2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2015-03-13 16:08:06 -04:00
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2011-01-25 13:33:31 -05:00
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2006-01-02 19:04:38 +01:00
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2006-01-11 13:30:43 +01:00
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2006-01-02 19:04:38 +01:00
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2006-01-11 13:30:43 +01:00
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2006-01-02 19:04:38 +01:00
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2006-01-11 13:30:43 +01:00
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2006-01-02 19:04:38 +01:00
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2015-01-09 15:27:07 +08:00
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tipc: reduce locking scope during packet reception
We convert packet/message reception according to the same principle
we have been using for message sending and timeout handling:
We move the function tipc_rcv() to node.c, hence handling the initial
packet reception at the link aggregation level. The function grabs
the node lock, selects the receiving link, and accesses it via a new
call tipc_link_rcv(). This function appends buffers to the input
queue for delivery upwards, but it may also append outgoing packets
to the xmit queue, just as we do during regular message sending. The
latter will happen when buffers are forwarded from the link backlog,
or when retransmission is requested.
Upon return of this function, and after having released the node lock,
tipc_rcv() delivers/tranmsits the contents of those queues, but it may
also perform actions such as link activation or reset, as indicated by
the return flags from the link.
This reduces the number of cpu cycles spent inside the node spinlock,
and reduces contention on that lock.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-07-16 16:54:31 -04:00
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2006-01-02 19:04:38 +01:00
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2011-04-08 10:50:52 -04:00
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2006-01-02 19:04:38 +01:00
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tipc: resolve race problem at unicast message reception
TIPC handles message cardinality and sequencing at the link layer,
before passing messages upwards to the destination sockets. During the
upcall from link to socket no locks are held. It is therefore possible,
and we see it happen occasionally, that messages arriving in different
threads and delivered in sequence still bypass each other before they
reach the destination socket. This must not happen, since it violates
the sequentiality guarantee.
We solve this by adding a new input buffer queue to the link structure.
Arriving messages are added safely to the tail of that queue by the
link, while the head of the queue is consumed, also safely, by the
receiving socket. Sequentiality is secured per socket by only allowing
buffers to be dequeued inside the socket lock. Since there may be multiple
simultaneous readers of the queue, we use a 'filter' parameter to reduce
the risk that they peek the same buffer from the queue, hence also
reducing the risk of contention on the receiving socket locks.
This solves the sequentiality problem, and seems to cause no measurable
performance degradation.
A nice side effect of this change is that lock handling in the functions
tipc_rcv() and tipc_bcast_rcv() now becomes uniform, something that
will enable future simplifications of those functions.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-02-05 08:36:41 -05:00
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2008-03-06 15:06:55 -08:00
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2010-11-30 12:00:53 +00:00
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2011-04-08 10:50:52 -04:00
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tipc: clean up handling of message priorities
Messages transferred by TIPC are assigned an "importance priority", -an
integer value indicating how to treat the message when there is link or
destination socket congestion.
There is no separate header field for this value. Instead, the message
user values have been chosen in ascending order according to perceived
importance, so that the message user field can be used for this.
This is not a good solution. First, we have many more users than the
needed priority levels, so we end up with treating more priority
levels than necessary. Second, the user field cannot always
accurately reflect the priority of the message. E.g., a message
fragment packet should really have the priority of the enveloped
user data message, and not the priority of the MSG_FRAGMENTER user.
Until now, we have been working around this problem in different ways,
but it is now time to implement a consistent way of handling such
priorities, although still within the constraint that we cannot
allocate any more bits in the regular data message header for this.
In this commit, we define a new priority level, TIPC_SYSTEM_IMPORTANCE,
that will be the only one used apart from the four (lower) user data
levels. All non-data messages map down to this priority. Furthermore,
we take some free bits from the MSG_FRAGMENTER header and allocate
them to store the priority of the enveloped message. We then adjust
the functions msg_importance()/msg_set_importance() so that they
read/set the correct header fields depending on user type.
This small protocol change is fully compatible, because the code at
the receiving end of a link currently reads the importance level
only from user data messages, where there is no change.
Reviewed-by: Erik Hugne <erik.hugne@ericsson.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-03-13 16:08:11 -04:00
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2011-04-08 10:50:52 -04:00
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2010-11-30 12:00:53 +00:00
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tipc: clean up handling of message priorities
Messages transferred by TIPC are assigned an "importance priority", -an
integer value indicating how to treat the message when there is link or
destination socket congestion.
There is no separate header field for this value. Instead, the message
user values have been chosen in ascending order according to perceived
importance, so that the message user field can be used for this.
This is not a good solution. First, we have many more users than the
needed priority levels, so we end up with treating more priority
levels than necessary. Second, the user field cannot always
accurately reflect the priority of the message. E.g., a message
fragment packet should really have the priority of the enveloped
user data message, and not the priority of the MSG_FRAGMENTER user.
Until now, we have been working around this problem in different ways,
but it is now time to implement a consistent way of handling such
priorities, although still within the constraint that we cannot
allocate any more bits in the regular data message header for this.
