 14:37:30 -08:00
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2009-04-26 20:25:54 +10:00
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2011-11-24 22:37:54 -05:00
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2009-04-26 20:25:54 +10:00
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2008-02-15 14:37:30 -08:00
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2012-06-12 16:20:35 +02:00
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2008-02-15 14:37:30 -08:00
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2012-06-12 16:20:35 +02:00
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2017-09-05 12:53:12 +02:00
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2011-12-09 08:06:57 -05:00
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2017-09-05 12:53:12 +02:00
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2011-12-09 08:06:57 -05:00
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2011-11-24 22:37:54 -05:00
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2008-02-15 14:37:30 -08:00
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|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
|
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|
|
2013-09-29 11:24:49 -04:00
|
|
|
|
2011-11-24 22:37:54 -05:00
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
|
|
|
|
2009-04-26 20:25:54 +10:00
|
|
|
|
|
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
|
|
|
|
2009-04-26 20:25:54 +10:00
|
|
|
|
|
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
|
|
|
|
2009-04-26 20:25:54 +10:00
|
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|
|
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
|
|
|
|
2011-01-07 17:50:10 +11:00
|
|
|
|
2009-04-26 20:25:54 +10:00
|
|
|
|
|
|
|
|
|
2011-11-24 22:37:54 -05:00
|
|
|
|
2009-04-26 20:25:54 +10:00
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
2011-11-24 22:37:54 -05:00
|
|
|
|
2013-09-29 11:24:49 -04:00
|
|
|
|
[PATCH] r/o bind mounts: track numbers of writers to mounts
This is the real meat of the entire series. It actually
implements the tracking of the number of writers to a mount.
However, it causes scalability problems because there can be
hundreds of cpus doing open()/close() on files on the same mnt at
the same time. Even an atomic_t in the mnt has massive scalaing
problems because the cacheline gets so terribly contended.
This uses a statically-allocated percpu variable. All want/drop
operations are local to a cpu as long that cpu operates on the same
mount, and there are no writer count imbalances. Writer count
imbalances happen when a write is taken on one cpu, and released
on another, like when an open/close pair is performed on two
Upon a remount,ro request, all of the data from the percpu
variables is collected (expensive, but very rare) and we determine
if there are any outstanding writers to the mount.
I've written a little benchmark to sit in a loop for a couple of
seconds in several cpus in parallel doing open/write/close loops.
http://sr71.net/~dave/linux/openbench.c
The code in here is a a worst-possible case for this patch. It
does opens on a _pair_ of files in two different mounts in parallel.
This should cause my code to lose its "operate on the same mount"
optimization completely. This worst-case scenario causes a 3%
degredation in the benchmark.
I could probably get rid of even this 3%, but it would be more
complex than what I have here, and I think this is getting into
acceptable territory. In practice, I expect writing more than 3
bytes to a file, as well as disk I/O to mask any effects that this
has.
(To get rid of that 3%, we could have an #defined number of mounts
in the percpu variable. So, instead of a CPU getting operate only
on percpu data when it accesses only one mount, it could stay on
percpu data when it only accesses N or fewer mounts.)
[AV] merged fix for __clear_mnt_mount() stepping on freed vfsmount
Acked-by: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Christoph Hellwig <hch@infradead.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2008-02-15 14:37:59 -08:00
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2008-02-15 14:37:30 -08:00
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2011-11-24 22:37:54 -05:00
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2008-02-15 14:38:00 -08:00
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2013-09-29 11:24:49 -04:00
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2011-11-24 22:37:54 -05:00
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2013-09-29 11:24:49 -04:00
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2008-02-15 14:38:00 -08:00
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2011-11-21 12:11:31 +01:00
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2011-11-21 12:11:33 +01:00
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2013-09-29 11:24:49 -04:00
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2011-11-21 12:11:31 +01:00
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2011-11-21 12:11:33 +01:00
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2011-11-21 12:11:31 +01:00
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2013-09-29 11:24:49 -04:00
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2011-11-21 12:11:31 +01:00
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2011-11-24 20:38:33 -05:00
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2005-04-16 15:20:36 -07:00
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2015-02-13 14:36:41 -08:00
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2009-04-26 20:25:54 +10:00
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2011-11-24 22:53:09 -05:00
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2009-04-26 20:25:54 +10:00
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2011-11-24 20:38:33 -05:00
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2005-04-16 15:20:36 -07:00
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2014-01-24 12:17:54 +00:00
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2013-09-29 22:06:07 -04:00
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2015-05-08 11:43:53 -04:00
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2013-09-29 22:06:07 -04:00
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2015-05-08 11:43:53 -04:00
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2013-09-29 22:06:07 -04:00
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2015-05-08 11:43:53 -04:00
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2013-09-29 22:06:07 -04:00
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2015-05-08 11:43:53 -04:00
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2013-09-29 22:06:07 -04:00
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2015-05-08 11:43:53 -04:00
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2013-09-29 22:06:07 -04:00
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2005-04-16 15:20:36 -07:00
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2013-10-01 16:11:26 -04:00
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2013-09-29 22:06:07 -04:00
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2005-04-16 15:20:36 -07:00
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2013-10-01 16:11:26 -04:00
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2005-04-16 15:20:36 -07:00
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2014-03-20 21:10:51 -04:00
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2013-10-01 16:11:26 -04:00
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2014-03-20 21:10:51 -04:00
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2013-10-01 16:11:26 -04:00
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2005-11-07 17:20:17 -05:00
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2012-06-25 12:55:28 +01:00
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2005-11-07 17:20:17 -05:00
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2016-11-20 19:45:28 -05:00
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2005-11-07 17:20:17 -05:00
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2011-11-24 18:22:03 -05:00
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2013-09-29 22:06:07 -04:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2013-09-29 22:06:07 -04:00
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2005-11-07 17:20:17 -05:00
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2013-10-04 19:15:13 -07:00
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2014-02-24 17:32:34 -08:00
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2013-03-15 10:53:28 -04:00
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2014-02-28 13:46:44 -05:00
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2013-03-15 10:53:28 -04:00
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2014-02-28 13:46:44 -05:00
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2013-03-15 10:53:28 -04:00
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2014-02-24 17:32:34 -08:00
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2017-01-03 14:18:43 +13:00
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2014-02-24 17:32:34 -08:00
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2017-01-03 14:18:43 +13:00
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2014-02-24 17:32:34 -08:00
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2013-03-15 10:53:28 -04:00
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2017-01-03 14:18:43 +13:00
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2013-03-15 10:53:28 -04:00
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2017-01-03 14:18:43 +13:00
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2013-09-05 14:39:11 +02:00
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2017-01-03 14:18:43 +13:00
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2013-03-15 10:53:28 -04:00
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2013-09-22 19:37:01 -07:00
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2013-03-15 10:53:28 -04:00
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2014-02-28 13:46:44 -05:00
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2013-03-15 10:53:28 -04:00
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2011-11-25 00:46:35 -05:00
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2005-04-16 15:20:36 -07:00
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2006-12-08 02:37:56 -08:00
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2005-04-16 15:20:36 -07:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2006-12-08 02:37:56 -08:00
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2005-11-07 17:15:49 -05:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
|
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2006-12-08 02:37:56 -08:00
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2005-11-07 17:15:49 -05:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2014-12-29 13:03:41 -06:00
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2011-11-24 19:41:16 -05:00
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2011-11-24 22:19:58 -05:00
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2011-11-24 22:25:07 -05:00
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2011-11-24 23:24:33 -05:00
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2014-03-20 21:10:51 -04:00
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2013-09-22 19:37:01 -07:00
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2013-03-15 10:53:28 -04:00
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2005-04-16 15:20:36 -07:00
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2014-12-29 13:03:41 -06:00
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2015-01-15 22:58:33 -06:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2013-03-15 10:53:28 -04:00
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2011-11-24 21:28:22 -05:00
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2005-11-07 17:19:50 -05:00
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2013-03-15 10:53:28 -04:00
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2011-11-25 03:19:09 -05:00
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2013-03-15 10:53:28 -04:00
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2011-11-25 03:19:09 -05:00
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2013-03-15 10:53:28 -04:00
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2013-09-22 19:37:01 -07:00
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2005-11-07 17:19:50 -05:00
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2017-01-20 18:28:35 +13:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2013-03-15 10:53:28 -04:00
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2005-04-16 15:20:36 -07:00
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2013-03-15 10:53:28 -04:00
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2017-01-20 18:28:35 +13:00
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2005-11-07 17:19:50 -05:00
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2017-01-20 18:28:35 +13:00
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2014-08-10 03:44:55 -04:00
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2017-01-20 18:28:35 +13:00
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2014-08-10 03:44:55 -04:00
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2005-11-07 17:19:50 -05:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2005-11-07 17:19:50 -05:00
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2017-01-20 18:28:35 +13:00
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2005-11-07 17:19:50 -05:00
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2011-11-24 22:19:58 -05:00
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2011-11-24 22:37:54 -05:00
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2005-11-07 17:19:50 -05:00
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2011-11-25 00:46:35 -05:00
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2005-11-07 17:19:50 -05:00
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2011-11-24 22:19:58 -05:00
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2005-11-07 17:19:50 -05:00
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2011-11-25 02:19:55 -05:00
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2012-06-09 00:59:08 -04:00
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2011-11-25 00:46:35 -05:00
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2011-01-14 22:30:21 -05:00
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2005-11-07 17:19:50 -05:00
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2016-09-28 00:27:17 -05:00
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2017-01-20 18:28:35 +13:00
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2006-12-08 02:37:56 -08:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 03:06:56 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 23:24:33 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 03:06:56 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 23:24:33 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 22:19:58 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 23:24:33 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 18:57:30 -05:00
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2005-11-07 17:21:20 -05:00
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2011-11-24 23:24:33 -05:00
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2005-11-07 17:21:20 -05:00
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2011-03-17 22:08:28 -04:00
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2011-11-24 20:38:33 -05:00
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2011-03-17 22:08:28 -04:00
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2017-07-17 08:45:35 +01:00
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2011-11-24 20:38:33 -05:00
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2011-03-17 22:08:28 -04:00
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2014-01-24 12:17:54 +00:00
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2011-03-17 22:08:28 -04:00
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2011-11-24 20:38:33 -05:00
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2011-11-24 22:25:07 -05:00
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2011-11-24 22:19:58 -05:00
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2013-09-29 11:24:49 -04:00
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2011-11-21 12:11:30 +01:00
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2013-09-29 11:24:49 -04:00
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2011-11-24 20:38:33 -05:00
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2011-03-17 22:08:28 -04:00
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2017-02-01 06:06:16 +13:00
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2017-07-17 08:45:35 +01:00
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2017-02-01 06:06:16 +13:00
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2011-11-24 21:24:27 -05:00
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2005-11-07 17:17:22 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 21:24:27 -05:00
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2012-06-25 12:55:18 +01:00
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2005-04-16 15:20:36 -07:00
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2012-06-25 12:55:18 +01:00
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2008-03-27 13:06:23 +01:00
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2012-07-31 13:13:04 -07:00
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2012-06-25 12:55:18 +01:00
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2005-11-07 17:19:50 -05:00
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2012-06-25 12:55:18 +01:00
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2005-04-16 15:20:36 -07:00
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2012-06-25 12:55:18 +01:00
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smarter propagate_mnt()
The current mainline has copies propagated to *all* nodes, then
tears down the copies we made for nodes that do not contain
counterparts of the desired mountpoint. That sets the right
propagation graph for the copies (at teardown time we move
the slaves of removed node to a surviving peer or directly
to master), but we end up paying a fairly steep price in
useless allocations. It's fairly easy to create a situation
where N calls of mount(2) create exactly N bindings, with
O(N^2) vfsmounts allocated and freed in process.
Fortunately, it is possible to avoid those allocations/freeings.
The trick is to create copies in the right order and find which
one would've eventually become a master with the current algorithm.
It turns out to be possible in O(nodes getting propagation) time
and with no extra allocations at all.
One part is that we need to make sure that eventual master will be
created before its slaves, so we need to walk the propagation
tree in a different order - by peer groups. And iterate through
the peers before dealing with the next group.
Another thing is finding the (earlier) copy that will be a master
of one we are about to create; to do that we are (temporary) marking
the masters of mountpoints we are attaching the copies to.
Either we are in a peer of the last mountpoint we'd dealt with,
or we have the following situation: we are attaching to mountpoint M,
the last copy S_0 had been attached to M_0 and there are sequences
S_0...S_n, M_0...M_n such that S_{i+1} is a master of S_{i},
S_{i} mounted on M{i} and we need to create a slave of the first S_{k}
such that M is getting propagation from M_{k}. It means that the master
of M_{k} will be among the sequence of masters of M. On the
other hand, the nearest marked node in that sequence will either
be the master of M_{k} or the master of M_{k-1} (the latter -
in the case if M_{k-1} is a slave of something M gets propagation
from, but in a wrong peer group).
So we go through the sequence of masters of M until we find
a marked one (P). Let N be the one before it. Then we go through
the sequence of masters of S_0 until we find one (say, S) mounted
on a node D that has P as master and check if D is a peer of N.
If it is, S will be the master of new copy, if not - the master of S
will be.
That's it for the hard part; the rest is fairly simple. Iterator
is in next_group(), handling of one prospective mountpoint is
propagate_one().
It seems to survive all tests and gives a noticably better performance
than the current mainline for setups that are seriously using shared
subtrees.
Cc: stable@vger.kernel.org
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2014-02-27 09:35:45 -05:00
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2013-03-22 04:08:05 -07:00
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|
mnt: Correct permission checks in do_remount
While invesgiating the issue where in "mount --bind -oremount,ro ..."
would result in later "mount --bind -oremount,rw" succeeding even if
the mount started off locked I realized that there are several
additional mount flags that should be locked and are not.
In particular MNT_NOSUID, MNT_NODEV, MNT_NOEXEC, and the atime
flags in addition to MNT_READONLY should all be locked. These
flags are all per superblock, can all be changed with MS_BIND,
and should not be changable if set by a more privileged user.
The following additions to the current logic are added in this patch.
- nosuid may not be clearable by a less privileged user.
- nodev may not be clearable by a less privielged user.
- noexec may not be clearable by a less privileged user.
- atime flags may not be changeable by a less privileged user.
The logic with atime is that always setting atime on access is a
global policy and backup software and auditing software could break if
atime bits are not updated (when they are configured to be updated),
and serious performance degradation could result (DOS attack) if atime
updates happen when they have been explicitly disabled. Therefore an
unprivileged user should not be able to mess with the atime bits set
by a more privileged user.
The additional restrictions are implemented with the addition of
MNT_LOCK_NOSUID, MNT_LOCK_NODEV, MNT_LOCK_NOEXEC, and MNT_LOCK_ATIME
mnt flags.
Taken together these changes and the fixes for MNT_LOCK_READONLY
should make it safe for an unprivileged user to create a user
namespace and to call "mount --bind -o remount,... ..." without
the danger of mount flags being changed maliciously.
