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This merge pulls the XFS master branch into the latest Linus master.
This results in a merge conflict whose best fix is not obvious.
I manually fixed the conflict, in "fs/xfs/xfs_iget.c".
Dave Chinner had done work that resulted in RCU freeing of inodes
separate from what Nick Piggin had done, and their results differed
slightly in xfs_inode_free(). The fix updates Nick's call_rcu()
with the use of VFS_I(), while incorporating needed updates to some
XFS inode fields implemented in Dave's series. Dave's RCU callback
function has also been removed.
Signed-off-by: Alex Elder <aelder@sgi.com>
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* 'for-2.6.38' of git://git.kernel.org/pub/scm/linux/kernel/git/tj/wq: (33 commits)
usb: don't use flush_scheduled_work()
speedtch: don't abuse struct delayed_work
media/video: don't use flush_scheduled_work()
media/video: explicitly flush request_module work
ioc4: use static work_struct for ioc4_load_modules()
init: don't call flush_scheduled_work() from do_initcalls()
s390: don't use flush_scheduled_work()
rtc: don't use flush_scheduled_work()
mmc: update workqueue usages
mfd: update workqueue usages
dvb: don't use flush_scheduled_work()
leds-wm8350: don't use flush_scheduled_work()
mISDN: don't use flush_scheduled_work()
macintosh/ams: don't use flush_scheduled_work()
vmwgfx: don't use flush_scheduled_work()
tpm: don't use flush_scheduled_work()
sonypi: don't use flush_scheduled_work()
hvsi: don't use flush_scheduled_work()
xen: don't use flush_scheduled_work()
gdrom: don't use flush_scheduled_work()
...
Fixed up trivial conflict in drivers/media/video/bt8xx/bttv-input.c
as per Tejun.
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This simple implementation just checks for no ACLs on the inode, and
if so, then the rcu-walk may proceed, otherwise fail it.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
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Signed-off-by: Nick Piggin <npiggin@kernel.dk>
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RCU free the struct inode. This will allow:
- Subsequent store-free path walking patch. The inode must be consulted for
permissions when walking, so an RCU inode reference is a must.
- sb_inode_list_lock to be moved inside i_lock because sb list walkers who want
to take i_lock no longer need to take sb_inode_list_lock to walk the list in
the first place. This will simplify and optimize locking.
- Could remove some nested trylock loops in dcache code
- Could potentially simplify things a bit in VM land. Do not need to take the
page lock to follow page->mapping.
The downsides of this is the performance cost of using RCU. In a simple
creat/unlink microbenchmark, performance drops by about 10% due to inability to
reuse cache-hot slab objects. As iterations increase and RCU freeing starts
kicking over, this increases to about 20%.
In cases where inode lifetimes are longer (ie. many inodes may be allocated
during the average life span of a single inode), a lot of this cache reuse is
not applicable, so the regression caused by this patch is smaller.
The cache-hot regression could largely be avoided by using SLAB_DESTROY_BY_RCU,
however this adds some complexity to list walking and store-free path walking,
so I prefer to implement this at a later date, if it is shown to be a win in
real situations. I haven't found a regression in any non-micro benchmark so I
doubt it will be a problem.
Signed-off-by: Nick Piggin <npiggin@kernel.dk>
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The only thing that the grant lock remains to protect is the grant head
manipulations when adding or removing space from the log. These calculations
are already based on atomic variables, so we can already update them safely
without locks. However, the grant head manpulations require atomic multi-step
calculations to be executed, which the algorithms currently don't allow.
To make these multi-step calculations atomic, convert the algorithms to
compare-and-exchange loops on the atomic variables. That is, we sample the old
value, perform the calculation and use atomic64_cmpxchg() to attempt to update
the head with the new value. If the head has not changed since we sampled it,
it will succeed and we are done. Otherwise, we rerun the calculation again from
a new sample of the head.