In this commit, we define a new priority level, TIPC_SYSTEM_IMPORTANCE,
that will be the only one used apart from the four (lower) user data
levels. All non-data messages map down to this priority. Furthermore,
we take some free bits from the MSG_FRAGMENTER header and allocate
them to store the priority of the enveloped message. We then adjust
the functions msg_importance()/msg_set_importance() so that they
read/set the correct header fields depending on user type.
This small protocol change is fully compatible, because the code at
the receiving end of a link currently reads the importance level
only from user data messages, where there is no change.
Reviewed-by: Erik Hugne <erik.hugne@ericsson.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-03-13 16:08:11 -04:00
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tipc: eliminate delayed link deletion at link failover
When a bearer is disabled manually, all its links have to be reset
and deleted. However, if there is a remaining, parallel link ready
to take over a deleted link's traffic, we currently delay the delete
of the removed link until the failover procedure is finished. This
is because the remaining link needs to access state from the reset
link, such as the last received packet number, and any partially
reassembled buffer, in order to perform a successful failover.
In this commit, we do instead move the state data over to the new
link, so that it can fulfill the procedure autonomously, without
accessing any data on the old link. This means that we can now
proceed and delete all pertaining links immediately when a bearer
is disabled. This saves us from some unnecessary complexity in such
situations.
We also choose to change the confusing definitions CHANGEOVER_PROTOCOL,
ORIGINAL_MSG and DUPLICATE_MSG to the more descriptive TUNNEL_PROTOCOL,
FAILOVER_MSG and SYNCH_MSG respectively.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-04-02 09:33:01 -04:00
|
|
|
|
tipc: clean up handling of message priorities
Messages transferred by TIPC are assigned an "importance priority", -an
integer value indicating how to treat the message when there is link or
destination socket congestion.
There is no separate header field for this value. Instead, the message
user values have been chosen in ascending order according to perceived
importance, so that the message user field can be used for this.
This is not a good solution. First, we have many more users than the
needed priority levels, so we end up with treating more priority
levels than necessary. Second, the user field cannot always
accurately reflect the priority of the message. E.g., a message
fragment packet should really have the priority of the enveloped
user data message, and not the priority of the MSG_FRAGMENTER user.
Until now, we have been working around this problem in different ways,
but it is now time to implement a consistent way of handling such
priorities, although still within the constraint that we cannot
allocate any more bits in the regular data message header for this.
In this commit, we define a new priority level, TIPC_SYSTEM_IMPORTANCE,
that will be the only one used apart from the four (lower) user data
levels. All non-data messages map down to this priority. Furthermore,
we take some free bits from the MSG_FRAGMENTER header and allocate
them to store the priority of the enveloped message. We then adjust
the functions msg_importance()/msg_set_importance() so that they
read/set the correct header fields depending on user type.
This small protocol change is fully compatible, because the code at
the receiving end of a link currently reads the importance level
only from user data messages, where there is no change.
Reviewed-by: Erik Hugne <erik.hugne@ericsson.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-03-13 16:08:11 -04:00
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|
2011-04-08 10:50:52 -04:00
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|
2011-05-31 15:03:18 -04:00
|
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|
2008-03-06 15:06:55 -08:00
|
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2006-01-02 19:04:38 +01:00
|
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|
2015-02-27 08:56:57 +01:00
|
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|
2011-10-07 15:19:11 -04:00
|
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|
2015-01-09 15:27:01 +08:00
|
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2015-03-13 16:08:06 -04:00
|
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2015-01-09 15:27:01 +08:00
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|
tipc: make struct tipc_link generic to support broadcast
Realizing that unicast is just a special case of broadcast, we also see
that we can go in the other direction, i.e., that modest changes to the
current unicast link can make it generic enough to support broadcast.
The following changes are introduced here:
- A new counter ("ackers") in struct tipc_link, to indicate how many
peers need to ack a packet before it can be released.
- A corresponding counter in the skb user area, to keep track of how
many peers a are left to ack before a buffer can be released.
- A new counter ("acked"), to keep persistent track of how far a peer
has acked at the moment, i.e., where in the transmission queue to
start updating buffers when the next ack arrives. This is to avoid
double acknowledgements from a peer, with inadvertent relase of
packets as a result.
- A more generic tipc_link_retrans() function, where retransmit starts
from a given sequence number, instead of the first packet in the
transmision queue. This is to minimize the number of retransmitted
packets on the broadcast media.
When the new functionality is taken into use in the next commits,
we expect it to have minimal effect on unicast mode performance.
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-10-22 08:51:38 -04:00
|
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|
2015-01-09 15:27:01 +08:00
|
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|
2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
|
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|
2015-01-09 15:27:01 +08:00
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2010-11-30 12:00:53 +00:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2006-01-02 19:04:38 +01:00
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2007-04-24 14:51:55 -07:00
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2008-04-26 22:42:14 -07:00
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2006-01-02 19:04:38 +01:00
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2008-06-04 17:54:48 -07:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2008-03-06 15:07:42 -08:00
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2006-01-02 19:04:38 +01:00
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2010-09-22 20:43:57 +00:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2010-12-31 18:59:32 +00:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2008-06-04 17:54:48 -07:00
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2006-01-02 19:04:38 +01:00
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2008-06-04 17:54:48 -07:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2015-03-25 12:07:25 -04:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2015-05-14 10:46:14 -04:00
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2006-01-02 19:04:38 +01:00
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2015-05-14 10:46:14 -04:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2015-05-14 10:46:14 -04:00
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2006-01-02 19:04:38 +01:00
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2015-05-14 10:46:14 -04:00
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2006-01-02 19:04:38 +01:00
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2015-05-14 10:46:14 -04:00
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2006-01-02 19:04:38 +01:00
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2015-05-14 10:46:14 -04:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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tipc: clean up handling of message priorities
Messages transferred by TIPC are assigned an "importance priority", -an
integer value indicating how to treat the message when there is link or
destination socket congestion.