Cc: stable@vger.kernel.org
Acked-by: Serge E. Hallyn <serge.hallyn@ubuntu.com>
Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2014-07-28 17:26:07 -07:00
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2013-03-22 04:08:05 -07:00
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2013-03-29 21:04:39 -07:00
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2014-10-07 17:11:46 -07:00
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2013-03-29 21:04:39 -07:00
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2012-06-25 12:55:18 +01:00
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2013-09-29 11:24:49 -04:00
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2012-06-25 12:55:18 +01:00
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2013-09-29 11:24:49 -04:00
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2012-06-25 12:55:18 +01:00
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2012-07-31 13:13:04 -07:00
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2012-06-25 12:55:18 +01:00
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2016-11-20 19:33:09 -05:00
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2012-06-25 12:55:18 +01:00
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2011-11-24 20:55:08 -05:00
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2008-03-27 13:06:23 +01:00
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2014-01-24 12:17:54 +00:00
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2008-03-27 13:06:23 +01:00
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2012-06-25 12:55:18 +01:00
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2005-04-16 15:20:36 -07:00
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2014-08-08 13:08:20 -04:00
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2017-08-07 17:44:45 +09:00
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2014-08-08 13:08:20 -04:00
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2017-08-07 17:44:45 +09:00
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2014-08-08 13:08:20 -04:00
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2011-11-25 00:33:11 -05:00
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|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
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2013-09-29 22:06:07 -04:00
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2011-01-14 22:30:21 -05:00
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fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
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2013-09-29 11:24:49 -04:00
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|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
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2013-09-29 22:06:07 -04:00
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2013-09-29 11:24:49 -04:00
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|
fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2013-09-29 22:06:07 -04:00
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2012-05-08 13:32:02 +09:30
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2011-11-21 12:11:30 +01:00
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2014-12-23 21:37:03 -06:00
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2013-09-29 11:24:49 -04:00
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2013-09-28 12:41:25 -04:00
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2014-08-08 13:08:20 -04:00
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fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
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2011-11-25 00:57:42 -05:00
|
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|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
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2011-11-25 00:57:42 -05:00
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|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
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2011-11-24 22:37:54 -05:00
|
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|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
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2016-11-24 08:03:41 +11:00
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2016-11-20 19:45:28 -05:00
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2005-11-07 17:13:39 -05:00
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2014-08-07 09:12:31 -04:00
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2005-11-07 17:13:39 -05:00
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2005-04-16 15:20:36 -07:00
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2008-03-27 13:06:24 +01:00
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2011-12-06 12:21:54 -05:00
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2005-04-16 15:20:36 -07:00
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fs: use seq_open_private() for proc_mounts
A patchset to remove support for passing pre-allocated struct seq_file to
seq_open(). Such feature is undocumented and prone to error.
In particular, if seq_release() is used in release handler, it will
kfree() a pointer which was not allocated by seq_open().
So this patchset drops support for pre-allocated struct seq_file: it's
only of use in proc_namespace.c and can be easily replaced by using
seq_open_private()/seq_release_private().
Additionally, it documents the use of file->private_data to hold pointer
to struct seq_file by seq_open().
This patch (of 3):
Since patch described below, from v2.6.15-rc1, seq_open() could use a
struct seq_file already allocated by the caller if the pointer to the
structure is stored in file->private_data before calling the function.
Commit 1abe77b0fc4b485927f1f798ae81a752677e1d05
Author: Al Viro <viro@zeniv.linux.org.uk>
Date: Mon Nov 7 17:15:34 2005 -0500
[PATCH] allow callers of seq_open do allocation themselves
Allow caller of seq_open() to kmalloc() seq_file + whatever else they
want and set ->private_data to it. seq_open() will then abstain from
doing allocation itself.
Such behavior is only used by mounts_open_common().
In order to drop support for such uncommon feature, proc_mounts is
converted to use seq_open_private(), which take care of allocating the
proc_mounts structure, making it available through ->private in struct
seq_file.
Conversely, proc_mounts is converted to use seq_release_private(), in
order to release the private structure allocated by seq_open_private().
Then, ->private is used directly instead of proc_mounts() macro to access
to the proc_mounts structure.
Link: http://lkml.kernel.org/r/cover.1433193673.git.ydroneaud@opteya.com
Signed-off-by: Yann Droneaud <ydroneaud@opteya.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-30 14:57:30 -07:00
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|
2005-04-16 15:20:36 -07:00
|
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|
2005-11-07 17:17:51 -05:00
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2014-02-27 14:40:10 -05:00
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2005-04-16 15:20:36 -07:00
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fs: use seq_open_private() for proc_mounts
A patchset to remove support for passing pre-allocated struct seq_file to
seq_open(). Such feature is undocumented and prone to error.
In particular, if seq_release() is used in release handler, it will
kfree() a pointer which was not allocated by seq_open().
So this patchset drops support for pre-allocated struct seq_file: it's
only of use in proc_namespace.c and can be easily replaced by using
seq_open_private()/seq_release_private().
Additionally, it documents the use of file->private_data to hold pointer
to struct seq_file by seq_open().
This patch (of 3):
Since patch described below, from v2.6.15-rc1, seq_open() could use a
struct seq_file already allocated by the caller if the pointer to the
structure is stored in file->private_data before calling the function.
Commit 1abe77b0fc4b485927f1f798ae81a752677e1d05
Author: Al Viro <viro@zeniv.linux.org.uk>
Date: Mon Nov 7 17:15:34 2005 -0500
[PATCH] allow callers of seq_open do allocation themselves
Allow caller of seq_open() to kmalloc() seq_file + whatever else they
want and set ->private_data to it. seq_open() will then abstain from
doing allocation itself.
Such behavior is only used by mounts_open_common().
In order to drop support for such uncommon feature, proc_mounts is
converted to use seq_open_private(), which take care of allocating the
proc_mounts structure, making it available through ->private in struct
seq_file.
Conversely, proc_mounts is converted to use seq_release_private(), in
order to release the private structure allocated by seq_open_private().
Then, ->private is used directly instead of proc_mounts() macro to access
to the proc_mounts structure.
Link: http://lkml.kernel.org/r/cover.1433193673.git.ydroneaud@opteya.com
Signed-off-by: Yann Droneaud <ydroneaud@opteya.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-30 14:57:30 -07:00
|
|
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|
2007-07-15 23:39:55 -07:00
|
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|
2014-02-27 14:40:10 -05:00
|
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|
2005-04-16 15:20:36 -07:00
|
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2005-11-07 17:17:51 -05:00
|
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2005-04-16 15:20:36 -07:00
|
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|
2011-12-06 12:21:54 -05:00
|
|
|
|
2008-03-27 13:06:25 +01:00
|
|
|
|
fs: use seq_open_private() for proc_mounts
A patchset to remove support for passing pre-allocated struct seq_file to
seq_open(). Such feature is undocumented and prone to error.
In particular, if seq_release() is used in release handler, it will
kfree() a pointer which was not allocated by seq_open().
So this patchset drops support for pre-allocated struct seq_file: it's
only of use in proc_namespace.c and can be easily replaced by using
seq_open_private()/seq_release_private().
Additionally, it documents the use of file->private_data to hold pointer
to struct seq_file by seq_open().
This patch (of 3):
Since patch described below, from v2.6.15-rc1, seq_open() could use a
struct seq_file already allocated by the caller if the pointer to the
structure is stored in file->private_data before calling the function.
Commit 1abe77b0fc4b485927f1f798ae81a752677e1d05
Author: Al Viro <viro@zeniv.linux.org.uk>
Date: Mon Nov 7 17:15:34 2005 -0500
[PATCH] allow callers of seq_open do allocation themselves
Allow caller of seq_open() to kmalloc() seq_file + whatever else they
want and set ->private_data to it. seq_open() will then abstain from
doing allocation itself.
Such behavior is only used by mounts_open_common().
In order to drop support for such uncommon feature, proc_mounts is
converted to use seq_open_private(), which take care of allocating the
proc_mounts structure, making it available through ->private in struct
seq_file.
Conversely, proc_mounts is converted to use seq_release_private(), in
order to release the private structure allocated by seq_open_private().
Then, ->private is used directly instead of proc_mounts() macro to access
to the proc_mounts structure.
Link: http://lkml.kernel.org/r/cover.1433193673.git.ydroneaud@opteya.com
Signed-off-by: Yann Droneaud <ydroneaud@opteya.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-06-30 14:57:30 -07:00
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2011-11-25 02:19:55 -05:00
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2011-12-06 12:21:54 -05:00
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2005-04-16 15:20:36 -07:00
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2008-03-27 13:06:24 +01:00
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2005-04-16 15:20:36 -07:00
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2011-12-06 12:21:54 -05:00
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2006-03-20 13:44:12 -05:00
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2008-03-27 13:06:24 +01:00
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2006-03-20 13:44:12 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 03:06:56 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 03:06:56 -05:00
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2005-11-07 17:17:22 -05:00
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2011-11-24 18:57:30 -05:00
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2011-11-25 03:06:56 -05:00
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2005-04-16 15:20:36 -07:00
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fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
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2013-09-29 11:24:49 -04:00
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2011-11-25 03:06:56 -05:00
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2011-11-24 22:37:54 -05:00
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2005-04-16 15:20:36 -07:00
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2013-09-29 11:24:49 -04:00
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2005-04-16 15:20:36 -07:00
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2006-03-27 01:14:51 -08:00
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2005-04-16 15:20:36 -07:00
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2006-03-27 01:14:51 -08:00
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2005-04-16 15:20:36 -07:00
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2006-03-27 01:14:51 -08:00
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2010-01-16 12:56:08 -05:00
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2013-09-29 11:24:49 -04:00
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2011-11-24 21:35:16 -05:00
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2006-03-27 01:14:51 -08:00
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2013-09-29 11:24:49 -04:00
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2010-01-16 12:56:08 -05:00
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2005-11-07 17:20:17 -05:00
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2005-04-16 15:20:36 -07:00
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2014-03-20 21:10:51 -04:00
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2013-03-16 14:35:16 -04:00
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2013-03-16 15:12:40 -04:00
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2005-11-07 17:17:04 -05:00
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2015-04-02 20:33:53 -05:00
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2013-03-16 15:12:40 -04:00
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2015-04-02 20:33:53 -05:00
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2013-03-16 15:12:40 -04:00
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2015-04-02 20:33:53 -05:00
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2013-09-29 22:06:07 -04:00
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2015-01-10 19:01:08 -05:00
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2005-11-07 17:17:04 -05:00
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2013-03-16 15:12:40 -04:00
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2013-03-16 14:35:16 -04:00
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2013-03-16 15:12:40 -04:00
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2013-03-16 14:35:16 -04:00
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2014-12-24 07:20:01 -06:00
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2015-04-01 18:30:06 -05:00
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2014-12-24 07:20:01 -06:00
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2015-07-17 14:15:30 -05:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2013-09-29 22:06:07 -04:00
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|
fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2014-12-24 07:20:01 -06:00
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2005-04-16 15:20:36 -07:00
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2014-12-18 13:10:48 -06:00
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2011-11-24 18:57:30 -05:00
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2005-04-16 15:20:36 -07:00
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2015-01-03 05:39:35 -06:00
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2014-12-18 13:10:48 -06:00
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2014-12-22 18:30:08 -06:00
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2014-12-18 13:10:48 -06:00
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2014-12-22 18:30:08 -06:00
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2005-04-16 15:20:36 -07:00
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2014-12-22 19:12:07 -06:00
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2014-12-18 13:10:48 -06:00
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2014-08-18 15:09:26 -04:00
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2014-12-18 13:10:48 -06:00
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2014-08-18 15:09:26 -04:00
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2014-12-18 13:10:48 -06:00
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2014-12-24 07:20:01 -06:00
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2011-01-15 20:08:44 -05:00
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2005-11-07 17:20:17 -05:00
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2014-12-18 13:10:48 -06:00
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2016-09-28 00:27:17 -05:00
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2014-12-23 21:37:03 -06:00
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2014-12-18 13:10:48 -06:00
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2011-11-25 00:22:05 -05:00
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2011-11-25 02:19:55 -05:00
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2016-09-28 00:27:17 -05:00
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2011-11-25 00:46:35 -05:00
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2014-12-24 07:20:01 -06:00
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2013-09-29 22:06:07 -04:00
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2015-01-10 19:01:08 -05:00
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2015-07-17 14:15:30 -05:00
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2014-12-23 21:37:03 -06:00
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2011-11-24 21:47:05 -05:00
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2014-08-30 18:32:05 -04:00
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2014-12-23 21:37:03 -06:00
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2008-03-21 23:59:49 -04:00
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2011-11-24 20:43:10 -05:00
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2005-04-16 15:20:36 -07:00
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2013-03-16 14:39:34 -04:00
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2008-03-22 00:46:23 -04:00
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2011-11-24 21:35:16 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 21:35:16 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 21:35:16 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 21:35:16 -05:00
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2005-04-16 15:20:36 -07:00
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fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
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|
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2013-09-29 11:24:49 -04:00
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|
2011-11-24 22:37:54 -05:00
|
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|
|
2013-09-29 11:24:49 -04:00
|
|
|
|
2005-04-16 15:20:36 -07:00
|
|
|
|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
|
|
|
|
2013-09-29 11:24:49 -04:00
|
|
|
|
2005-04-16 15:20:36 -07:00
|
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2011-11-25 00:57:42 -05:00
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2005-04-16 15:20:36 -07:00
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2008-04-24 07:21:56 -04:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 21:35:16 -05:00
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2005-04-16 15:20:36 -07:00
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2014-10-08 12:32:47 -07:00
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2005-04-16 15:20:36 -07:00
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2017-07-17 08:45:34 +01:00
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2017-07-17 08:45:35 +01:00
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2005-04-16 15:20:36 -07:00
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2013-03-16 15:12:40 -04:00
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2013-09-29 11:24:49 -04:00
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2005-11-07 17:15:49 -05:00
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2005-04-16 15:20:36 -07:00
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2013-09-29 22:06:07 -04:00
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2011-11-25 02:19:55 -05:00
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2014-12-24 07:20:01 -06:00
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2005-04-16 15:20:36 -07:00
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2013-09-29 22:06:07 -04:00
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2014-12-24 07:20:01 -06:00
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2013-09-29 22:06:07 -04:00
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2005-04-16 15:20:36 -07:00
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2013-09-29 11:24:49 -04:00
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2013-03-16 14:35:16 -04:00
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2005-04-16 15:20:36 -07:00
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2013-10-03 01:31:18 -07:00
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2017-01-03 14:18:43 +13:00
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2013-10-03 01:31:18 -07:00
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2015-01-19 11:48:45 -06:00
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2013-10-03 01:31:18 -07:00
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2016-04-15 14:24:41 -07:00
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2013-10-03 01:31:18 -07:00
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2014-12-23 21:37:03 -06:00
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2015-07-17 14:54:27 -05:00
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2014-12-23 21:37:03 -06:00
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2015-04-01 18:30:06 -05:00
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2013-10-03 01:31:18 -07:00
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2017-01-03 14:18:43 +13:00
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2013-10-03 01:31:18 -07:00
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2017-07-04 17:25:09 +01:00
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2013-02-22 22:45:42 -05:00
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2015-11-16 09:49:34 -05:00
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2015-11-11 17:22:33 -06:00
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2015-11-16 09:49:34 -05:00
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2005-04-16 15:20:36 -07:00
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2009-01-14 14:14:12 +01:00
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2005-04-16 15:20:36 -07:00
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2008-07-22 09:59:21 -04:00
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2011-11-25 00:33:11 -05:00
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2005-04-16 15:20:36 -07:00
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2010-02-10 12:15:53 +01:00
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2005-04-16 15:20:36 -07:00
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2010-02-10 12:15:53 +01:00
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2013-02-22 22:45:42 -05:00
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2010-02-10 12:15:53 +01:00
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2013-09-08 14:03:27 -04:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 00:33:11 -05:00
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2005-04-16 15:20:36 -07:00
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2008-07-22 09:59:21 -04:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 00:46:35 -05:00
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2005-04-16 15:20:36 -07:00
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2013-03-29 21:04:39 -07:00
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2014-10-04 14:44:03 -07:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 00:33:11 -05:00
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2005-04-16 15:20:36 -07:00
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2008-02-14 19:34:31 -08:00
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2008-07-22 09:59:21 -04:00
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2011-11-25 00:33:11 -05:00
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2005-04-16 15:20:36 -07:00
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2005-11-07 17:16:09 -05:00
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2005-04-16 15:20:36 -07:00
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2009-01-14 14:14:12 +01:00
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2005-04-16 15:20:36 -07:00
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2005-11-07 17:16:09 -05:00
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2005-04-16 15:20:36 -07:00
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2013-03-30 01:35:18 -07:00
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2010-03-07 18:49:36 -08:00
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2013-03-30 01:35:18 -07:00
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|
take the targets of /proc/*/ns/* symlinks to separate fs
New pseudo-filesystem: nsfs. Targets of /proc/*/ns/* live there now.