This allows us to remove the grant lock from around all the grant head space
manipulations, and that effectively removes the grant lock from the log
completely. Hence we can remove the grant lock completely from the log at this
point.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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The log grant ticket wait queues are currently protected by the log
grant lock. However, the queues are functionally independent from
each other, and operations on them only require serialisation
against other queue operations now that all of the other log
variables they use are atomic values.
Hence, we can make them independent of the grant lock by introducing
new locks just to protect the lists operations. because the lists
are independent, we can use a lock per list and ensure that reserve
and write head queuing do not contend.
To ensure forced shutdowns work correctly in conjunction with the
new fast paths, ensure that we check whether the log has been shut
down in the grant functions once we hold the relevant spin locks but
before we go to sleep. This is needed to co-ordinate correctly with
the wakeups that are issued on the ticket queues so we don't leave
any processes sleeping on the queues during a shutdown.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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cancel_delayed_work_sync()
cancel_rearming_delayed_work[queue]() has been superceded by
cancel_delayed_work_sync() quite some time ago. Convert all the
in-kernel users. The conversions are completely equivalent and
trivial.
Signed-off-by: Tejun Heo <tj@kernel.org>
Acked-by: "David S. Miller" <davem@davemloft.net>
Acked-by: Greg Kroah-Hartman <gregkh@suse.de>
Acked-by: Evgeniy Polyakov <zbr@ioremap.net>
Cc: Jeff Garzik <jgarzik@pobox.com>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: Mauro Carvalho Chehab <mchehab@infradead.org>
Cc: netdev@vger.kernel.org
Cc: Anton Vorontsov <cbou@mail.ru>
Cc: David Woodhouse <dwmw2@infradead.org>
Cc: "J. Bruce Fields" <bfields@fieldses.org>
Cc: Neil Brown <neilb@suse.de>
Cc: Alex Elder <aelder@sgi.com>
Cc: xfs-masters@oss.sgi.com
Cc: Christoph Lameter <cl@linux-foundation.org>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Cc: Andrew Morton <akpm@linux-foundation.org>
Cc: netfilter-devel@vger.kernel.org
Cc: Trond Myklebust <Trond.Myklebust@netapp.com>
Cc: linux-nfs@vger.kernel.org
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Now that we don't mark VFS inodes dirty anymore for internal
timestamp changes, but rely on the transaction subsystem to push
them out, we need to explicitly log the source inode in rename after
updating it's timestamps to make sure the changes actually get
forced out by sync/fsync or an AIL push.
We already account for the fourth inode in the log reservation, as a
rename of directories needs to update the nlink field, so just
adding the xfs_trans_log_inode call is enough.
This fixes the xfsqa 065 regression introduced by:
"xfs: don't use vfs writeback for pure metadata modifications"
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Alex Elder <aelder@sgi.com>
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Convert the log grant heads to atomic64_t types in preparation for
converting the accounting algorithms to atomic operations. his patch
just converts the variables; the algorithmic changes are in a
separate patch for clarity.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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log->l_tail_lsn is currently protected by the log grant lock. The
lock is only needed for serialising readers against writers, so we
don't really need the lock if we make the l_tail_lsn variable an
atomic. Converting the l_tail_lsn variable to an atomic64_t means we
can start to peel back the grant lock from various operations.
Also, provide functions to safely crack an atomic LSN variable into
it's component pieces and to recombined the components into an
atomic variable. Use them where appropriate.
This also removes the need for explicitly holding a spinlock to read
the l_tail_lsn on 32 bit platforms.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
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log->l_last_sync_lsn is updated in only one critical spot - log
buffer Io completion - and is protected by the grant lock here. This
requires the grant lock to be taken for every log buffer IO
completion. Converting the l_last_sync_lsn variable to an atomic64_t
means that we do not need to take the grant lock in log buffer IO
completion to update it.
This also removes the need for explicitly holding a spinlock to read
the l_last_sync_lsn on 32 bit platforms.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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The xlog_grant_push_ail() currently takes the grant lock internally to sample
the tail lsn, last sync lsn and the reserve grant head. Most of the callers
already hold the grant lock but have to drop it before calling
xlog_grant_push_ail(). This is a left over from when the AIL tail pushing was
done in line and hence xlog_grant_push_ail had to drop the grant lock. AIL push
is now done in another thread and hence we can safely hold the grant lock over
the entire xlog_grant_push_ail call.