There is no separate header field for this value. Instead, the message
user values have been chosen in ascending order according to perceived
importance, so that the message user field can be used for this.
This is not a good solution. First, we have many more users than the
needed priority levels, so we end up with treating more priority
levels than necessary. Second, the user field cannot always
accurately reflect the priority of the message. E.g., a message
fragment packet should really have the priority of the enveloped
user data message, and not the priority of the MSG_FRAGMENTER user.
Until now, we have been working around this problem in different ways,
but it is now time to implement a consistent way of handling such
priorities, although still within the constraint that we cannot
allocate any more bits in the regular data message header for this.
In this commit, we define a new priority level, TIPC_SYSTEM_IMPORTANCE,
that will be the only one used apart from the four (lower) user data
levels. All non-data messages map down to this priority. Furthermore,
we take some free bits from the MSG_FRAGMENTER header and allocate
them to store the priority of the enveloped message. We then adjust
the functions msg_importance()/msg_set_importance() so that they
read/set the correct header fields depending on user type.
This small protocol change is fully compatible, because the code at
the receiving end of a link currently reads the importance level
only from user data messages, where there is no change.
Reviewed-by: Erik Hugne <erik.hugne@ericsson.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-03-13 16:08:11 -04:00
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tipc: improve link congestion algorithm
The link congestion algorithm used until now implies two problems.
- It is too generous towards lower-level messages in situations of high
load by giving "absolute" bandwidth guarantees to the different
priority levels. LOW traffic is guaranteed 10%, MEDIUM is guaranted
20%, HIGH is guaranteed 30%, and CRITICAL is guaranteed 40% of the
available bandwidth. But, in the absence of higher level traffic, the
ratio between two distinct levels becomes unreasonable. E.g. if there
is only LOW and MEDIUM traffic on a system, the former is guaranteed
1/3 of the bandwidth, and the latter 2/3. This again means that if
there is e.g. one LOW user and 10 MEDIUM users, the former will have
33.3% of the bandwidth, and the others will have to compete for the
remainder, i.e. each will end up with 6.7% of the capacity.
- Packets of type MSG_BUNDLER are created at SYSTEM importance level,
but only after the packets bundled into it have passed the congestion
test for their own respective levels. Since bundled packets don't
result in incrementing the level counter for their own importance,
only occasionally for the SYSTEM level counter, they do in practice
obtain SYSTEM level importance. Hence, the current implementation
provides a gap in the congestion algorithm that in the worst case
may lead to a link reset.
We now refine the congestion algorithm as follows:
- A message is accepted to the link backlog only if its own level
counter, and all superior level counters, permit it.
- The importance of a created bundle packet is set according to its
contents. A bundle packet created from messges at levels LOW to
CRITICAL is given importance level CRITICAL, while a bundle created
from a SYSTEM level message is given importance SYSTEM. In the latter
case only subsequent SYSTEM level messages are allowed to be bundled
into it.
This solves the first problem described above, by making the bandwidth
guarantee relative to the total number of users at all levels; only
the upper limit for each level remains absolute. In the example
described above, the single LOW user would use 1/11th of the bandwidth,
the same as each of the ten MEDIUM users, but he still has the same
guarantee against starvation as the latter ones.
The fix also solves the second problem. If the CRITICAL level is filled
up by bundle packets of that level, no lower level packets will be
accepted any more.
Suggested-by: Gergely Kiss <gergely.kiss@ericsson.com>
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-05-14 10:46:17 -04:00
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2015-10-14 09:23:18 -04:00
|
|
|
|
tipc: clean up handling of message priorities
Messages transferred by TIPC are assigned an "importance priority", -an
integer value indicating how to treat the message when there is link or
destination socket congestion.
There is no separate header field for this value. Instead, the message
user values have been chosen in ascending order according to perceived
importance, so that the message user field can be used for this.
This is not a good solution. First, we have many more users than the
needed priority levels, so we end up with treating more priority
levels than necessary. Second, the user field cannot always
accurately reflect the priority of the message. E.g., a message
fragment packet should really have the priority of the enveloped
user data message, and not the priority of the MSG_FRAGMENTER user.
Until now, we have been working around this problem in different ways,
but it is now time to implement a consistent way of handling such
priorities, although still within the constraint that we cannot
allocate any more bits in the regular data message header for this.
In this commit, we define a new priority level, TIPC_SYSTEM_IMPORTANCE,
that will be the only one used apart from the four (lower) user data
levels. All non-data messages map down to this priority. Furthermore,
we take some free bits from the MSG_FRAGMENTER header and allocate
them to store the priority of the enveloped message. We then adjust
the functions msg_importance()/msg_set_importance() so that they
read/set the correct header fields depending on user type.