It's not mountable (not even registered, so it's not in /proc/filesystems,
etc.). Files on it *are* bindable - we explicitly permit that in do_loopback().
This stuff lives in fs/nsfs.c now; proc_ns_fget() moved there as well.
get_proc_ns() is a macro now (it's simply returning ->i_private; would
have been an inline, if not for header ordering headache).
proc_ns_inode() is an ex-parrot. The interface used in procfs is
ns_get_path(path, task, ops) and ns_get_name(buf, size, task, ops).
Dentries and inodes are never hashed; a non-counting reference to dentry
is stashed in ns_common (removed by ->d_prune()) and reused by ns_get_path()
if present. See ns_get_path()/ns_prune_dentry/nsfs_evict() for details
of that mechanism.
As the result, proc_ns_follow_link() has stopped poking in nd->path.mnt;
it does nd_jump_link() on a consistent <vfsmount,dentry> pair it gets
from ns_get_path().
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2014-11-01 10:57:28 -04:00
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2013-03-30 01:35:18 -07:00
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2014-11-01 00:00:23 -04:00
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2013-03-30 01:35:18 -07:00
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2014-11-01 03:13:17 -04:00
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2010-03-07 18:49:36 -08:00
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2011-11-24 21:24:27 -05:00
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2005-11-07 17:17:22 -05:00
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2005-04-16 15:20:36 -07:00
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2013-03-15 10:53:28 -04:00
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2005-04-16 15:20:36 -07:00
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2013-03-30 01:35:18 -07:00
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2012-06-25 12:55:18 +01:00
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2005-11-07 17:21:20 -05:00
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2005-11-07 17:17:22 -05:00
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2012-06-25 12:55:18 +01:00
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2011-11-24 22:25:07 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 23:24:33 -05:00
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2011-11-24 18:57:30 -05:00
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2008-04-29 00:59:40 -07:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 03:06:56 -05:00
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2013-03-30 01:35:18 -07:00
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2005-11-07 17:21:20 -05:00
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2011-11-24 22:19:58 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 21:24:27 -05:00
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2013-03-15 10:53:28 -04:00
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2011-11-24 21:24:27 -05:00
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2012-06-25 12:55:18 +01:00
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2013-09-29 11:24:49 -04:00
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2011-11-25 02:19:55 -05:00
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2017-01-20 18:28:35 +13:00
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2013-09-29 11:24:49 -04:00
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2005-04-16 15:20:36 -07:00
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2012-06-25 12:55:18 +01:00
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2005-04-16 15:20:36 -07:00
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2013-09-29 11:24:49 -04:00
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2014-12-24 07:20:01 -06:00
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2013-09-29 11:24:49 -04:00
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2005-04-16 15:20:36 -07:00
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2012-06-25 12:55:18 +01:00
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2005-04-16 15:20:36 -07:00
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2012-06-25 12:55:18 +01:00
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2016-11-20 19:45:28 -05:00
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2007-06-07 12:20:32 -04:00
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2011-11-24 20:55:08 -05:00
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2013-03-16 15:12:40 -04:00
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2015-01-07 14:28:26 -06:00
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2013-03-16 14:49:45 -04:00
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2012-06-25 12:55:18 +01:00
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2013-08-14 12:44:39 +03:00
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2012-06-25 12:55:18 +01:00
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2007-06-07 12:20:32 -04:00
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2013-03-16 15:12:40 -04:00
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2013-09-29 11:24:49 -04:00
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2014-12-24 07:20:01 -06:00
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2013-09-29 11:24:49 -04:00
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2013-03-16 14:42:19 -04:00
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2007-06-07 12:20:32 -04:00
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2014-10-24 00:14:36 +02:00
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2016-11-20 19:45:28 -05:00
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2014-10-24 00:14:36 +02:00
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2010-01-30 22:51:25 -05:00
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2011-11-25 02:19:55 -05:00
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2010-01-30 22:51:25 -05:00
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2011-11-25 02:19:55 -05:00
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2010-01-30 22:51:25 -05:00
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2011-11-24 19:54:23 -05:00
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2008-03-27 13:06:23 +01:00
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2011-11-24 18:57:30 -05:00
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2008-03-27 13:06:23 +01:00
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2011-11-25 03:06:56 -05:00
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2011-11-25 01:05:37 -05:00
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2011-11-24 19:54:23 -05:00
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2008-03-27 13:06:23 +01:00
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2011-11-24 19:54:23 -05:00
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2008-03-27 13:06:23 +01:00
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2011-11-24 18:57:30 -05:00
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2008-03-27 13:06:23 +01:00
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2011-11-25 03:06:56 -05:00
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2011-11-25 01:05:37 -05:00
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2011-11-24 19:54:23 -05:00
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2008-03-27 13:06:23 +01:00
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2011-11-24 19:54:23 -05:00
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2008-03-27 13:06:23 +01:00
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2016-09-28 00:27:17 -05:00
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2005-11-07 17:19:50 -05:00
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2005-11-07 17:20:03 -05:00
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2005-11-07 17:19:50 -05:00
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2005-11-07 17:21:20 -05:00
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2005-11-07 17:19:50 -05:00
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2005-11-07 17:21:01 -05:00
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2005-11-07 17:21:20 -05:00
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2005-11-07 17:21:01 -05:00
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2005-11-07 17:20:03 -05:00
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2005-11-07 17:21:01 -05:00
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2005-11-07 17:19:50 -05:00
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2011-11-24 19:59:16 -05:00
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2013-03-15 10:53:28 -04:00
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2005-11-07 17:19:50 -05:00
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2014-03-20 21:10:51 -04:00
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2016-09-28 00:27:17 -05:00
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2017-01-20 18:28:35 +13:00
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2011-11-24 18:57:30 -05:00
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2014-03-20 21:10:51 -04:00
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2008-03-27 13:06:23 +01:00
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2005-11-07 17:19:50 -05:00
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2017-01-20 18:28:35 +13:00
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2016-09-28 00:27:17 -05:00
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2011-11-25 01:05:37 -05:00
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2011-11-24 19:59:16 -05:00
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2008-03-27 13:06:23 +01:00
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2014-03-21 10:14:08 -04:00
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|
smarter propagate_mnt()
The current mainline has copies propagated to *all* nodes, then
tears down the copies we made for nodes that do not contain
counterparts of the desired mountpoint. That sets the right
propagation graph for the copies (at teardown time we move
the slaves of removed node to a surviving peer or directly
to master), but we end up paying a fairly steep price in
useless allocations. It's fairly easy to create a situation
where N calls of mount(2) create exactly N bindings, with
O(N^2) vfsmounts allocated and freed in process.
Fortunately, it is possible to avoid those allocations/freeings.
The trick is to create copies in the right order and find which
one would've eventually become a master with the current algorithm.
It turns out to be possible in O(nodes getting propagation) time
and with no extra allocations at all.
One part is that we need to make sure that eventual master will be
created before its slaves, so we need to walk the propagation
tree in a different order - by peer groups. And iterate through
the peers before dealing with the next group.
Another thing is finding the (earlier) copy that will be a master
of one we are about to create; to do that we are (temporary) marking
the masters of mountpoints we are attaching the copies to.
Either we are in a peer of the last mountpoint we'd dealt with,
or we have the following situation: we are attaching to mountpoint M,
the last copy S_0 had been attached to M_0 and there are sequences
S_0...S_n, M_0...M_n such that S_{i+1} is a master of S_{i},
S_{i} mounted on M{i} and we need to create a slave of the first S_{k}
such that M is getting propagation from M_{k}. It means that the master
of M_{k} will be among the sequence of masters of M. On the
other hand, the nearest marked node in that sequence will either
be the master of M_{k} or the master of M_{k-1} (the latter -
in the case if M_{k-1} is a slave of something M gets propagation
from, but in a wrong peer group).
So we go through the sequence of masters of M until we find
a marked one (P). Let N be the one before it. Then we go through
the sequence of masters of S_0 until we find one (say, S) mounted
on a node D that has P as master and check if D is a peer of N.
If it is, S will be the master of new copy, if not - the master of S
will be.
That's it for the hard part; the rest is fairly simple. Iterator
is in next_group(), handling of one prospective mountpoint is
propagate_one().
It seems to survive all tests and gives a noticably better performance
than the current mainline for setups that are seriously using shared
subtrees.
Cc: stable@vger.kernel.org
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2014-02-27 09:35:45 -05:00
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2014-03-21 10:14:08 -04:00
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2011-11-25 03:06:56 -05:00
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2011-11-24 20:43:10 -05:00
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|
2014-03-21 10:14:08 -04:00
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2005-11-07 17:19:50 -05:00
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2008-03-21 20:48:19 -04:00
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|
2011-11-24 19:59:16 -05:00
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2013-03-15 10:53:28 -04:00
|
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2011-11-25 00:46:35 -05:00
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|
2005-11-07 17:20:03 -05:00
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2013-03-15 10:53:28 -04:00
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2017-01-20 18:28:35 +13:00
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2005-11-07 17:20:03 -05:00
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2005-11-07 17:19:50 -05:00
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2014-03-20 21:10:51 -04:00
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2014-03-20 20:34:43 -04:00
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2014-03-20 21:10:51 -04:00
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2017-01-20 18:28:35 +13:00
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2005-11-07 17:19:50 -05:00
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2017-01-20 18:28:35 +13:00
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2013-09-29 11:24:49 -04:00
|
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|
fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
|
|
|
|
2005-11-07 17:19:50 -05:00
|
|
|
|
2008-03-27 13:06:23 +01:00
|
|
|
|
|
|
|
|
|
smarter propagate_mnt()
The current mainline has copies propagated to *all* nodes, then
tears down the copies we made for nodes that do not contain
counterparts of the desired mountpoint. That sets the right
propagation graph for the copies (at teardown time we move
the slaves of removed node to a surviving peer or directly
to master), but we end up paying a fairly steep price in
useless allocations. It's fairly easy to create a situation
where N calls of mount(2) create exactly N bindings, with
O(N^2) vfsmounts allocated and freed in process.
Fortunately, it is possible to avoid those allocations/freeings.
The trick is to create copies in the right order and find which
one would've eventually become a master with the current algorithm.
It turns out to be possible in O(nodes getting propagation) time
and with no extra allocations at all.
One part is that we need to make sure that eventual master will be
created before its slaves, so we need to walk the propagation
tree in a different order - by peer groups. And iterate through
the peers before dealing with the next group.
Another thing is finding the (earlier) copy that will be a master
of one we are about to create; to do that we are (temporary) marking
the masters of mountpoints we are attaching the copies to.
Either we are in a peer of the last mountpoint we'd dealt with,
or we have the following situation: we are attaching to mountpoint M,
the last copy S_0 had been attached to M_0 and there are sequences
S_0...S_n, M_0...M_n such that S_{i+1} is a master of S_{i},
S_{i} mounted on M{i} and we need to create a slave of the first S_{k}
such that M is getting propagation from M_{k}. It means that the master
of M_{k} will be among the sequence of masters of M. On the
other hand, the nearest marked node in that sequence will either
be the master of M_{k} or the master of M_{k-1} (the latter -
in the case if M_{k-1} is a slave of something M gets propagation
from, but in a wrong peer group).
So we go through the sequence of masters of M until we find
a marked one (P). Let N be the one before it. Then we go through
the sequence of masters of S_0 until we find one (say, S) mounted
on a node D that has P as master and check if D is a peer of N.
If it is, S will be the master of new copy, if not - the master of S
will be.
That's it for the hard part; the rest is fairly simple. Iterator
is in next_group(), handling of one prospective mountpoint is
propagate_one().
It seems to survive all tests and gives a noticably better performance
than the current mainline for setups that are seriously using shared
subtrees.
Cc: stable@vger.kernel.org
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2014-02-27 09:35:45 -05:00
|
|
|
|
|
|
|
|
|
2016-09-28 00:27:17 -05:00
|
|
|
|
2014-12-24 07:20:01 -06:00
|
|
|
|
smarter propagate_mnt()
The current mainline has copies propagated to *all* nodes, then
tears down the copies we made for nodes that do not contain
counterparts of the desired mountpoint. That sets the right
propagation graph for the copies (at teardown time we move
the slaves of removed node to a surviving peer or directly
to master), but we end up paying a fairly steep price in
useless allocations. It's fairly easy to create a situation
where N calls of mount(2) create exactly N bindings, with
O(N^2) vfsmounts allocated and freed in process.
Fortunately, it is possible to avoid those allocations/freeings.
The trick is to create copies in the right order and find which
one would've eventually become a master with the current algorithm.
It turns out to be possible in O(nodes getting propagation) time
and with no extra allocations at all.
One part is that we need to make sure that eventual master will be
created before its slaves, so we need to walk the propagation
tree in a different order - by peer groups. And iterate through
the peers before dealing with the next group.
Another thing is finding the (earlier) copy that will be a master
of one we are about to create; to do that we are (temporary) marking
the masters of mountpoints we are attaching the copies to.
Either we are in a peer of the last mountpoint we'd dealt with,
or we have the following situation: we are attaching to mountpoint M,
the last copy S_0 had been attached to M_0 and there are sequences
S_0...S_n, M_0...M_n such that S_{i+1} is a master of S_{i},
S_{i} mounted on M{i} and we need to create a slave of the first S_{k}
such that M is getting propagation from M_{k}. It means that the master
of M_{k} will be among the sequence of masters of M. On the
other hand, the nearest marked node in that sequence will either
be the master of M_{k} or the master of M_{k-1} (the latter -
in the case if M_{k-1} is a slave of something M gets propagation
from, but in a wrong peer group).