Push the grant lock outside of xlog_grant_push_ail() to simplify the locking
and synchronisation needed for tail pushing. This will reduce traffic on the
grant lock by itself, but this is only one step in preparing for the complete
removal of the grant lock.
While there, clean up the formatting of xlog_grant_push_ail() to match the
rest of the XFS code.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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The log grant queues are one of the few places left using sv_t
constructs for waiting. Given we are touching this code, we should
convert them to plain wait queues. While there, convert all the
other sv_t users in the log code as well.
Seeing as this removes the last users of the sv_t type, remove the
header file defining the wrapper and the fragments that still
reference it.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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Prepare for switching the grant heads to atomic variables by
combining the two 32 bit values that make up the grant head into a
single 64 bit variable. Provide wrapper functions to combine and
split the grant heads appropriately for calculations and use them as
necessary.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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The log grant space calculations are repeated for both write and
reserve grant heads. To make it simpler to convert the calculations
toa different algorithm, factor them so both the gratn heads use the
same calculation functions. Once this is done we can drop the
wrappers that are used in only a couple of place to update both
grant heads at once as they don't provide any particular value.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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Factor repeated debug code out of grant head manipulation functions into a
separate function. This removes ifdef DEBUG spagetti from the code and makes
the code easier to follow.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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The grant write and reserve queues use a roll-your-own double linked
list, so convert it to a standard list_head structure and convert
all the list traversals to use list_for_each_entry(). We can also
get rid of the XLOG_TIC_IN_Q flag as we can use the list_empty()
check to tell if the ticket is in a list or not.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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We now have two copies of AIL delete operations that are mostly
duplicate functionality. The single log item deletes can be
implemented via the bulk updates by turning xfs_trans_ail_delete()
into a simple wrapper. This removes all the duplicate delete
functionality and associated helpers.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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We now have two copies of AIL insert operations that are mostly
duplicate functionality. The single log item updates can be
implemented via the bulk updates by turning xfs_trans_ail_update()
into a simple wrapper. This removes all the duplicate insert
functionality and associated helpers.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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When inode buffer IO completes, usually all of the inodes are removed from the
AIL. This involves processing them one at a time and taking the AIL lock once
for every inode. When all CPUs are processing inode IO completions, this causes
excessive amount sof contention on the AIL lock.
Instead, change the way we process inode IO completion in the buffer
IO done callback. Allow the inode IO done callback to walk the list
of IO done callbacks and pull all the inodes off the buffer in one
go and then process them as a batch.
Once all the inodes for removal are collected, take the AIL lock
once and do a bulk removal operation to minimise traffic on the AIL
lock.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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To allow buffer iodone callbacks to consume multiple items off the
callback list, first we need to convert the xfs_buf_do_callbacks()
to consume items and always pull the next item from the head of the
list.
The means the item list walk is never dependent on knowing the
next item on the list and hence allows callbacks to remove items
from the list as well. This allows callbacks to do bulk operations
by scanning the list for identical callbacks, consuming them all
and then processing them in bulk, negating the need for multiple
callbacks of that type.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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The xfaild often tries to rest to wait for congestion to pass of for
IO to complete, but is regularly woken in tail-pushing situations.
In severe cases, the xfsaild is getting woken tens of thousands of
times a second. Reduce the number needless wakeups by only waking
the xfsaild if the new target is larger than the old one. Further
make short sleeps uninterruptible as they occur when the xfsaild has
decided it needs to back off to allow some IO to complete and being
woken early is counter-productive.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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When inserting items into the AIL from the transaction committed
callbacks, we take the AIL lock for every single item that is to be
inserted. For a CIL checkpoint commit, this can be tens of thousands
of individual inserts, yet almost all of the items will be inserted
at the same point in the AIL because they have the same index.