This small protocol change is fully compatible, because the code at
the receiving end of a link currently reads the importance level
only from user data messages, where there is no change.
Reviewed-by: Erik Hugne <erik.hugne@ericsson.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-03-13 16:08:11 -04:00
|
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|
|
tipc: improve link congestion algorithm
The link congestion algorithm used until now implies two problems.
- It is too generous towards lower-level messages in situations of high
load by giving "absolute" bandwidth guarantees to the different
priority levels. LOW traffic is guaranteed 10%, MEDIUM is guaranted
20%, HIGH is guaranteed 30%, and CRITICAL is guaranteed 40% of the
available bandwidth. But, in the absence of higher level traffic, the
ratio between two distinct levels becomes unreasonable. E.g. if there
is only LOW and MEDIUM traffic on a system, the former is guaranteed
1/3 of the bandwidth, and the latter 2/3. This again means that if
there is e.g. one LOW user and 10 MEDIUM users, the former will have
33.3% of the bandwidth, and the others will have to compete for the
remainder, i.e. each will end up with 6.7% of the capacity.
- Packets of type MSG_BUNDLER are created at SYSTEM importance level,
but only after the packets bundled into it have passed the congestion
test for their own respective levels. Since bundled packets don't
result in incrementing the level counter for their own importance,
only occasionally for the SYSTEM level counter, they do in practice
obtain SYSTEM level importance. Hence, the current implementation
provides a gap in the congestion algorithm that in the worst case
may lead to a link reset.
We now refine the congestion algorithm as follows:
- A message is accepted to the link backlog only if its own level
counter, and all superior level counters, permit it.
- The importance of a created bundle packet is set according to its
contents. A bundle packet created from messges at levels LOW to
CRITICAL is given importance level CRITICAL, while a bundle created
from a SYSTEM level message is given importance SYSTEM. In the latter
case only subsequent SYSTEM level messages are allowed to be bundled
into it.
This solves the first problem described above, by making the bandwidth
guarantee relative to the total number of users at all levels; only
the upper limit for each level remains absolute. In the example
described above, the single LOW user would use 1/11th of the bandwidth,
the same as each of the ten MEDIUM users, but he still has the same
guarantee against starvation as the latter ones.
The fix also solves the second problem. If the CRITICAL level is filled
up by bundle packets of that level, no lower level packets will be
accepted any more.
Suggested-by: Gergely Kiss <gergely.kiss@ericsson.com>
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-05-14 10:46:17 -04:00
|
|
|
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|
|
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|
|
|
|
2015-10-14 09:23:18 -04:00
|
|
|
|
tipc: improve link congestion algorithm
The link congestion algorithm used until now implies two problems.
- It is too generous towards lower-level messages in situations of high
load by giving "absolute" bandwidth guarantees to the different
priority levels. LOW traffic is guaranteed 10%, MEDIUM is guaranted
20%, HIGH is guaranteed 30%, and CRITICAL is guaranteed 40% of the
available bandwidth. But, in the absence of higher level traffic, the
ratio between two distinct levels becomes unreasonable. E.g. if there
is only LOW and MEDIUM traffic on a system, the former is guaranteed
1/3 of the bandwidth, and the latter 2/3. This again means that if
there is e.g. one LOW user and 10 MEDIUM users, the former will have
33.3% of the bandwidth, and the others will have to compete for the
remainder, i.e. each will end up with 6.7% of the capacity.
- Packets of type MSG_BUNDLER are created at SYSTEM importance level,
but only after the packets bundled into it have passed the congestion
test for their own respective levels. Since bundled packets don't
result in incrementing the level counter for their own importance,
only occasionally for the SYSTEM level counter, they do in practice
obtain SYSTEM level importance. Hence, the current implementation
provides a gap in the congestion algorithm that in the worst case
may lead to a link reset.
We now refine the congestion algorithm as follows:
- A message is accepted to the link backlog only if its own level
counter, and all superior level counters, permit it.
- The importance of a created bundle packet is set according to its
contents. A bundle packet created from messges at levels LOW to
CRITICAL is given importance level CRITICAL, while a bundle created
from a SYSTEM level message is given importance SYSTEM. In the latter
case only subsequent SYSTEM level messages are allowed to be bundled
into it.
This solves the first problem described above, by making the bandwidth
guarantee relative to the total number of users at all levels; only
the upper limit for each level remains absolute. In the example
described above, the single LOW user would use 1/11th of the bandwidth,
the same as each of the ten MEDIUM users, but he still has the same
guarantee against starvation as the latter ones.
The fix also solves the second problem. If the CRITICAL level is filled
up by bundle packets of that level, no lower level packets will be
accepted any more.
Suggested-by: Gergely Kiss <gergely.kiss@ericsson.com>
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-05-14 10:46:17 -04:00
|
|
|
|
tipc: clean up handling of message priorities
Messages transferred by TIPC are assigned an "importance priority", -an
integer value indicating how to treat the message when there is link or
destination socket congestion.
There is no separate header field for this value. Instead, the message
user values have been chosen in ascending order according to perceived
importance, so that the message user field can be used for this.