So we go through the sequence of masters of M until we find
a marked one (P). Let N be the one before it. Then we go through
the sequence of masters of S_0 until we find one (say, S) mounted
on a node D that has P as master and check if D is a peer of N.
If it is, S will be the master of new copy, if not - the master of S
will be.
That's it for the hard part; the rest is fairly simple. Iterator
is in next_group(), handling of one prospective mountpoint is
propagate_one().
It seems to survive all tests and gives a noticably better performance
than the current mainline for setups that are seriously using shared
subtrees.
Cc: stable@vger.kernel.org
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2014-02-27 09:35:45 -05:00
|
|
|
|
|
|
|
|
|
2014-03-21 10:14:08 -04:00
|
|
|
|
2008-03-27 13:06:23 +01:00
|
|
|
|
2016-09-28 00:27:17 -05:00
|
|
|
|
2017-01-20 18:28:35 +13:00
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
|
2008-03-27 13:06:23 +01:00
|
|
|
|
2005-11-07 17:19:50 -05:00
|
|
|
|
|
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2011-03-18 08:55:38 -04:00
|
|
|
|
|
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2011-03-18 08:55:38 -04:00
|
|
|
|
2016-01-22 15:40:57 -05:00
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2016-01-22 15:40:57 -05:00
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2011-03-18 08:55:38 -04:00
|
|
|
|
2013-03-16 15:12:40 -04:00
|
|
|
|
2011-03-18 08:55:38 -04:00
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2017-01-03 14:18:43 +13:00
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2013-03-16 15:12:40 -04:00
|
|
|
|
2016-01-22 15:40:57 -05:00
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
|
|
|
|
|
|
|
|
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|
|
2013-03-16 15:12:40 -04:00
|
|
|
|
2016-01-22 15:40:57 -05:00
|
|
|
|
2011-03-18 08:55:38 -04:00
|
|
|
|
|
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2011-03-18 08:55:38 -04:00
|
|
|
|
|
|
|
|
|
|
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2011-03-18 08:55:38 -04:00
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2017-01-03 14:18:43 +13:00
|
|
|
|
|
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2017-01-03 14:18:43 +13:00
|
|
|
|
|
|
|
|
|
2013-03-16 14:49:45 -04:00
|
|
|
|
2016-01-22 15:40:57 -05:00
|
|
|
|
2011-03-18 08:55:38 -04:00
|
|
|
|
|
|
|
|
|
2013-03-15 10:53:28 -04:00
|
|
|
|
2005-04-16 15:20:36 -07:00
|
|
|
|
2017-07-17 08:45:35 +01:00
|
|
|
|
2005-04-16 15:20:36 -07:00
|
|
|
|
|
|
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|
|
VFS: (Scripted) Convert S_ISLNK/DIR/REG(dentry->d_inode) to d_is_*(dentry)
Convert the following where appropriate:
(1) S_ISLNK(dentry->d_inode) to d_is_symlink(dentry).
(2) S_ISREG(dentry->d_inode) to d_is_reg(dentry).
(3) S_ISDIR(dentry->d_inode) to d_is_dir(dentry). This is actually more
complicated than it appears as some calls should be converted to
d_can_lookup() instead. The difference is whether the directory in
question is a real dir with a ->lookup op or whether it's a fake dir with
a ->d_automount op.
In some circumstances, we can subsume checks for dentry->d_inode not being
NULL into this, provided we the code isn't in a filesystem that expects
d_inode to be NULL if the dirent really *is* negative (ie. if we're going to
use d_inode() rather than d_backing_inode() to get the inode pointer).
Note that the dentry type field may be set to something other than
DCACHE_MISS_TYPE when d_inode is NULL in the case of unionmount, where the VFS
manages the fall-through from a negative dentry to a lower layer. In such a
case, the dentry type of the negative union dentry is set to the same as the
type of the lower dentry.
However, if you know d_inode is not NULL at the call site, then you can use
the d_is_xxx() functions even in a filesystem.
There is one further complication: a 0,0 chardev dentry may be labelled
DCACHE_WHITEOUT_TYPE rather than DCACHE_SPECIAL_TYPE. Strictly, this was
intended for special directory entry types that don't have attached inodes.
The following perl+coccinelle script was used:
use strict;
my @callers;
open($fd, 'git grep -l \'S_IS[A-Z].*->d_inode\' |') ||
die "Can't grep for S_ISDIR and co. callers";
@callers = <$fd>;
close($fd);
unless (@callers) {
print "No matches\n";
exit(0);
}
my @cocci = (
'@@',
'expression E;',
'@@',
'',
'- S_ISLNK(E->d_inode->i_mode)',
'+ d_is_symlink(E)',
'',
'@@',
'expression E;',
'@@',
'',
'- S_ISDIR(E->d_inode->i_mode)',
'+ d_is_dir(E)',
'',
'@@',
'expression E;',
'@@',
'',
'- S_ISREG(E->d_inode->i_mode)',
'+ d_is_reg(E)' );
my $coccifile = "tmp.sp.cocci";
open($fd, ">$coccifile") || die $coccifile;
print($fd "$_\n") || die $coccifile foreach (@cocci);
close($fd);
foreach my $file (@callers) {
chomp $file;
print "Processing ", $file, "\n";
system("spatch", "--sp-file", $coccifile, $file, "--in-place", "--no-show-diff") == 0 ||
die "spatch failed";
}
[AV: overlayfs parts skipped]
Signed-off-by: David Howells <dhowells@redhat.com>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2015-01-29 12:02:35 +00:00
|
|
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|
|
|
|
2005-04-16 15:20:36 -07:00
|
|
|
|
|
|
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|
2013-03-15 10:53:28 -04:00
|
|
|
|
2005-04-16 15:20:36 -07:00
|
|
|
|
|
|
|
|
|
2010-08-26 11:07:22 -07:00
|
|
|
|
|
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|
|
|
|
|
|
|
|
|
|
|
|
|
2017-07-17 08:45:35 +01:00
|
|
|
|
2010-08-26 11:07:22 -07:00
|
|
|
|
2017-07-17 08:45:35 +01:00
|
|
|
|
2010-08-26 11:07:22 -07:00
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2005-11-07 17:19:07 -05:00
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|
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|
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|
|
|
2017-07-17 08:45:35 +01:00
|
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|
2005-11-07 17:19:07 -05:00
|
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|
|
2011-11-24 18:57:30 -05:00
|
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2011-11-24 19:54:23 -05:00
|
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|
2017-07-17 08:45:35 +01:00
|
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|
|
2010-08-26 11:07:22 -07:00
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2008-03-27 13:06:23 +01:00
|
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2005-11-07 17:19:07 -05:00
|
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2008-08-02 00:51:11 -04:00
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2005-11-07 17:19:07 -05:00
|
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|
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2017-07-17 08:45:35 +01:00
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|
|
2010-08-26 11:07:22 -07:00
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2013-03-16 15:12:40 -04:00
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2008-03-27 13:06:23 +01:00
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2013-09-29 11:24:49 -04:00
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|
2011-11-25 03:06:56 -05:00
|
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|
|
2011-11-24 20:43:10 -05:00
|
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|
2013-09-29 11:24:49 -04:00
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2008-03-27 13:06:23 +01:00
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2013-03-16 15:12:40 -04:00
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2008-03-27 13:06:23 +01:00
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2005-11-07 17:19:07 -05:00
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2013-03-29 21:04:39 -07:00
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2005-04-16 15:20:36 -07:00
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2012-10-11 11:42:01 -04:00
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2008-02-08 04:22:12 -08:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2013-03-15 10:53:28 -04:00
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2013-02-22 22:49:10 -05:00
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2005-04-16 15:20:36 -07:00
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2011-09-26 20:36:09 -04:00
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2005-04-16 15:20:36 -07:00
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2010-03-07 18:49:36 -08:00
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2013-03-30 01:35:18 -07:00
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2017-07-04 17:25:09 +01:00
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2010-03-07 18:49:36 -08:00
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2013-03-15 10:53:28 -04:00
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2011-03-18 08:55:38 -04:00
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2011-11-24 21:24:27 -05:00
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2013-03-15 10:53:28 -04:00
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2011-11-24 21:24:27 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 01:05:37 -05:00
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2011-03-18 08:55:38 -04:00
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2005-11-07 17:21:20 -05:00
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take the targets of /proc/*/ns/* symlinks to separate fs
New pseudo-filesystem: nsfs. Targets of /proc/*/ns/* live there now.
It's not mountable (not even registered, so it's not in /proc/filesystems,
etc.). Files on it *are* bindable - we explicitly permit that in do_loopback().
This stuff lives in fs/nsfs.c now; proc_ns_fget() moved there as well.
get_proc_ns() is a macro now (it's simply returning ->i_private; would
have been an inline, if not for header ordering headache).
proc_ns_inode() is an ex-parrot. The interface used in procfs is
ns_get_path(path, task, ops) and ns_get_name(buf, size, task, ops).
Dentries and inodes are never hashed; a non-counting reference to dentry
is stashed in ns_common (removed by ->d_prune()) and reused by ns_get_path()
if present. See ns_get_path()/ns_prune_dentry/nsfs_evict() for details
of that mechanism.
As the result, proc_ns_follow_link() has stopped poking in nd->path.mnt;
it does nd_jump_link() on a consistent <vfsmount,dentry> pair it gets
from ns_get_path().
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2014-11-01 10:57:28 -04:00
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2011-03-18 08:55:38 -04:00
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2005-04-16 15:20:36 -07:00
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2013-03-29 21:04:39 -07:00
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2005-11-07 17:15:04 -05:00
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2013-03-30 01:35:18 -07:00
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2005-11-07 17:15:04 -05:00
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2011-11-24 21:24:27 -05:00
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2005-11-07 17:15:04 -05:00
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2012-06-25 12:55:18 +01:00
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2013-04-09 17:33:29 +04:00
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2012-06-25 12:55:18 +01:00
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2005-11-07 17:15:04 -05:00
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2013-03-29 21:04:39 -07:00
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2013-03-15 10:53:28 -04:00
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2005-11-07 17:15:04 -05:00
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2013-09-29 11:24:49 -04:00
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2014-12-24 07:20:01 -06:00
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2013-09-29 11:24:49 -04:00
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2005-11-07 17:16:29 -05:00
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2011-03-18 08:55:38 -04:00
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2013-03-15 10:53:28 -04:00
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2005-11-07 17:15:04 -05:00
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2008-08-02 00:51:11 -04:00
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2005-04-16 15:20:36 -07:00
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2008-02-15 14:38:00 -08:00
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2011-11-24 22:37:54 -05:00
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2008-02-15 14:38:00 -08:00
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2011-11-24 22:37:54 -05:00
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2008-02-15 14:38:00 -08:00
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2005-04-16 15:20:36 -07:00
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2017-07-17 08:45:35 +01:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2011-11-25 00:46:35 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 00:46:35 -05:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2005-04-16 15:20:36 -07:00
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2014-07-28 17:10:56 -07:00
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mnt: Correct permission checks in do_remount
While invesgiating the issue where in "mount --bind -oremount,ro ..."
would result in later "mount --bind -oremount,rw" succeeding even if
the mount started off locked I realized that there are several
additional mount flags that should be locked and are not.
In particular MNT_NOSUID, MNT_NODEV, MNT_NOEXEC, and the atime
flags in addition to MNT_READONLY should all be locked. These
flags are all per superblock, can all be changed with MS_BIND,
and should not be changable if set by a more privileged user.
The following additions to the current logic are added in this patch.
- nosuid may not be clearable by a less privileged user.
- nodev may not be clearable by a less privielged user.
- noexec may not be clearable by a less privileged user.
- atime flags may not be changeable by a less privileged user.
The logic with atime is that always setting atime on access is a
global policy and backup software and auditing software could break if
atime bits are not updated (when they are configured to be updated),
and serious performance degradation could result (DOS attack) if atime
updates happen when they have been explicitly disabled. Therefore an
unprivileged user should not be able to mess with the atime bits set
by a more privileged user.
The additional restrictions are implemented with the addition of
MNT_LOCK_NOSUID, MNT_LOCK_NODEV, MNT_LOCK_NOEXEC, and MNT_LOCK_ATIME
mnt flags.
Taken together these changes and the fixes for MNT_LOCK_READONLY
should make it safe for an unprivileged user to create a user
namespace and to call "mount --bind -o remount,... ..." without
the danger of mount flags being changed maliciously.
Cc: stable@vger.kernel.org
Acked-by: Serge E. Hallyn <serge.hallyn@ubuntu.com>
Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2014-07-28 17:26:07 -07:00
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2016-05-18 13:50:06 -05:00
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mnt: Correct permission checks in do_remount
While invesgiating the issue where in "mount --bind -oremount,ro ..."
would result in later "mount --bind -oremount,rw" succeeding even if
the mount started off locked I realized that there are several
additional mount flags that should be locked and are not.
In particular MNT_NOSUID, MNT_NODEV, MNT_NOEXEC, and the atime
flags in addition to MNT_READONLY should all be locked. These
flags are all per superblock, can all be changed with MS_BIND,
and should not be changable if set by a more privileged user.
The following additions to the current logic are added in this patch.
- nosuid may not be clearable by a less privileged user.
- nodev may not be clearable by a less privielged user.
- noexec may not be clearable by a less privileged user.
- atime flags may not be changeable by a less privileged user.
The logic with atime is that always setting atime on access is a
global policy and backup software and auditing software could break if
atime bits are not updated (when they are configured to be updated),
and serious performance degradation could result (DOS attack) if atime
updates happen when they have been explicitly disabled. Therefore an
unprivileged user should not be able to mess with the atime bits set
by a more privileged user.
The additional restrictions are implemented with the addition of
MNT_LOCK_NOSUID, MNT_LOCK_NODEV, MNT_LOCK_NOEXEC, and MNT_LOCK_ATIME
mnt flags.
Taken together these changes and the fixes for MNT_LOCK_READONLY
should make it safe for an unprivileged user to create a user
namespace and to call "mount --bind -o remount,... ..." without
the danger of mount flags being changed maliciously.
Cc: stable@vger.kernel.org
Acked-by: Serge E. Hallyn <serge.hallyn@ubuntu.com>
Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2014-07-28 17:26:07 -07:00
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2011-03-03 16:09:14 -05:00
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2005-04-16 15:20:36 -07:00
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2017-07-17 08:45:35 +01:00
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2013-02-22 22:49:10 -05:00
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2009-05-08 13:36:58 -04:00
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2017-07-17 08:45:35 +01:00
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2010-01-16 13:01:26 -05:00
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2013-09-29 11:24:49 -04:00
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2014-07-28 16:26:53 -07:00
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2011-11-25 00:46:35 -05:00
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2013-09-29 11:24:49 -04:00
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2008-09-26 19:01:20 -07:00
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2013-09-16 22:41:01 -04:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 20:01:19 -05:00
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2005-11-07 17:21:20 -05:00
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2011-11-24 18:57:30 -05:00
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2011-11-25 03:06:56 -05:00
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2011-11-25 01:05:37 -05:00
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2005-11-07 17:21:20 -05:00
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2012-10-11 11:42:01 -04:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2011-11-24 21:47:05 -05:00
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2011-11-24 19:59:16 -05:00
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2013-03-15 10:53:28 -04:00
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2013-02-22 22:49:10 -05:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2005-04-16 15:20:36 -07:00
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2013-03-15 10:53:28 -04:00
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Add a dentry op to allow processes to be held during pathwalk transit
Add a dentry op (d_manage) to permit a filesystem to hold a process and make it
sleep when it tries to transit away from one of that filesystem's directories
during a pathwalk. The operation is keyed off a new dentry flag
(DCACHE_MANAGE_TRANSIT).