To reduce the overhead and contention on the AIL lock for such
operations, introduce a "bulk insert" operation which allows a list
of log items with the same LSN to be inserted in a single operation
via a list splice. To do this, we need to pre-sort the log items
being committed into a temporary list for insertion.
The complexity is that not every log item will end up with the same
LSN, and not every item is actually inserted into the AIL. Items
that don't match the commit LSN will be inserted and unpinned as per
the current one-at-a-time method (relatively rare), while items that
are not to be inserted will be unpinned and freed immediately. Items
that are to be inserted at the given commit lsn are placed in a
temporary array and inserted into the AIL in bulk each time the
array fills up.
As a result of this, we trade off AIL hold time for a significant
reduction in traffic. lock_stat output shows that the worst case
hold time is unchanged, but contention from AIL inserts drops by an
order of magnitude and the number of lock traversal decreases
significantly.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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xfs_ail_delete() has a needlessly complex interface. It returns the log item
that was passed in for deletion (which the callers then assert is identical to
the one passed in), and callers of xfs_ail_delete() still need to invalidate
current traversal cursors.
Make xfs_ail_delete() return void, move the cursor invalidation inside it, and
clean up the callers just to use the log item pointer they passed in.
While cleaning up, remove the messy and unnecessary "/* ARGUSED */" comments
around all these functions.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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EFI/EFD interactions are protected from races by the AIL lock. They
are the only type of log items that require the the AIL lock to
serialise internal state, so they need to be separated from the AIL
lock before we can do bulk insert operations on the AIL.
To acheive this, convert the counter of the number of extents in the
EFI to an atomic so it can be safely manipulated by EFD processing
without locks. Also, convert the EFI state flag manipulations to use
atomic bit operations so no locks are needed to record state
changes. Finally, use the state bits to determine when it is safe to
free the EFI and clean up the code to do this neatly.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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XFS_EFI_CANCELED has not been set in the code base since
xfs_efi_cancel() was removed back in 2006 by commit
065d312e15902976d256ddaf396a7950ec0350a8 ("[XFS] Remove unused
iop_abort log item operation), and even then xfs_efi_cancel() was
never called. I haven't tracked it back further than that (beyond
git history), but it indicates that the handling of EFIs in
cancelled transactions has been broken for a long time.
Basically, when we get an IOP_UNPIN(lip, 1); call from
xfs_trans_uncommit() (i.e. remove == 1), if we don't free the log
item descriptor we leak it. Fix the behviour to be correct and kill
the XFS_EFI_CANCELED flag.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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Now that the buffer reclaim infrastructure can handle different reclaim
priorities for different types of buffers, reconnect the hooks in the
XFS code that has been sitting dormant since it was ported to Linux. This
should finally give use reclaim prioritisation that is on a par with the
functionality that Irix provided XFS 15 years ago.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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Introduce a per-buftarg LRU for memory reclaim to operate on. This
is the last piece we need to put in place so that we can fully
control the buffer lifecycle. This allows XFS to be responsibile for
maintaining the working set of buffers under memory pressure instead
of relying on the VM reclaim not to take pages we need out from
underneath us.
The implementation introduces a b_lru_ref counter into the buffer.
This is currently set to 1 whenever the buffer is referenced and so is used to
determine if the buffer should be added to the LRU or not when freed.
Effectively it allows lazy LRU initialisation of the buffer so we do not need
to touch the LRU list and locks in xfs_buf_find().
Instead, when the buffer is being released and we drop the last
reference to it, we check the b_lru_ref count and if it is none zero
we re-add the buffer reference and add the inode to the LRU. The
b_lru_ref counter is decremented by the shrinker, and whenever the
shrinker comes across a buffer with a zero b_lru_ref counter, if
released the LRU reference on the buffer. In the absence of a lookup
race, this will result in the buffer being freed.