This is not a good solution. First, we have many more users than the
needed priority levels, so we end up with treating more priority
levels than necessary. Second, the user field cannot always
accurately reflect the priority of the message. E.g., a message
fragment packet should really have the priority of the enveloped
user data message, and not the priority of the MSG_FRAGMENTER user.
Until now, we have been working around this problem in different ways,
but it is now time to implement a consistent way of handling such
priorities, although still within the constraint that we cannot
allocate any more bits in the regular data message header for this.
In this commit, we define a new priority level, TIPC_SYSTEM_IMPORTANCE,
that will be the only one used apart from the four (lower) user data
levels. All non-data messages map down to this priority. Furthermore,
we take some free bits from the MSG_FRAGMENTER header and allocate
them to store the priority of the enveloped message. We then adjust
the functions msg_importance()/msg_set_importance() so that they
read/set the correct header fields depending on user type.
This small protocol change is fully compatible, because the code at
the receiving end of a link currently reads the importance level
only from user data messages, where there is no change.
Reviewed-by: Erik Hugne <erik.hugne@ericsson.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-03-13 16:08:11 -04:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2015-03-25 12:07:25 -04:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2011-05-31 15:03:18 -04:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2010-11-30 12:00:53 +00:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2011-04-08 10:50:52 -04:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2011-04-08 10:50:52 -04:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2011-04-08 10:50:52 -04:00
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2006-01-02 19:04:38 +01:00
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2011-04-08 11:04:15 -04:00
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tipc: eliminate delayed link deletion at link failover
When a bearer is disabled manually, all its links have to be reset
and deleted. However, if there is a remaining, parallel link ready
to take over a deleted link's traffic, we currently delay the delete
of the removed link until the failover procedure is finished. This
is because the remaining link needs to access state from the reset
link, such as the last received packet number, and any partially
reassembled buffer, in order to perform a successful failover.
In this commit, we do instead move the state data over to the new
link, so that it can fulfill the procedure autonomously, without
accessing any data on the old link. This means that we can now
proceed and delete all pertaining links immediately when a bearer
is disabled. This saves us from some unnecessary complexity in such
situations.
We also choose to change the confusing definitions CHANGEOVER_PROTOCOL,
ORIGINAL_MSG and DUPLICATE_MSG to the more descriptive TUNNEL_PROTOCOL,
FAILOVER_MSG and SYNCH_MSG respectively.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-04-02 09:33:01 -04:00
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2011-04-08 11:04:15 -04:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2008-06-04 17:47:30 -07:00
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2006-01-02 19:04:38 +01:00
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2008-06-04 17:47:30 -07:00
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2006-01-02 19:04:38 +01:00
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2011-10-28 16:26:41 -04:00
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2015-03-13 16:08:05 -04:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2015-10-22 08:51:41 -04:00
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2006-01-02 19:04:38 +01:00
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2015-07-30 18:24:19 -04:00
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2006-01-02 19:04:38 +01:00
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2015-07-30 18:24:19 -04:00
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2006-01-02 19:04:38 +01:00
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2011-05-25 13:28:27 -04:00
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2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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tipc: reduce locking scope during packet reception
We convert packet/message reception according to the same principle
we have been using for message sending and timeout handling:
We move the function tipc_rcv() to node.c, hence handling the initial
packet reception at the link aggregation level. The function grabs
the node lock, selects the receiving link, and accesses it via a new
call tipc_link_rcv(). This function appends buffers to the input
queue for delivery upwards, but it may also append outgoing packets
to the xmit queue, just as we do during regular message sending. The
latter will happen when buffers are forwarded from the link backlog,
or when retransmission is requested.
Upon return of this function, and after having released the node lock,
tipc_rcv() delivers/tranmsits the contents of those queues, but it may
also perform actions such as link activation or reset, as indicated by
the return flags from the link.
This reduces the number of cpu cycles spent inside the node spinlock,
and reduces contention on that lock.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-07-16 16:54:31 -04:00
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2006-01-02 19:04:38 +01:00
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tipc: reduce locking scope during packet reception
We convert packet/message reception according to the same principle
we have been using for message sending and timeout handling:
We move the function tipc_rcv() to node.c, hence handling the initial
packet reception at the link aggregation level. The function grabs
the node lock, selects the receiving link, and accesses it via a new
call tipc_link_rcv(). This function appends buffers to the input
queue for delivery upwards, but it may also append outgoing packets
to the xmit queue, just as we do during regular message sending. The
latter will happen when buffers are forwarded from the link backlog,
or when retransmission is requested.
Upon return of this function, and after having released the node lock,
tipc_rcv() delivers/tranmsits the contents of those queues, but it may
also perform actions such as link activation or reset, as indicated by
the return flags from the link.
This reduces the number of cpu cycles spent inside the node spinlock,
and reduces contention on that lock.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-07-16 16:54:31 -04:00
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2006-01-02 19:04:38 +01:00
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2011-05-25 13:28:27 -04:00
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2006-01-02 19:04:38 +01:00
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2011-02-28 15:30:20 -05:00
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2006-01-02 19:04:38 +01:00
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2011-02-28 15:30:20 -05:00
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2006-01-02 19:04:38 +01:00
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tipc: guarantee peer bearer id exchange after reboot
When a link endpoint is going down locally, e.g., because its interface
is being stopped, it will spontaneously send out a RESET message to
its peer, informing it about this fact. This saves the peer from
detecting the failure via probing, and hence gives both speedier and
less resource consuming failure detection on the peer side.