The filesystem is allowed to be selective about which processes it holds and
which it permits to continue on or prohibits from transiting from each flagged
directory. This will allow autofs to hold up client processes whilst letting
its userspace daemon through to maintain the directory or the stuff behind it
or mounted upon it.
The ->d_manage() dentry operation:
int (*d_manage)(struct path *path, bool mounting_here);
takes a pointer to the directory about to be transited away from and a flag
indicating whether the transit is undertaken by do_add_mount() or
do_move_mount() skipping through a pile of filesystems mounted on a mountpoint.
It should return 0 if successful and to let the process continue on its way;
-EISDIR to prohibit the caller from skipping to overmounted filesystems or
automounting, and to use this directory; or some other error code to return to
the user.
->d_manage() is called with namespace_sem writelocked if mounting_here is true
and no other locks held, so it may sleep. However, if mounting_here is true,
it may not initiate or wait for a mount or unmount upon the parameter
directory, even if the act is actually performed by userspace.
Within fs/namei.c, follow_managed() is extended to check with d_manage() first
on each managed directory, before transiting away from it or attempting to
automount upon it.
follow_down() is renamed follow_down_one() and should only be used where the
filesystem deliberately intends to avoid management steps (e.g. autofs).
A new follow_down() is added that incorporates the loop done by all other
callers of follow_down() (do_add/move_mount(), autofs and NFSD; whilst AFS, NFS
and CIFS do use it, their use is removed by converting them to use
d_automount()). The new follow_down() calls d_manage() as appropriate. It
also takes an extra parameter to indicate if it is being called from mount code
(with namespace_sem writelocked) which it passes to d_manage(). follow_down()
ignores automount points so that it can be used to mount on them.
__follow_mount_rcu() is made to abort rcu-walk mode if it hits a directory with
DCACHE_MANAGE_TRANSIT set on the basis that we're probably going to have to
sleep. It would be possible to enter d_manage() in rcu-walk mode too, and have
that determine whether to abort or not itself. That would allow the autofs
daemon to continue on in rcu-walk mode.
Note that DCACHE_MANAGE_TRANSIT on a directory should be cleared when it isn't
required as every tranist from that directory will cause d_manage() to be
invoked. It can always be set again when necessary.
==========================
WHAT THIS MEANS FOR AUTOFS
==========================
Autofs currently uses the lookup() inode op and the d_revalidate() dentry op to
trigger the automounting of indirect mounts, and both of these can be called
with i_mutex held.
autofs knows that the i_mutex will be held by the caller in lookup(), and so
can drop it before invoking the daemon - but this isn't so for d_revalidate(),
since the lock is only held on _some_ of the code paths that call it. This
means that autofs can't risk dropping i_mutex from its d_revalidate() function
before it calls the daemon.
The bug could manifest itself as, for example, a process that's trying to
validate an automount dentry that gets made to wait because that dentry is
expired and needs cleaning up:
mkdir S ffffffff8014e05a 0 32580 24956
Call Trace:
[<ffffffff885371fd>] :autofs4:autofs4_wait+0x674/0x897
[<ffffffff80127f7d>] avc_has_perm+0x46/0x58
[<ffffffff8009fdcf>] autoremove_wake_function+0x0/0x2e
[<ffffffff88537be6>] :autofs4:autofs4_expire_wait+0x41/0x6b
[<ffffffff88535cfc>] :autofs4:autofs4_revalidate+0x91/0x149
[<ffffffff80036d96>] __lookup_hash+0xa0/0x12f
[<ffffffff80057a2f>] lookup_create+0x46/0x80
[<ffffffff800e6e31>] sys_mkdirat+0x56/0xe4
versus the automount daemon which wants to remove that dentry, but can't
because the normal process is holding the i_mutex lock:
automount D ffffffff8014e05a 0 32581 1 32561
Call Trace:
[<ffffffff80063c3f>] __mutex_lock_slowpath+0x60/0x9b
[<ffffffff8000ccf1>] do_path_lookup+0x2ca/0x2f1
[<ffffffff80063c89>] .text.lock.mutex+0xf/0x14
[<ffffffff800e6d55>] do_rmdir+0x77/0xde
[<ffffffff8005d229>] tracesys+0x71/0xe0
[<ffffffff8005d28d>] tracesys+0xd5/0xe0
which means that the system is deadlocked.
This patch allows autofs to hold up normal processes whilst the daemon goes
ahead and does things to the dentry tree behind the automouter point without
risking a deadlock as almost no locks are held in d_manage() and none in
d_automount().
Signed-off-by: David Howells <dhowells@redhat.com>
Was-Acked-by: Ian Kent <raven@themaw.net>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2011-01-14 18:45:26 +00:00
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2011-11-25 00:46:35 -05:00
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2011-11-25 01:05:37 -05:00
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2011-11-25 00:46:35 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 01:05:37 -05:00
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2005-04-16 15:20:36 -07:00
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2013-03-29 21:04:39 -07:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2005-11-07 17:20:03 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 21:47:05 -05:00
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2005-11-07 17:20:03 -05:00
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2005-04-16 15:20:36 -07:00
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VFS: (Scripted) Convert S_ISLNK/DIR/REG(dentry->d_inode) to d_is_*(dentry)
Convert the following where appropriate:
(1) S_ISLNK(dentry->d_inode) to d_is_symlink(dentry).
(2) S_ISREG(dentry->d_inode) to d_is_reg(dentry).
(3) S_ISDIR(dentry->d_inode) to d_is_dir(dentry). This is actually more
complicated than it appears as some calls should be converted to
d_can_lookup() instead. The difference is whether the directory in
question is a real dir with a ->lookup op or whether it's a fake dir with
a ->d_automount op.
In some circumstances, we can subsume checks for dentry->d_inode not being
NULL into this, provided we the code isn't in a filesystem that expects
d_inode to be NULL if the dirent really *is* negative (ie. if we're going to
use d_inode() rather than d_backing_inode() to get the inode pointer).
Note that the dentry type field may be set to something other than
DCACHE_MISS_TYPE when d_inode is NULL in the case of unionmount, where the VFS
manages the fall-through from a negative dentry to a lower layer. In such a
case, the dentry type of the negative union dentry is set to the same as the
type of the lower dentry.
However, if you know d_inode is not NULL at the call site, then you can use
the d_is_xxx() functions even in a filesystem.
There is one further complication: a 0,0 chardev dentry may be labelled
DCACHE_WHITEOUT_TYPE rather than DCACHE_SPECIAL_TYPE. Strictly, this was
intended for special directory entry types that don't have attached inodes.
The following perl+coccinelle script was used:
use strict;
my @callers;
open($fd, 'git grep -l \'S_IS[A-Z].*->d_inode\' |') ||
die "Can't grep for S_ISDIR and co. callers";
@callers = <$fd>;
close($fd);
unless (@callers) {
print "No matches\n";
exit(0);
}
my @cocci = (
'@@',
'expression E;',
'@@',
'',
'- S_ISLNK(E->d_inode->i_mode)',
'+ d_is_symlink(E)',
'',
'@@',
'expression E;',
'@@',
'',
'- S_ISDIR(E->d_inode->i_mode)',
'+ d_is_dir(E)',
'',
'@@',
'expression E;',
'@@',
'',
'- S_ISREG(E->d_inode->i_mode)',
'+ d_is_reg(E)' );
my $coccifile = "tmp.sp.cocci";
open($fd, ">$coccifile") || die $coccifile;
print($fd "$_\n") || die $coccifile foreach (@cocci);
close($fd);
foreach my $file (@callers) {
chomp $file;
print "Processing ", $file, "\n";
system("spatch", "--sp-file", $coccifile, $file, "--in-place", "--no-show-diff") == 0 ||
die "spatch failed";
}
[AV: overlayfs parts skipped]
Signed-off-by: David Howells <dhowells@redhat.com>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2015-01-29 12:02:35 +00:00
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2005-11-07 17:20:03 -05:00
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2011-11-25 01:05:37 -05:00
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2005-11-07 17:20:03 -05:00
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2005-11-07 17:21:20 -05:00
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2011-11-25 01:05:37 -05:00
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2005-11-07 17:21:20 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 01:05:37 -05:00
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2011-11-24 21:47:05 -05:00
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2005-11-07 17:20:03 -05:00
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2005-04-16 15:20:36 -07:00
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2013-03-15 10:53:28 -04:00
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2008-02-14 19:34:32 -08:00
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2005-11-07 17:20:03 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 00:22:05 -05:00
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2005-04-16 15:20:36 -07:00
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2013-03-15 10:53:28 -04:00
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2005-04-16 15:20:36 -07:00
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2008-03-21 20:48:19 -04:00
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2008-08-02 00:51:11 -04:00
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2005-04-16 15:20:36 -07:00
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2011-03-17 22:08:28 -04:00
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2011-11-25 00:30:56 -05:00
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2011-03-17 22:08:28 -04:00
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2013-03-15 10:53:28 -04:00
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2011-03-17 22:08:28 -04:00
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smarter propagate_mnt()
The current mainline has copies propagated to *all* nodes, then
tears down the copies we made for nodes that do not contain
counterparts of the desired mountpoint. That sets the right
propagation graph for the copies (at teardown time we move
the slaves of removed node to a surviving peer or directly
to master), but we end up paying a fairly steep price in
useless allocations. It's fairly easy to create a situation
where N calls of mount(2) create exactly N bindings, with
O(N^2) vfsmounts allocated and freed in process.
Fortunately, it is possible to avoid those allocations/freeings.
The trick is to create copies in the right order and find which
one would've eventually become a master with the current algorithm.
It turns out to be possible in O(nodes getting propagation) time
and with no extra allocations at all.
One part is that we need to make sure that eventual master will be
created before its slaves, so we need to walk the propagation
tree in a different order - by peer groups. And iterate through
the peers before dealing with the next group.
Another thing is finding the (earlier) copy that will be a master
of one we are about to create; to do that we are (temporary) marking
the masters of mountpoints we are attaching the copies to.
Either we are in a peer of the last mountpoint we'd dealt with,
or we have the following situation: we are attaching to mountpoint M,
the last copy S_0 had been attached to M_0 and there are sequences
S_0...S_n, M_0...M_n such that S_{i+1} is a master of S_{i},
S_{i} mounted on M{i} and we need to create a slave of the first S_{k}
such that M is getting propagation from M_{k}. It means that the master
of M_{k} will be among the sequence of masters of M. On the
other hand, the nearest marked node in that sequence will either
be the master of M_{k} or the master of M_{k-1} (the latter -
in the case if M_{k-1} is a slave of something M gets propagation
from, but in a wrong peer group).
So we go through the sequence of masters of M until we find
a marked one (P). Let N be the one before it. Then we go through
the sequence of masters of S_0 until we find one (say, S) mounted
on a node D that has P as master and check if D is a peer of N.
If it is, S will be the master of new copy, if not - the master of S
will be.
That's it for the hard part; the rest is fairly simple. Iterator
is in next_group(), handling of one prospective mountpoint is
propagate_one().
It seems to survive all tests and gives a noticably better performance
than the current mainline for setups that are seriously using shared
subtrees.
Cc: stable@vger.kernel.org
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2014-02-27 09:35:45 -05:00
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2011-03-17 22:08:28 -04:00
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2013-03-15 10:53:28 -04:00
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2011-03-17 22:08:28 -04:00
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2013-03-15 10:53:28 -04:00
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2011-03-17 22:08:28 -04:00
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2013-03-15 10:53:28 -04:00
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2012-09-21 08:19:02 -04:00
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2013-03-15 10:53:28 -04:00
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2012-09-21 08:19:02 -04:00
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2011-03-17 22:08:28 -04:00
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2011-11-25 00:30:56 -05:00
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2011-03-17 22:08:28 -04:00
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VFS: (Scripted) Convert S_ISLNK/DIR/REG(dentry->d_inode) to d_is_*(dentry)
Convert the following where appropriate:
(1) S_ISLNK(dentry->d_inode) to d_is_symlink(dentry).
(2) S_ISREG(dentry->d_inode) to d_is_reg(dentry).
(3) S_ISDIR(dentry->d_inode) to d_is_dir(dentry). This is actually more
complicated than it appears as some calls should be converted to
d_can_lookup() instead. The difference is whether the directory in
question is a real dir with a ->lookup op or whether it's a fake dir with
a ->d_automount op.
In some circumstances, we can subsume checks for dentry->d_inode not being
NULL into this, provided we the code isn't in a filesystem that expects
d_inode to be NULL if the dirent really *is* negative (ie. if we're going to
use d_inode() rather than d_backing_inode() to get the inode pointer).
Note that the dentry type field may be set to something other than
DCACHE_MISS_TYPE when d_inode is NULL in the case of unionmount, where the VFS
manages the fall-through from a negative dentry to a lower layer. In such a
case, the dentry type of the negative union dentry is set to the same as the
type of the lower dentry.
However, if you know d_inode is not NULL at the call site, then you can use
the d_is_xxx() functions even in a filesystem.
There is one further complication: a 0,0 chardev dentry may be labelled
DCACHE_WHITEOUT_TYPE rather than DCACHE_SPECIAL_TYPE. Strictly, this was
intended for special directory entry types that don't have attached inodes.
The following perl+coccinelle script was used:
use strict;
my @callers;
open($fd, 'git grep -l \'S_IS[A-Z].*->d_inode\' |') ||
die "Can't grep for S_ISDIR and co. callers";
@callers = <$fd>;
close($fd);
unless (@callers) {
print "No matches\n";
exit(0);
}
my @cocci = (
'@@',
'expression E;',
'@@',
'',
'- S_ISLNK(E->d_inode->i_mode)',
'+ d_is_symlink(E)',
'',
'@@',
'expression E;',
'@@',
'',
'- S_ISDIR(E->d_inode->i_mode)',
'+ d_is_dir(E)',
'',
'@@',
'expression E;',
'@@',
'',
'- S_ISREG(E->d_inode->i_mode)',
'+ d_is_reg(E)' );
my $coccifile = "tmp.sp.cocci";
open($fd, ">$coccifile") || die $coccifile;
print($fd "$_\n") || die $coccifile foreach (@cocci);
close($fd);
foreach my $file (@callers) {
chomp $file;
print "Processing ", $file, "\n";
system("spatch", "--sp-file", $coccifile, $file, "--in-place", "--no-show-diff") == 0 ||
die "spatch failed";
}
[AV: overlayfs parts skipped]
Signed-off-by: David Howells <dhowells@redhat.com>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2015-01-29 12:02:35 +00:00
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2011-03-17 22:08:28 -04:00
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2011-11-25 00:30:56 -05:00
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2013-03-15 10:53:28 -04:00
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2011-03-17 22:08:28 -04:00
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2013-03-15 10:53:28 -04:00
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2011-03-17 22:08:28 -04:00
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2011-01-17 01:47:59 -05:00
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2016-06-09 16:06:06 -05:00
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2015-05-08 23:22:29 -05:00
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2005-04-16 15:20:36 -07:00
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2017-07-17 08:45:35 +01:00
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2012-10-11 11:42:01 -04:00
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2005-04-16 15:20:36 -07:00
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2012-07-26 21:42:03 -07:00
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2005-04-16 15:20:36 -07:00
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2011-01-17 01:41:58 -05:00
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2005-04-16 15:20:36 -07:00
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2012-07-26 21:42:03 -07:00
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2005-04-16 15:20:36 -07:00
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2012-07-26 21:42:03 -07:00
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2017-07-17 08:45:35 +01:00
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2012-07-26 21:42:03 -07:00
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2005-04-16 15:20:36 -07:00
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2016-06-09 16:06:06 -05:00
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2011-11-25 00:30:56 -05:00
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2011-01-17 01:41:58 -05:00
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2005-04-16 15:20:36 -07:00
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2011-01-17 01:35:23 -05:00
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2011-11-25 00:22:05 -05:00
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2011-01-17 01:35:23 -05:00
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2011-11-25 00:22:05 -05:00
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2011-01-17 01:35:23 -05:00
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2011-01-17 01:47:59 -05:00
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2011-01-17 01:35:23 -05:00
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2011-11-25 00:30:56 -05:00
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2011-01-17 01:47:59 -05:00
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2011-11-25 00:22:05 -05:00
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2013-03-16 15:12:40 -04:00
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2011-11-25 00:22:05 -05:00
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2013-03-16 15:12:40 -04:00
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2011-01-17 01:35:23 -05:00
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2011-01-17 01:47:59 -05:00
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2011-01-17 01:35:23 -05:00
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2011-01-14 19:10:03 +00:00
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2013-03-16 15:12:40 -04:00
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2011-01-14 19:10:03 +00:00
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2011-11-25 00:22:05 -05:00
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2011-01-14 19:10:03 +00:00
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2013-03-16 15:12:40 -04:00
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2011-01-14 19:10:03 +00:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 21:07:43 -05:00
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2005-04-16 15:20:36 -07:00
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2013-03-16 15:12:40 -04:00
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2013-09-29 11:24:49 -04:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 00:22:05 -05:00
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2011-11-25 00:57:42 -05:00
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2011-11-24 21:35:16 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 00:22:05 -05:00
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2005-04-16 15:20:36 -07:00
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2008-03-22 00:21:53 -04:00
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2011-11-25 00:22:05 -05:00
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2011-11-25 00:46:35 -05:00
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2014-12-24 07:20:01 -06:00
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2008-03-22 00:21:53 -04:00
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2013-09-29 11:24:49 -04:00
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2013-03-16 14:42:19 -04:00
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2006-06-09 09:34:17 -04:00
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2011-11-24 21:15:14 -05:00
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2006-06-09 09:34:17 -04:00
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2011-11-24 21:15:14 -05:00
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2006-06-09 09:34:17 -04:00
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2011-11-24 23:24:33 -05:00
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2006-06-09 09:34:17 -04:00
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2011-11-24 23:24:33 -05:00
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2006-06-09 09:34:17 -04:00
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2011-11-24 23:24:33 -05:00
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2006-06-09 09:34:17 -04:00
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2011-11-24 21:15:14 -05:00
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2005-04-16 15:20:36 -07:00
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2006-06-09 09:34:17 -04:00
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2011-11-24 23:24:33 -05:00
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2006-06-09 09:34:17 -04:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 21:35:16 -05:00
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2011-11-25 00:22:05 -05:00
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2006-06-09 09:34:17 -04:00
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2005-04-16 15:20:36 -07:00
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2006-06-09 09:34:17 -04:00
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2011-11-24 23:24:33 -05:00
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2011-11-24 22:19:58 -05:00
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2006-06-09 09:34:17 -04:00
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fs: brlock vfsmount_lock
fs: brlock vfsmount_lock
Use a brlock for the vfsmount lock. It must be taken for write whenever
modifying the mount hash or associated fields, and may be taken for read when
performing mount hash lookups.
A new lock is added for the mnt-id allocator, so it doesn't need to take
the heavy vfsmount write-lock.
The number of atomics should remain the same for fastpath rlock cases, though
code would be slightly slower due to per-cpu access. Scalability is not not be
much improved in common cases yet, due to other locks (ie. dcache_lock) getting
in the way. However path lookups crossing mountpoints should be one case where
scalability is improved (currently requiring the global lock).
The slowpath is slower due to use of brlock. On a 64 core, 64 socket, 32 node
Altix system (high latency to remote nodes), a simple umount microbenchmark
(mount --bind mnt mnt2 ; umount mnt2 loop 1000 times), before this patch it
took 6.8s, afterwards took 7.1s, about 5% slower.
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
2010-08-18 04:37:39 +10:00
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2013-09-29 22:06:07 -04:00
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2006-06-09 09:34:17 -04:00
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2013-03-16 14:39:34 -04:00
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2006-06-09 09:34:17 -04:00
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2011-11-24 21:07:43 -05:00
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2006-06-09 09:34:17 -04:00
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2008-03-22 00:46:23 -04:00
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2008-03-22 00:21:53 -04:00
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2011-11-24 21:07:43 -05:00
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2011-11-25 00:22:05 -05:00
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2011-11-25 00:46:35 -05:00
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2014-12-24 07:20:01 -06:00
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2008-03-22 00:21:53 -04:00
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2005-04-16 15:20:36 -07:00
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2005-11-07 17:16:09 -05:00
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2005-04-16 15:20:36 -07:00
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2015-12-14 18:44:44 -05:00
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2005-04-16 15:20:36 -07:00
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2015-12-14 18:44:44 -05:00
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2005-11-07 17:16:09 -05:00
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2005-04-16 15:20:36 -07:00
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2015-12-14 18:44:44 -05:00
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2005-04-16 15:20:36 -07:00
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2015-12-14 18:44:44 -05:00
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2005-04-16 15:20:36 -07:00
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2015-12-14 18:44:44 -05:00
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2005-04-16 15:20:36 -07:00
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2015-12-14 18:44:44 -05:00
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2005-04-16 15:20:36 -07:00
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2015-12-14 18:44:44 -05:00
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2005-04-16 15:20:36 -07:00
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2014-08-28 11:26:03 -06:00
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fs: fix overflow in sys_mount() for in-kernel calls
sys_mount() reads/copies a whole page for its "type" parameter. When
do_mount_root() passes a kernel address that points to an object which is
smaller than a whole page, copy_mount_options() will happily go past this
memory object, possibly dereferencing "wild" pointers that could be in any
state (hence the kmemcheck warning, which shows that parts of the next
page are not even allocated).
(The likelihood of something going wrong here is pretty low -- first of
all this only applies to kernel calls to sys_mount(), which are mostly
found in the boot code. Secondly, I guess if the page was not mapped,
exact_copy_from_user() _would_ in fact handle it correctly because of its
access_ok(), etc. checks.)
But it is much nicer to avoid the dubious reads altogether, by stopping as
soon as we find a NUL byte. Is there a good reason why we can't do
something like this, using the already existing strndup_from_user()?
[akpm@linux-foundation.org: make copy_mount_string() static]
[AV: fix compat mount breakage, which involves undoing akpm's change above]
Reported-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Vegard Nossum <vegard.nossum@gmail.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: al <al@dizzy.pdmi.ras.ru>
2009-09-18 13:05:45 -07:00
|
|
|
|
2014-08-28 11:26:03 -06:00
|
|
|
|
fs: fix overflow in sys_mount() for in-kernel calls
sys_mount() reads/copies a whole page for its "type" parameter. When
do_mount_root() passes a kernel address that points to an object which is
smaller than a whole page, copy_mount_options() will happily go past this
memory object, possibly dereferencing "wild" pointers that could be in any
state (hence the kmemcheck warning, which shows that parts of the next
page are not even allocated).
(The likelihood of something going wrong here is pretty low -- first of
all this only applies to kernel calls to sys_mount(), which are mostly
found in the boot code. Secondly, I guess if the page was not mapped,
exact_copy_from_user() _would_ in fact handle it correctly because of its
access_ok(), etc. checks.)
But it is much nicer to avoid the dubious reads altogether, by stopping as
soon as we find a NUL byte. Is there a good reason why we can't do
something like this, using the already existing strndup_from_user()?
[akpm@linux-foundation.org: make copy_mount_string() static]
[AV: fix compat mount breakage, which involves undoing akpm's change above]
Reported-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Vegard Nossum <vegard.nossum@gmail.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: al <al@dizzy.pdmi.ras.ru>
2009-09-18 13:05:45 -07:00
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2005-04-16 15:20:36 -07:00
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2014-09-14 22:15:10 +09:00
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2012-10-11 11:42:01 -04:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2017-07-17 08:45:35 +01:00
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2005-04-16 15:20:36 -07:00
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2017-07-17 08:45:35 +01:00
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2009-10-04 21:49:49 +09:00
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2014-09-14 22:15:10 +09:00
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2009-10-04 21:49:49 +09:00
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2013-05-04 14:40:51 -04:00
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2017-07-17 08:45:35 +01:00
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2015-11-16 09:49:34 -05:00
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2009-10-04 21:49:49 +09:00
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2009-04-19 18:40:43 +02:00
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2009-03-26 17:53:14 +00:00
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2005-04-16 15:20:36 -07:00
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2006-01-09 20:52:17 -08:00
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2009-03-26 17:49:56 +00:00
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2017-07-17 08:45:35 +01:00
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2008-02-15 14:38:00 -08:00
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2006-01-09 20:52:17 -08:00
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2014-07-28 17:36:04 -07:00
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2017-07-17 08:45:35 +01:00
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2017-10-08 00:28:21 -04:00
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2017-10-11 07:01:31 +02:00
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2017-10-08 00:28:21 -04:00
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2005-04-16 15:20:36 -07:00
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2017-07-17 08:45:35 +01:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2005-11-07 17:21:20 -05:00
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2008-08-02 00:51:11 -04:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2005-04-16 15:20:36 -07:00
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2017-07-17 08:45:35 +01:00
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2005-04-16 15:20:36 -07:00
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2008-08-02 00:51:11 -04:00
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2005-04-16 15:20:36 -07:00
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2016-08-08 14:37:37 -05:00
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2012-07-26 21:08:32 -07:00
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2014-11-01 00:45:45 -04:00
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2016-08-08 14:37:37 -05:00
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2012-07-26 21:08:32 -07:00
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2010-03-07 18:49:36 -08:00
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2012-07-26 21:08:32 -07:00
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2009-06-22 15:09:13 -04:00
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2016-08-08 14:37:37 -05:00
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2011-06-15 10:21:48 -07:00
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2009-06-22 15:09:13 -04:00
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2016-08-08 14:37:37 -05:00
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2016-09-22 13:08:36 -05:00
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2016-08-08 14:37:37 -05:00
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2009-06-22 15:09:13 -04:00
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2016-08-08 14:37:37 -05:00
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2009-06-22 15:09:13 -04:00
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2016-08-08 14:37:37 -05:00
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2014-11-01 00:45:45 -04:00
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2011-06-15 10:21:48 -07:00
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2016-08-08 14:37:37 -05:00
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2011-06-15 10:21:48 -07:00
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2014-11-01 02:32:53 -04:00
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2010-03-07 18:49:36 -08:00
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2009-06-22 15:09:13 -04:00
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2012-07-26 21:08:32 -07:00
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2016-08-08 14:37:37 -05:00
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2016-09-28 00:27:17 -05:00
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2009-06-22 15:09:13 -04:00
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2016-06-20 20:42:34 +02:00
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2013-09-28 20:47:57 -04:00
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2005-04-16 15:20:36 -07:00
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2006-12-08 02:37:56 -08:00
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2008-05-10 20:44:54 -04:00
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2011-11-24 18:57:30 -05:00
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2013-09-28 20:47:57 -04:00
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2011-11-24 20:55:08 -05:00
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2012-07-31 13:13:04 -07:00
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2005-04-16 15:20:36 -07:00
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2013-09-28 20:47:57 -04:00
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2012-07-26 21:08:32 -07:00
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2009-06-22 15:09:13 -04:00
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2005-04-16 15:20:36 -07:00
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2013-03-16 15:12:40 -04:00
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2005-04-16 15:20:36 -07:00
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2013-03-30 01:35:18 -07:00
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2013-09-28 20:47:57 -04:00
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2013-03-22 04:08:05 -07:00
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2012-07-31 13:13:04 -07:00
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2012-06-25 12:55:18 +01:00
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2013-03-16 14:49:45 -04:00
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2012-07-26 21:08:32 -07:00
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2012-06-25 12:55:18 +01:00
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2005-04-16 15:20:36 -07:00
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2011-12-06 13:32:36 -05:00
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2011-11-25 02:19:55 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 03:06:56 -05:00
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2011-11-24 20:55:08 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 00:46:35 -05:00
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2016-09-28 00:27:17 -05:00
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2013-09-28 20:47:57 -04:00
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2011-11-24 18:57:30 -05:00
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2005-04-16 15:20:36 -07:00
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2013-09-28 20:47:57 -04:00
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2011-11-24 18:57:30 -05:00
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2005-04-16 15:20:36 -07:00
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2011-11-25 03:06:56 -05:00
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2013-03-30 01:35:18 -07:00
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2005-04-16 15:20:36 -07:00
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2013-03-16 14:49:45 -04:00
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2005-04-16 15:20:36 -07:00
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2011-01-14 22:30:21 -05:00
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2005-04-16 15:20:36 -07:00
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2011-01-14 22:30:21 -05:00
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2005-04-16 15:20:36 -07:00
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2006-02-07 12:59:00 -08:00
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2005-04-16 15:20:36 -07:00
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2009-06-22 15:09:13 -04:00
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2011-11-25 02:19:55 -05:00
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2009-06-22 15:09:13 -04:00
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2012-07-26 21:08:32 -07:00
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2009-06-22 15:09:13 -04:00
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2011-11-25 02:19:55 -05:00
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2011-12-06 13:32:36 -05:00
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2016-09-28 00:27:17 -05:00
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2013-05-04 15:18:53 -04:00
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2011-11-16 16:12:14 -05:00
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2011-11-25 02:19:55 -05:00
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2009-06-22 15:09:13 -04:00
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2011-11-16 21:43:59 -05:00
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2011-11-22 12:31:21 -05:00
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2011-11-16 21:43:59 -05:00
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2011-11-22 12:31:21 -05:00
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2011-11-16 21:43:59 -05:00
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2011-11-22 12:31:21 -05:00
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2011-11-16 21:43:59 -05:00
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2009-01-14 14:14:12 +01:00
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2005-04-16 15:20:36 -07:00
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fs: fix overflow in sys_mount() for in-kernel calls
sys_mount() reads/copies a whole page for its "type" parameter. When
do_mount_root() passes a kernel address that points to an object which is
smaller than a whole page, copy_mount_options() will happily go past this
memory object, possibly dereferencing "wild" pointers that could be in any
state (hence the kmemcheck warning, which shows that parts of the next
page are not even allocated).