According to the link FSM, a receiver of a RESET message, ignoring the
reason for it, must now consider the sender ready to come back up, and
starts periodically sending out ACTIVATE messages to the peer in order
to re-establish the link. Also, according to the FSM, the receiver of
an ACTIVATE message can now go directly to state ESTABLISHED and start
sending regular traffic packets. This is a well-proven and robust FSM.
However, in the case of a reboot, there is a small possibilty that link
endpoint on the rebooted node may have been re-created with a new bearer
identity between the moment it sent its (pre-boot) RESET and the moment
it receives the ACTIVATE from the peer. The new bearer identity cannot
be known by the peer according to this scenario, since traffic headers
don't convey such information. This is a problem, because both endpoints
need to know the correct value of the peer's bearer id at any moment in
time in order to be able to produce correct link events for their users.
The only way to guarantee this is to enforce a full setup message
exchange (RESET + ACTIVATE) even after the reboot, since those messages
carry the bearer idientity in their header.
In this commit we do this by introducing and setting a "stopping" bit in
the header of the spontaneously generated RESET messages, informing the
peer that the sender will not be immediately ready to re-establish the
link. A receiver seeing this bit must act as if this were a locally
detected connectivity failure, and hence has to go through a full two-
way setup message exchange before any link can be re-established.
Although never reported, this problem seems to have always been around.
This protocol addition is fully backwards compatible.
Acked-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2016-04-15 13:33:03 -04:00
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|
2011-10-07 15:19:11 -04:00
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|
|
2015-02-27 08:56:57 +01:00
|
|
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|
2011-10-07 15:19:11 -04:00
|
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|
2006-01-02 19:04:38 +01:00
|
|
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|
2007-02-09 23:25:21 +09:00
|
|
|
|
2006-01-02 19:04:38 +01:00
|
|
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|
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|
|
2015-07-30 18:24:19 -04:00
|
|
|
|
2006-01-02 19:04:38 +01:00
|
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|
2015-07-30 18:24:19 -04:00
|
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2006-01-02 19:04:38 +01:00
|
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tipc: redesign connection-level flow control
There are two flow control mechanisms in TIPC; one at link level that
handles network congestion, burst control, and retransmission, and one
at connection level which' only remaining task is to prevent overflow
in the receiving socket buffer. In TIPC, the latter task has to be
solved end-to-end because messages can not be thrown away once they
have been accepted and delivered upwards from the link layer, i.e, we
can never permit the receive buffer to overflow.
Currently, this algorithm is message based. A counter in the receiving
socket keeps track of number of consumed messages, and sends a dedicated
acknowledge message back to the sender for each 256 consumed message.
A counter at the sending end keeps track of the sent, not yet
acknowledged messages, and blocks the sender if this number ever reaches
512 unacknowledged messages. When the missing acknowledge arrives, the
socket is then woken up for renewed transmission. This works well for
keeping the message flow running, as it almost never happens that a
sender socket is blocked this way.
A problem with the current mechanism is that it potentially is very
memory consuming. Since we don't distinguish between small and large
messages, we have to dimension the socket receive buffer according
to a worst-case of both. I.e., the window size must be chosen large
enough to sustain a reasonable throughput even for the smallest
messages, while we must still consider a scenario where all messages
are of maximum size. Hence, the current fix window size of 512 messages
and a maximum message size of 66k results in a receive buffer of 66 MB
when truesize(66k) = 131k is taken into account. It is possible to do
much better.
This commit introduces an algorithm where we instead use 1024-byte
blocks as base unit. This unit, always rounded upwards from the
actual message size, is used when we advertise windows as well as when
we count and acknowledge transmitted data. The advertised window is
based on the configured receive buffer size in such a way that even
the worst-case truesize/msgsize ratio always is covered. Since the
smallest possible message size (from a flow control viewpoint) now is
1024 bytes, we can safely assume this ratio to be less than four, which
is the value we are now using.
This way, we have been able to reduce the default receive buffer size
from 66 MB to 2 MB with maintained performance.
In order to keep this solution backwards compatible, we introduce a
new capability bit in the discovery protocol, and use this throughout
the message sending/reception path to always select the right unit.
Acked-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2016-05-02 11:58:47 -04:00
|
|
|
|
2006-01-02 19:04:38 +01:00
|
|
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|
|
tipc: redesign connection-level flow control
There are two flow control mechanisms in TIPC; one at link level that
handles network congestion, burst control, and retransmission, and one
at connection level which' only remaining task is to prevent overflow
in the receiving socket buffer. In TIPC, the latter task has to be
solved end-to-end because messages can not be thrown away once they
have been accepted and delivered upwards from the link layer, i.e, we
can never permit the receive buffer to overflow.
Currently, this algorithm is message based. A counter in the receiving
socket keeps track of number of consumed messages, and sends a dedicated
acknowledge message back to the sender for each 256 consumed message.
A counter at the sending end keeps track of the sent, not yet
acknowledged messages, and blocks the sender if this number ever reaches
512 unacknowledged messages. When the missing acknowledge arrives, the
socket is then woken up for renewed transmission. This works well for
keeping the message flow running, as it almost never happens that a
sender socket is blocked this way.