(The likelihood of something going wrong here is pretty low -- first of
all this only applies to kernel calls to sys_mount(), which are mostly
found in the boot code. Secondly, I guess if the page was not mapped,
exact_copy_from_user() _would_ in fact handle it correctly because of its
access_ok(), etc. checks.)
But it is much nicer to avoid the dubious reads altogether, by stopping as
soon as we find a NUL byte. Is there a good reason why we can't do
something like this, using the already existing strndup_from_user()?
[akpm@linux-foundation.org: make copy_mount_string() static]
[AV: fix compat mount breakage, which involves undoing akpm's change above]
Reported-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Vegard Nossum <vegard.nossum@gmail.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: al <al@dizzy.pdmi.ras.ru>
2009-09-18 13:05:45 -07:00
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2015-12-14 18:44:44 -05:00
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2005-04-16 15:20:36 -07:00
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2014-08-28 11:26:03 -06:00
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|
fs: fix overflow in sys_mount() for in-kernel calls
sys_mount() reads/copies a whole page for its "type" parameter. When
do_mount_root() passes a kernel address that points to an object which is
smaller than a whole page, copy_mount_options() will happily go past this
memory object, possibly dereferencing "wild" pointers that could be in any
state (hence the kmemcheck warning, which shows that parts of the next
page are not even allocated).
(The likelihood of something going wrong here is pretty low -- first of
all this only applies to kernel calls to sys_mount(), which are mostly
found in the boot code. Secondly, I guess if the page was not mapped,
exact_copy_from_user() _would_ in fact handle it correctly because of its
access_ok(), etc. checks.)
But it is much nicer to avoid the dubious reads altogether, by stopping as
soon as we find a NUL byte. Is there a good reason why we can't do
something like this, using the already existing strndup_from_user()?
[akpm@linux-foundation.org: make copy_mount_string() static]
[AV: fix compat mount breakage, which involves undoing akpm's change above]
Reported-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Vegard Nossum <vegard.nossum@gmail.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: al <al@dizzy.pdmi.ras.ru>
2009-09-18 13:05:45 -07:00
|
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2005-04-16 15:20:36 -07:00
|
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|
|
2014-08-28 11:26:03 -06:00
|
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|
|
|
|
|
|
|
|
|
|
|
|
fs: fix overflow in sys_mount() for in-kernel calls
sys_mount() reads/copies a whole page for its "type" parameter. When
do_mount_root() passes a kernel address that points to an object which is
smaller than a whole page, copy_mount_options() will happily go past this
memory object, possibly dereferencing "wild" pointers that could be in any
state (hence the kmemcheck warning, which shows that parts of the next
page are not even allocated).
(The likelihood of something going wrong here is pretty low -- first of
all this only applies to kernel calls to sys_mount(), which are mostly
found in the boot code. Secondly, I guess if the page was not mapped,
exact_copy_from_user() _would_ in fact handle it correctly because of its
access_ok(), etc. checks.)
But it is much nicer to avoid the dubious reads altogether, by stopping as
soon as we find a NUL byte. Is there a good reason why we can't do
something like this, using the already existing strndup_from_user()?
[akpm@linux-foundation.org: make copy_mount_string() static]
[AV: fix compat mount breakage, which involves undoing akpm's change above]
Reported-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Vegard Nossum <vegard.nossum@gmail.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: al <al@dizzy.pdmi.ras.ru>
2009-09-18 13:05:45 -07:00
|
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2005-04-16 15:20:36 -07:00
|
|
|
|
2015-12-14 18:44:44 -05:00
|
|
|
|
|
|
|
|
|
|
|
|
|
|
fs: fix overflow in sys_mount() for in-kernel calls
sys_mount() reads/copies a whole page for its "type" parameter. When
do_mount_root() passes a kernel address that points to an object which is
smaller than a whole page, copy_mount_options() will happily go past this
memory object, possibly dereferencing "wild" pointers that could be in any
state (hence the kmemcheck warning, which shows that parts of the next
page are not even allocated).
(The likelihood of something going wrong here is pretty low -- first of
all this only applies to kernel calls to sys_mount(), which are mostly
found in the boot code. Secondly, I guess if the page was not mapped,
exact_copy_from_user() _would_ in fact handle it correctly because of its
access_ok(), etc. checks.)
But it is much nicer to avoid the dubious reads altogether, by stopping as
soon as we find a NUL byte. Is there a good reason why we can't do
something like this, using the already existing strndup_from_user()?
[akpm@linux-foundation.org: make copy_mount_string() static]
[AV: fix compat mount breakage, which involves undoing akpm's change above]
Reported-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Vegard Nossum <vegard.nossum@gmail.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: al <al@dizzy.pdmi.ras.ru>
2009-09-18 13:05:45 -07:00
|
|
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|
2005-04-16 15:20:36 -07:00
|
|
|
|
2015-12-14 18:44:44 -05:00
|
|
|
|
2005-04-16 15:20:36 -07:00
|
|
|
|
2015-12-14 18:44:44 -05:00
|
|
|
|
fs: fix overflow in sys_mount() for in-kernel calls
sys_mount() reads/copies a whole page for its "type" parameter. When
do_mount_root() passes a kernel address that points to an object which is
smaller than a whole page, copy_mount_options() will happily go past this
memory object, possibly dereferencing "wild" pointers that could be in any
state (hence the kmemcheck warning, which shows that parts of the next
page are not even allocated).
(The likelihood of something going wrong here is pretty low -- first of
all this only applies to kernel calls to sys_mount(), which are mostly
found in the boot code. Secondly, I guess if the page was not mapped,
exact_copy_from_user() _would_ in fact handle it correctly because of its
access_ok(), etc. checks.)
But it is much nicer to avoid the dubious reads altogether, by stopping as
soon as we find a NUL byte. Is there a good reason why we can't do
something like this, using the already existing strndup_from_user()?
[akpm@linux-foundation.org: make copy_mount_string() static]
[AV: fix compat mount breakage, which involves undoing akpm's change above]
Reported-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Vegard Nossum <vegard.nossum@gmail.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: al <al@dizzy.pdmi.ras.ru>
2009-09-18 13:05:45 -07:00
|
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2005-04-16 15:20:36 -07:00
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2011-11-23 19:34:49 -05:00
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2013-09-29 22:06:07 -04:00
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2011-11-23 19:34:49 -05:00
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2011-11-24 22:00:28 -05:00
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2011-11-23 19:34:49 -05:00
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2011-11-24 22:00:28 -05:00
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2011-11-24 22:25:07 -05:00
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2011-11-24 22:19:58 -05:00
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2011-11-23 19:34:49 -05:00
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2011-11-24 22:00:28 -05:00
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2011-11-23 19:34:49 -05:00
|
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|
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2016-11-14 22:14:35 +01:00
|
|
|
|
2011-11-23 19:34:49 -05:00
|
|
|
|
2015-11-17 14:40:10 +08:00
|
|
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2013-09-29 22:06:07 -04:00
|
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2011-11-24 22:00:28 -05:00
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2013-09-29 22:06:07 -04:00
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2011-11-23 19:34:49 -05:00
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2005-04-16 15:20:36 -07:00
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2006-01-08 01:03:18 -08:00
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2005-04-16 15:20:36 -07:00
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2009-01-14 14:14:16 +01:00
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2005-04-16 15:20:36 -07:00
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2008-07-22 09:59:21 -04:00
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2013-03-15 10:53:28 -04:00
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2005-04-16 15:20:36 -07:00
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2013-02-22 22:45:42 -05:00
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2005-04-16 15:20:36 -07:00
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2008-07-22 09:59:21 -04:00
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2005-04-16 15:20:36 -07:00
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2008-07-22 09:59:21 -04:00
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2005-04-16 15:20:36 -07:00
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2008-07-22 09:59:21 -04:00
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2011-03-18 08:55:38 -04:00
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2005-04-16 15:20:36 -07:00
|
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2010-08-10 11:41:36 +02:00
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2013-03-15 10:53:28 -04:00
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2011-03-18 08:55:38 -04:00
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2005-04-16 15:20:36 -07:00
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2011-11-24 19:41:16 -05:00
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2013-03-15 10:53:28 -04:00
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2011-11-25 01:05:37 -05:00
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2011-03-18 08:55:38 -04:00
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2011-11-25 00:46:35 -05:00
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2011-03-18 08:55:38 -04:00
|
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2013-03-29 21:04:39 -07:00
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2005-04-16 15:20:36 -07:00
|
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|
2009-05-04 03:32:03 +04:00
|
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|
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2011-03-18 08:55:38 -04:00
|
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2005-04-16 15:20:36 -07:00
|
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2013-03-15 10:53:28 -04:00
|
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2011-03-18 08:55:38 -04:00
|
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2005-04-16 15:20:36 -07:00
|
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2008-03-22 18:00:39 -04:00
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2011-03-18 08:55:38 -04:00
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2011-11-24 21:47:05 -05:00
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2011-03-18 08:55:38 -04:00
|
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2013-03-15 10:53:28 -04:00
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2008-07-22 09:59:21 -04:00
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2011-03-18 08:55:38 -04:00
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2011-11-24 21:47:05 -05:00
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2011-03-18 08:55:38 -04:00
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2008-02-14 19:34:32 -08:00
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2013-03-15 10:53:28 -04:00
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2011-03-18 08:55:38 -04:00
|
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|
|
2014-10-08 10:42:27 -07:00
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2013-03-15 10:53:28 -04:00
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2013-09-29 11:24:49 -04:00
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2011-11-24 19:41:16 -05:00
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2013-03-29 21:04:39 -07:00
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2008-02-14 19:34:32 -08:00
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2013-03-15 10:53:28 -04:00
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2008-02-14 19:34:32 -08:00
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2013-03-15 10:53:28 -04:00
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2006-12-08 02:37:56 -08:00
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2014-10-08 10:42:57 -07:00
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|
2017-01-03 14:18:43 +13:00
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2013-09-29 11:24:49 -04:00
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2008-07-22 09:59:21 -04:00
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2005-04-16 15:20:36 -07:00
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2011-03-18 08:55:38 -04:00
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2013-03-15 10:53:28 -04:00
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2011-03-18 08:55:38 -04:00
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2008-03-22 18:00:39 -04:00
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2011-03-18 08:55:38 -04:00
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2008-07-22 09:59:21 -04:00
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2005-04-16 15:20:36 -07:00
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2008-07-22 09:59:21 -04:00
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2005-04-16 15:20:36 -07:00
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2006-12-08 02:37:56 -08:00
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2008-02-14 19:34:39 -08:00
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2012-07-26 21:42:03 -07:00
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2005-04-16 15:20:36 -07:00
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2012-07-26 21:42:03 -07:00
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2005-04-16 15:20:36 -07:00
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|
fs: scale mntget/mntput
The problem that this patch aims to fix is vfsmount refcounting scalability.
We need to take a reference on the vfsmount for every successful path lookup,
which often go to the same mount point.
The fundamental difficulty is that a "simple" reference count can never be made
scalable, because any time a reference is dropped, we must check whether that
was the last reference. To do that requires communication with all other CPUs
that may have taken a reference count.
We can make refcounts more scalable in a couple of ways, involving keeping
distributed counters, and checking for the global-zero condition less
frequently.
- check the global sum once every interval (this will delay zero detection
for some interval, so it's probably a showstopper for vfsmounts).
- keep a local count and only taking the global sum when local reaches 0 (this
is difficult for vfsmounts, because we can't hold preempt off for the life of
a reference, so a counter would need to be per-thread or tied strongly to a
particular CPU which requires more locking).
- keep a local difference of increments and decrements, which allows us to sum
the total difference and hence find the refcount when summing all CPUs. Then,
keep a single integer "long" refcount for slow and long lasting references,
and only take the global sum of local counters when the long refcount is 0.
This last scheme is what I implemented here. Attached mounts and process root
and working directory references are "long" references, and everything else is
a short reference.
This allows scalable vfsmount references during path walking over mounted
subtrees and unattached (lazy umounted) mounts with processes still running
in them.
This results in one fewer atomic op in the fastpath: mntget is now just a
per-CPU inc, rather than an atomic inc; and mntput just requires a spinlock
and non-atomic decrement in the common case. However code is otherwise bigger
and heavier, so single threaded performance is basically a wash.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
2011-01-07 17:50:11 +11:00
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2009-06-23 17:29:49 -04:00
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2005-04-16 15:20:36 -07:00
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2006-12-08 02:37:56 -08:00
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2011-12-06 13:32:36 -05:00
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umount: Do not allow unmounting rootfs.
Andrew Vagin <avagin@parallels.com> writes:
> #define _GNU_SOURCE
> #include <sys/types.h>
> #include <sys/stat.h>
> #include <fcntl.h>
> #include <sched.h>
> #include <unistd.h>
> #include <sys/mount.h>
>
> int main(int argc, char **argv)
> {
> int fd;
>
> fd = open("/proc/self/ns/mnt", O_RDONLY);
> if (fd < 0)
> return 1;
> while (1) {
> if (umount2("/", MNT_DETACH) ||
> setns(fd, CLONE_NEWNS))
> break;
> }
>
> return 0;
> }
>
> root@ubuntu:/home/avagin# gcc -Wall nsenter.c -o nsenter
> root@ubuntu:/home/avagin# strace ./nsenter
> execve("./nsenter", ["./nsenter"], [/* 22 vars */]) = 0
> ...
> open("/proc/self/ns/mnt", O_RDONLY) = 3
> umount("/", MNT_DETACH) = 0
> setns(3, 131072) = 0
> umount("/", MNT_DETACH
>
causes:
> [ 260.548301] ------------[ cut here ]------------
> [ 260.550941] kernel BUG at /build/buildd/linux-3.13.0/fs/pnode.c:372!