A problem with the current mechanism is that it potentially is very
memory consuming. Since we don't distinguish between small and large
messages, we have to dimension the socket receive buffer according
to a worst-case of both. I.e., the window size must be chosen large
enough to sustain a reasonable throughput even for the smallest
messages, while we must still consider a scenario where all messages
are of maximum size. Hence, the current fix window size of 512 messages
and a maximum message size of 66k results in a receive buffer of 66 MB
when truesize(66k) = 131k is taken into account. It is possible to do
much better.
This commit introduces an algorithm where we instead use 1024-byte
blocks as base unit. This unit, always rounded upwards from the
actual message size, is used when we advertise windows as well as when
we count and acknowledge transmitted data. The advertised window is
based on the configured receive buffer size in such a way that even
the worst-case truesize/msgsize ratio always is covered. Since the
smallest possible message size (from a flow control viewpoint) now is
1024 bytes, we can safely assume this ratio to be less than four, which
is the value we are now using.
This way, we have been able to reduce the default receive buffer size
from 66 MB to 2 MB with maintained performance.
In order to keep this solution backwards compatible, we introduce a
new capability bit in the discovery protocol, and use this throughout
the message sending/reception path to always select the right unit.
Acked-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2016-05-02 11:58:47 -04:00
|
|
|
|
2006-01-02 19:04:38 +01:00
|
|
|
|
|
|
|
|
|
|
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|
|
|
|
|
|
tipc: redesign connection-level flow control
There are two flow control mechanisms in TIPC; one at link level that
handles network congestion, burst control, and retransmission, and one
at connection level which' only remaining task is to prevent overflow
in the receiving socket buffer. In TIPC, the latter task has to be
solved end-to-end because messages can not be thrown away once they
have been accepted and delivered upwards from the link layer, i.e, we
can never permit the receive buffer to overflow.
Currently, this algorithm is message based. A counter in the receiving
socket keeps track of number of consumed messages, and sends a dedicated
acknowledge message back to the sender for each 256 consumed message.
A counter at the sending end keeps track of the sent, not yet
acknowledged messages, and blocks the sender if this number ever reaches
512 unacknowledged messages. When the missing acknowledge arrives, the
socket is then woken up for renewed transmission. This works well for
keeping the message flow running, as it almost never happens that a
sender socket is blocked this way.
A problem with the current mechanism is that it potentially is very
memory consuming. Since we don't distinguish between small and large
messages, we have to dimension the socket receive buffer according
to a worst-case of both. I.e., the window size must be chosen large
enough to sustain a reasonable throughput even for the smallest
messages, while we must still consider a scenario where all messages
are of maximum size. Hence, the current fix window size of 512 messages
and a maximum message size of 66k results in a receive buffer of 66 MB
when truesize(66k) = 131k is taken into account. It is possible to do
much better.
This commit introduces an algorithm where we instead use 1024-byte
blocks as base unit. This unit, always rounded upwards from the
actual message size, is used when we advertise windows as well as when
we count and acknowledge transmitted data. The advertised window is
based on the configured receive buffer size in such a way that even
the worst-case truesize/msgsize ratio always is covered. Since the
smallest possible message size (from a flow control viewpoint) now is
1024 bytes, we can safely assume this ratio to be less than four, which
is the value we are now using.
This way, we have been able to reduce the default receive buffer size
from 66 MB to 2 MB with maintained performance.
In order to keep this solution backwards compatible, we introduce a
new capability bit in the discovery protocol, and use this throughout
the message sending/reception path to always select the right unit.
Acked-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2016-05-02 11:58:47 -04:00
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2007-02-09 23:25:21 +09:00
|
|
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|
2006-01-02 19:04:38 +01:00
|
|
|
|
2010-09-22 20:43:57 +00:00
|
|
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|
2006-01-02 19:04:38 +01:00
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2007-02-09 23:25:21 +09:00
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2006-01-02 19:04:38 +01:00
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2015-07-30 18:24:19 -04:00
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tipc: reduce locking scope during packet reception
We convert packet/message reception according to the same principle
we have been using for message sending and timeout handling:
We move the function tipc_rcv() to node.c, hence handling the initial
packet reception at the link aggregation level. The function grabs
the node lock, selects the receiving link, and accesses it via a new
call tipc_link_rcv(). This function appends buffers to the input
queue for delivery upwards, but it may also append outgoing packets
to the xmit queue, just as we do during regular message sending. The
latter will happen when buffers are forwarded from the link backlog,
or when retransmission is requested.
Upon return of this function, and after having released the node lock,
tipc_rcv() delivers/tranmsits the contents of those queues, but it may
also perform actions such as link activation or reset, as indicated by
the return flags from the link.
This reduces the number of cpu cycles spent inside the node spinlock,
and reduces contention on that lock.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-07-16 16:54:31 -04:00
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2015-07-30 18:24:19 -04:00
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tipc: reduce locking scope during packet reception
We convert packet/message reception according to the same principle
we have been using for message sending and timeout handling:
We move the function tipc_rcv() to node.c, hence handling the initial
packet reception at the link aggregation level. The function grabs
the node lock, selects the receiving link, and accesses it via a new
call tipc_link_rcv(). This function appends buffers to the input
queue for delivery upwards, but it may also append outgoing packets
to the xmit queue, just as we do during regular message sending. The
latter will happen when buffers are forwarded from the link backlog,
or when retransmission is requested.
Upon return of this function, and after having released the node lock,
tipc_rcv() delivers/tranmsits the contents of those queues, but it may
also perform actions such as link activation or reset, as indicated by
the return flags from the link.
This reduces the number of cpu cycles spent inside the node spinlock,
and reduces contention on that lock.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-07-16 16:54:31 -04:00
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2015-07-30 18:24:19 -04:00
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2015-07-16 16:54:30 -04:00
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2015-07-30 18:24:19 -04:00
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2015-07-16 16:54:30 -04:00
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2016-04-07 10:09:14 -04:00
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2015-01-09 15:27:01 +08:00
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2015-03-13 16:08:06 -04:00
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tipc: introduce new tipc_sk_respond() function
Currently, we use the code sequence
if (msg_reverse())
tipc_link_xmit_skb()
at numerous locations in socket.c. The preparation of arguments
for these calls, as well as the sequence itself, makes the code
unecessarily complex.
In this commit, we introduce a new function, tipc_sk_respond(),
that performs this call combination. We also replace some, but not
yet all, of these explicit call sequences with calls to the new
function. Notably, we let the function tipc_sk_proto_rcv() use
the new function to directly send out PROBE_REPLY messages,
instead of deferring this to the calling tipc_sk_rcv() function,
as we do now.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-07-22 10:11:19 -04:00
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2015-02-05 08:36:36 -05:00
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2015-01-09 15:27:10 +08:00
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tipc: split up function tipc_msg_eval()
The function tipc_msg_eval() is in reality doing two related, but
different tasks. First it tries to find a new destination for named
messages, in case there was no first lookup, or if the first lookup
failed. Second, it does what its name suggests, evaluating the validity
of the message and its destination, and returning an appropriate error
code depending on the result.
This is confusing, and in this commit we choose to break it up into two
functions. A new function, tipc_msg_lookup_dest(), first attempts to find
a new destination, if the message is of the right type. If this lookup
fails, or if the message should not be subject to a second lookup, the
already existing tipc_msg_reverse() is called. This function performs
prepares the message for rejection, if applicable.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-02-05 08:36:39 -05:00
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2014-05-14 05:39:12 -04:00
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2015-05-14 10:46:18 -04:00
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tipc: resolve race problem at unicast message reception
TIPC handles message cardinality and sequencing at the link layer,
before passing messages upwards to the destination sockets. During the
upcall from link to socket no locks are held. It is therefore possible,
and we see it happen occasionally, that messages arriving in different
threads and delivered in sequence still bypass each other before they
reach the destination socket. This must not happen, since it violates
the sequentiality guarantee.
We solve this by adding a new input buffer queue to the link structure.
Arriving messages are added safely to the tail of that queue by the
link, while the head of the queue is consumed, also safely, by the
receiving socket. Sequentiality is secured per socket by only allowing
buffers to be dequeued inside the socket lock. Since there may be multiple
simultaneous readers of the queue, we use a 'filter' parameter to reduce
the risk that they peek the same buffer from the queue, hence also
reducing the risk of contention on the receiving socket locks.
This solves the sequentiality problem, and seems to cause no measurable
performance degradation.
A nice side effect of this change is that lock handling in the functions
tipc_rcv() and tipc_bcast_rcv() now becomes uniform, something that
will enable future simplifications of those functions.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-02-05 08:36:41 -05:00
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2015-02-05 08:36:36 -05:00
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2015-01-09 15:27:10 +08:00
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2015-07-22 10:11:20 -04:00
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2015-10-22 08:51:39 -04:00
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2015-10-15 14:52:43 -04:00
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2014-07-16 20:41:00 -04:00
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2015-05-14 10:46:14 -04:00
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2015-02-05 08:36:44 -05:00
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tipc: resolve race problem at unicast message reception
TIPC handles message cardinality and sequencing at the link layer,
before passing messages upwards to the destination sockets. During the
upcall from link to socket no locks are held. It is therefore possible,
and we see it happen occasionally, that messages arriving in different
threads and delivered in sequence still bypass each other before they
reach the destination socket. This must not happen, since it violates
the sequentiality guarantee.
We solve this by adding a new input buffer queue to the link structure.
Arriving messages are added safely to the tail of that queue by the
link, while the head of the queue is consumed, also safely, by the
receiving socket. Sequentiality is secured per socket by only allowing
buffers to be dequeued inside the socket lock. Since there may be multiple
simultaneous readers of the queue, we use a 'filter' parameter to reduce
the risk that they peek the same buffer from the queue, hence also
reducing the risk of contention on the receiving socket locks.
This solves the sequentiality problem, and seems to cause no measurable
performance degradation.
A nice side effect of this change is that lock handling in the functions
tipc_rcv() and tipc_bcast_rcv() now becomes uniform, something that
will enable future simplifications of those functions.
Reviewed-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Jon Maloy <jon.maloy@ericsson.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2015-02-05 08:36:41 -05:00
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2015-07-30 18:24:23 -04:00
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2006-01-02 19:04:38 +01:00
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