> [ 260.552068] invalid opcode: 0000 [#1] SMP
> [ 260.552068] Modules linked in: xt_CHECKSUM iptable_mangle xt_tcpudp xt_addrtype xt_conntrack ipt_MASQUERADE iptable_nat nf_conntrack_ipv4 nf_defrag_ipv4 nf_nat_ipv4 nf_nat nf_conntrack bridge stp llc dm_thin_pool dm_persistent_data dm_bufio dm_bio_prison iptable_filter ip_tables x_tables crct10dif_pclmul crc32_pclmul ghash_clmulni_intel binfmt_misc nfsd auth_rpcgss nfs_acl aesni_intel nfs lockd aes_x86_64 sunrpc fscache lrw gf128mul glue_helper ablk_helper cryptd serio_raw ppdev parport_pc lp parport btrfs xor raid6_pq libcrc32c psmouse floppy
> [ 260.552068] CPU: 0 PID: 1723 Comm: nsenter Not tainted 3.13.0-30-generic #55-Ubuntu
> [ 260.552068] Hardware name: Bochs Bochs, BIOS Bochs 01/01/2011
> [ 260.552068] task: ffff8800376097f0 ti: ffff880074824000 task.ti: ffff880074824000
> [ 260.552068] RIP: 0010:[<ffffffff811e9483>] [<ffffffff811e9483>] propagate_umount+0x123/0x130
> [ 260.552068] RSP: 0018:ffff880074825e98 EFLAGS: 00010246
> [ 260.552068] RAX: ffff88007c741140 RBX: 0000000000000002 RCX: ffff88007c741190
> [ 260.552068] RDX: ffff88007c741190 RSI: ffff880074825ec0 RDI: ffff880074825ec0
> [ 260.552068] RBP: ffff880074825eb0 R08: 00000000000172e0 R09: ffff88007fc172e0
> [ 260.552068] R10: ffffffff811cc642 R11: ffffea0001d59000 R12: ffff88007c741140
> [ 260.552068] R13: ffff88007c741140 R14: ffff88007c741140 R15: 0000000000000000
> [ 260.552068] FS: 00007fd5c7e41740(0000) GS:ffff88007fc00000(0000) knlGS:0000000000000000
> [ 260.552068] CS: 0010 DS: 0000 ES: 0000 CR0: 0000000080050033
> [ 260.552068] CR2: 00007fd5c7968050 CR3: 0000000070124000 CR4: 00000000000406f0
> [ 260.552068] Stack:
> [ 260.552068] 0000000000000002 0000000000000002 ffff88007c631000 ffff880074825ed8
> [ 260.552068] ffffffff811dcfac ffff88007c741140 0000000000000002 ffff88007c741160
> [ 260.552068] ffff880074825f38 ffffffff811dd12b ffffffff811cc642 0000000075640000
> [ 260.552068] Call Trace:
> [ 260.552068] [<ffffffff811dcfac>] umount_tree+0x20c/0x260
> [ 260.552068] [<ffffffff811dd12b>] do_umount+0x12b/0x300
> [ 260.552068] [<ffffffff811cc642>] ? final_putname+0x22/0x50
> [ 260.552068] [<ffffffff811cc849>] ? putname+0x29/0x40
> [ 260.552068] [<ffffffff811dd88c>] SyS_umount+0xdc/0x100
> [ 260.552068] [<ffffffff8172aeff>] tracesys+0xe1/0xe6
> [ 260.552068] Code: 89 50 08 48 8b 50 08 48 89 02 49 89 45 08 e9 72 ff ff ff 0f 1f 44 00 00 4c 89 e6 4c 89 e7 e8 f5 f6 ff ff 48 89 c3 e9 39 ff ff ff <0f> 0b 66 2e 0f 1f 84 00 00 00 00 00 90 66 66 66 66 90 55 b8 01
> [ 260.552068] RIP [<ffffffff811e9483>] propagate_umount+0x123/0x130
> [ 260.552068] RSP <ffff880074825e98>
> [ 260.611451] ---[ end trace 11c33d85f1d4c652 ]--
Which in practice is totally uninteresting. Only the global root user can
do it, and it is just a stupid thing to do.
However that is no excuse to allow a silly way to oops the kernel.
We can avoid this silly problem by setting MNT_LOCKED on the rootfs
mount point and thus avoid needing any special cases in the unmount
code.
Signed-off-by: "Eric W. Biederman" <ebiederm@xmission.com>
2014-10-07 12:19:53 -07:00
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2008-02-14 19:34:39 -08:00
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2005-04-16 15:20:36 -07:00
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2007-10-16 23:26:30 -07:00
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2005-04-16 15:20:36 -07:00
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2006-09-29 01:58:57 -07:00
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2005-04-16 15:20:36 -07:00
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2011-11-23 12:14:10 -05:00
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2007-07-20 10:11:58 +09:00
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2005-04-16 15:20:36 -07:00
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2014-02-28 13:46:44 -05:00
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2014-03-20 21:10:51 -04:00
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2014-02-28 13:46:44 -05:00
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2017-07-06 15:39:11 -07:00
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2014-02-28 13:46:44 -05:00
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2017-07-06 15:39:11 -07:00
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2014-02-28 13:46:44 -05:00
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2005-04-16 15:20:36 -07:00
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2013-03-15 10:53:28 -04:00
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2005-04-16 15:20:36 -07:00
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2013-11-28 14:54:43 -05:00
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2006-09-29 01:58:57 -07:00
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2008-04-30 00:55:09 -07:00
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2007-10-29 14:17:23 -06:00
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2008-04-30 00:55:09 -07:00
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2005-04-16 15:20:36 -07:00
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2009-06-22 15:09:13 -04:00
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2005-04-16 15:20:36 -07:00
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2010-02-05 02:21:06 -05:00
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2009-06-22 15:09:13 -04:00
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2013-09-16 21:19:20 -04:00
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2012-07-26 21:08:32 -07:00
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2005-04-16 15:20:36 -07:00
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2011-03-17 22:08:28 -04:00
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2011-07-19 09:32:38 -07:00
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2017-07-17 08:45:35 +01:00
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2011-07-19 09:32:38 -07:00
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2012-06-09 00:59:08 -04:00
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2011-07-19 09:32:38 -07:00
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2011-03-17 22:08:28 -04:00
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2011-07-19 09:32:38 -07:00
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2012-06-09 00:59:08 -04:00
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2013-09-29 22:06:07 -04:00
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2011-07-19 09:32:38 -07:00
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fix apparmor dereferencing potentially freed dentry, sanitize __d_path() API
__d_path() API is asking for trouble and in case of apparmor d_namespace_path()
getting just that. The root cause is that when __d_path() misses the root
it had been told to look for, it stores the location of the most remote ancestor
in *root. Without grabbing references. Sure, at the moment of call it had
been pinned down by what we have in *path. And if we raced with umount -l, we
could have very well stopped at vfsmount/dentry that got freed as soon as
prepend_path() dropped vfsmount_lock.
It is safe to compare these pointers with pre-existing (and known to be still
alive) vfsmount and dentry, as long as all we are asking is "is it the same
address?". Dereferencing is not safe and apparmor ended up stepping into
that. d_namespace_path() really wants to examine the place where we stopped,
even if it's not connected to our namespace. As the result, it looked
at ->d_sb->s_magic of a dentry that might've been already freed by that point.
All other callers had been careful enough to avoid that, but it's really
a bad interface - it invites that kind of trouble.
The fix is fairly straightforward, even though it's bigger than I'd like:
* prepend_path() root argument becomes const.
* __d_path() is never called with NULL/NULL root. It was a kludge
to start with. Instead, we have an explicit function - d_absolute_root().
Same as __d_path(), except that it doesn't get root passed and stops where
it stops. apparmor and tomoyo are using it.
* __d_path() returns NULL on path outside of root. The main
caller is show_mountinfo() and that's precisely what we pass root for - to
skip those outside chroot jail. Those who don't want that can (and do)
use d_path().
* __d_path() root argument becomes const. Everyone agrees, I hope.
* apparmor does *NOT* try to use __d_path() or any of its variants
when it sees that path->mnt is an internal vfsmount. In that case it's
definitely not mounted anywhere and dentry_path() is exactly what we want
there. Handling of sysctl()-triggered weirdness is moved to that place.
* if apparmor is asked to do pathname relative to chroot jail
and __d_path() tells it we it's not in that jail, the sucker just calls
d_absolute_path() instead. That's the other remaining caller of __d_path(),
BTW.
* seq_path_root() does _NOT_ return -ENAMETOOLONG (it's stupid anyway -
the normal seq_file logics will take care of growing the buffer and redoing
the call of ->show() just fine). However, if it gets path not reachable
from root, it returns SEQ_SKIP. The only caller adjusted (i.e. stopped
ignoring the return value as it used to do).
Reviewed-by: John Johansen <john.johansen@canonical.com>
ACKed-by: John Johansen <john.johansen@canonical.com>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
Cc: stable@vger.kernel.org
2011-12-05 08:43:34 -05:00
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2011-11-25 00:46:35 -05:00
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fix apparmor dereferencing potentially freed dentry, sanitize __d_path() API
__d_path() API is asking for trouble and in case of apparmor d_namespace_path()
getting just that. The root cause is that when __d_path() misses the root
it had been told to look for, it stores the location of the most remote ancestor
in *root. Without grabbing references. Sure, at the moment of call it had
been pinned down by what we have in *path. And if we raced with umount -l, we
could have very well stopped at vfsmount/dentry that got freed as soon as
prepend_path() dropped vfsmount_lock.
It is safe to compare these pointers with pre-existing (and known to be still
alive) vfsmount and dentry, as long as all we are asking is "is it the same
address?". Dereferencing is not safe and apparmor ended up stepping into
that. d_namespace_path() really wants to examine the place where we stopped,
even if it's not connected to our namespace. As the result, it looked
at ->d_sb->s_magic of a dentry that might've been already freed by that point.
All other callers had been careful enough to avoid that, but it's really
a bad interface - it invites that kind of trouble.
The fix is fairly straightforward, even though it's bigger than I'd like:
* prepend_path() root argument becomes const.
* __d_path() is never called with NULL/NULL root. It was a kludge
to start with. Instead, we have an explicit function - d_absolute_root().
Same as __d_path(), except that it doesn't get root passed and stops where
it stops. apparmor and tomoyo are using it.
* __d_path() returns NULL on path outside of root. The main
caller is show_mountinfo() and that's precisely what we pass root for - to
skip those outside chroot jail. Those who don't want that can (and do)
use d_path().
* __d_path() root argument becomes const. Everyone agrees, I hope.
* apparmor does *NOT* try to use __d_path() or any of its variants
when it sees that path->mnt is an internal vfsmount. In that case it's
definitely not mounted anywhere and dentry_path() is exactly what we want
there. Handling of sysctl()-triggered weirdness is moved to that place.
* if apparmor is asked to do pathname relative to chroot jail
and __d_path() tells it we it's not in that jail, the sucker just calls
d_absolute_path() instead. That's the other remaining caller of __d_path(),
BTW.
* seq_path_root() does _NOT_ return -ENAMETOOLONG (it's stupid anyway -
the normal seq_file logics will take care of growing the buffer and redoing
the call of ->show() just fine). However, if it gets path not reachable
from root, it returns SEQ_SKIP. The only caller adjusted (i.e. stopped
ignoring the return value as it used to do).
Reviewed-by: John Johansen <john.johansen@canonical.com>
ACKed-by: John Johansen <john.johansen@canonical.com>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>
Cc: stable@vger.kernel.org
2011-12-05 08:43:34 -05:00
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2010-03-07 18:49:36 -08:00
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2013-03-15 01:45:51 -07:00
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2016-06-09 16:06:06 -05:00
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2013-03-24 14:28:27 -07:00
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2015-05-08 23:49:47 -05:00
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2013-03-24 14:28:27 -07:00
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2013-03-30 19:57:41 -07:00
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2013-03-24 14:28:27 -07:00
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2013-09-16 21:37:36 -04:00
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2013-03-24 14:28:27 -07:00
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2013-03-30 19:57:41 -07:00
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2015-06-04 09:43:11 -05:00
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2016-06-09 16:06:06 -05:00
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2013-03-30 19:57:41 -07:00
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2015-05-08 16:36:50 -05:00
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2016-06-15 06:59:49 -05:00
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2015-06-04 09:43:11 -05:00
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2016-06-10 12:21:40 -05:00
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2017-07-17 08:45:34 +01:00
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2016-06-10 12:21:40 -05:00
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2015-05-08 23:49:47 -05:00
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2015-06-04 09:43:11 -05:00
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2015-05-08 23:49:47 -05:00
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2015-06-04 09:43:11 -05:00
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2015-05-08 23:49:47 -05:00
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2015-01-07 08:10:09 -06:00
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2013-03-30 19:57:41 -07:00
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2015-01-07 08:10:09 -06:00
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2016-05-27 14:50:05 -05:00
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2015-01-07 08:10:09 -06:00
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2015-05-13 20:51:09 -05:00
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2013-03-30 19:57:41 -07:00
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2013-03-24 14:28:27 -07:00
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2015-05-08 23:49:47 -05:00
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2015-06-04 09:43:11 -05:00
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2013-03-30 19:57:41 -07:00
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2013-03-24 14:28:27 -07:00
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2013-03-30 19:57:41 -07:00
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2013-09-16 21:37:36 -04:00
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2013-03-30 19:57:41 -07:00
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2013-03-24 14:28:27 -07:00
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2016-06-09 16:06:06 -05:00
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2016-06-15 06:59:49 -05:00
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2016-06-09 16:06:06 -05:00
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2016-06-15 06:59:49 -05:00
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2016-06-09 16:06:06 -05:00
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fs: Treat foreign mounts as nosuid
If a process gets access to a mount from a different user
namespace, that process should not be able to take advantage of
setuid files or selinux entrypoints from that filesystem. Prevent
this by treating mounts from other mount namespaces and those not
owned by current_user_ns() or an ancestor as nosuid.
This will make it safer to allow more complex filesystems to be
mounted in non-root user namespaces.
This does not remove the need for MNT_LOCK_NOSUID. The setuid,
setgid, and file capability bits can no longer be abused if code in
a user namespace were to clear nosuid on an untrusted filesystem,
but this patch, by itself, is insufficient to protect the system
from abuse of files that, when execed, would increase MAC privilege.
As a more concrete explanation, any task that can manipulate a
vfsmount associated with a given user namespace already has
capabilities in that namespace and all of its descendents. If they
can cause a malicious setuid, setgid, or file-caps executable to
appear in that mount, then that executable will only allow them to
elevate privileges in exactly the set of namespaces in which they
are already privileges.
On the other hand, if they can cause a malicious executable to
appear with a dangerous MAC label, running it could change the
caller's security context in a way that should not have been
possible, even inside the namespace in which the task is confined.
As a hardening measure, this would have made CVE-2014-5207 much
more difficult to exploit.
Signed-off-by: Andy Lutomirski <luto@amacapital.net>
Signed-off-by: Seth Forshee <seth.forshee@canonical.com>
Acked-by: James Morris <james.l.morris@oracle.com>
Acked-by: Serge Hallyn <serge.hallyn@canonical.com>
Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
2016-06-23 16:41:05 -05:00
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2014-11-01 00:37:32 -04:00
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2010-03-07 18:49:36 -08:00
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2014-11-01 00:00:23 -04:00
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2010-03-07 18:49:36 -08:00
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2014-02-03 19:13:49 -08:00
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2010-03-07 18:49:36 -08:00
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2014-11-01 00:00:23 -04:00
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2010-03-07 18:49:36 -08:00
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2014-02-03 19:13:49 -08:00
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2010-03-07 18:49:36 -08:00
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2014-11-01 00:37:32 -04:00
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2010-03-07 18:49:36 -08:00
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2014-11-01 00:00:23 -04:00
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2010-03-07 18:49:36 -08:00
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2014-11-01 00:37:32 -04:00
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2010-03-07 18:49:36 -08:00
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2017-04-15 17:31:22 -04:00
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2010-03-07 18:49:36 -08:00
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2017-04-15 17:31:22 -04:00
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2010-03-07 18:49:36 -08:00
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2012-07-26 21:42:03 -07:00
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2013-03-20 12:49:49 -07:00
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2012-09-13 16:38:03 +08:00
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2010-03-07 18:49:36 -08:00
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2017-04-15 17:31:22 -04:00
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2010-03-07 18:49:36 -08:00
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2017-04-15 17:31:22 -04:00
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2010-03-07 18:49:36 -08:00
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2017-06-08 17:32:29 -07:00
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2010-03-07 18:49:36 -08:00
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2016-09-06 00:47:13 -07:00
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2010-03-07 18:49:36 -08:00
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2016-09-06 00:47:13 -07:00
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2010-03-07 18:49:36 -08:00
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