This counting mechanism is used instead of a reference flag so that
it is simple to re-introduce buffer-type specific reclaim reference
counts to prioritise reclaim more effectively. We still have all
those hooks in the XFS code, so this will provide the infrastructure
to re-implement that functionality.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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Recent tests writing lots of small files showed the flusher thread
being CPU bound and taking a long time to do allocations on a debug
kernel. perf showed this as the prime reason:
samples pcnt function DSO
_______ _____ ___________________________ _________________
224648.00 36.8% xfs_error_test [kernel.kallsyms]
86045.00 14.1% xfs_btree_check_sblock [kernel.kallsyms]
39778.00 6.5% prandom32 [kernel.kallsyms]
37436.00 6.1% xfs_btree_increment [kernel.kallsyms]
29278.00 4.8% xfs_btree_get_rec [kernel.kallsyms]
27717.00 4.5% random32 [kernel.kallsyms]
Walking btree blocks during allocation checking them requires each
block (a cache hit, so no I/O) call xfs_error_test(), which then
does a random32() call as the first operation. IOWs, ~50% of the
CPU is being consumed just testing whether we need to inject an
error, even though error injection is not active.
Kill this overhead when error injection is not active by adding a
global counter of active error traps and only calling into
xfs_error_test when fault injection is active.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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When an inode has been marked stale because the cluster is being
freed, we don't want to (re-)insert this inode into the AIL. There
is a race condition where the cluster buffer may be unpinned before
the inode is inserted into the AIL during transaction committed
processing. If the buffer is unpinned before the inode item has been
committed and inserted, then it is possible for the buffer to be
released and hence processthe stale inode callbacks before the inode
is inserted into the AIL.
In this case, we then insert a clean, stale inode into the AIL which
will never get removed by an IO completion. It will, however, get
reclaimed and that triggers an assert in xfs_inode_free()
complaining about freeing an inode still in the AIL.
This race can be avoided by not moving stale inodes forward in the AIL
during transaction commit completion processing. This closes the
race condition by ensuring we never insert clean stale inodes into
the AIL. It is safe to do this because a dirty stale inode, by
definition, must already be in the AIL.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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There is an assumption in the parts of XFS that flushing a dirty
file will make all the delayed allocation blocks disappear from an
inode. That is, that after calling xfs_flush_pages() then
ip->i_delayed_blks will be zero.
This is an invalid assumption as we may have specualtive
preallocation beyond EOF and they are recorded in
ip->i_delayed_blks. A flush of the dirty pages of an inode will not
change the state of these blocks beyond EOF, so a non-zero
deeelalloc block count after a flush is valid.
The bmap code has an invalid ASSERT() that needs to be removed, and
the swapext code has a bug in that while it swaps the data forks
around, it fails to swap the i_delayed_blks counter associated with
the fork and hence can get the block accounting wrong.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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As reported by Nick Piggin, XFS is suffering from long pauses under
highly concurrent workloads when hosted on ramdisks. The problem is
that an inode buffer is stuck in the pinned state in memory and as a
result either the inode buffer or one of the inodes within the
buffer is stopping the tail of the log from being moved forward.
The system remains in this state until a periodic log force issued
by xfssyncd causes the buffer to be unpinned. The main problem is
that these are stale buffers, and are hence held locked until the
transaction/checkpoint that marked them state has been committed to
disk. When the filesystem gets into this state, only the xfssyncd
can cause the async transactions to be committed to disk and hence
unpin the inode buffer.
This problem was encountered when scaling the busy extent list, but
only the blocking lock interface was fixed to solve the problem.
Extend the same fix to the buffer trylock operations - if we fail to
lock a pinned, stale buffer, then force the log immediately so that
when the next attempt to lock it comes around, it will have been
unpinned.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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Since the move to the new truncate sequence we call xfs_setattr to
truncate down excessively instanciated blocks. As shown by the testcase
in kernel.org BZ #22452 that doesn't work too well. Due to the confusion
of the internal inode size, and the VFS inode i_size it zeroes data that
it shouldn't.
But full blown truncate seems like overkill here. We only instanciate
delayed allocations in the write path, and given that we never released
the iolock we can't have converted them to real allocations yet either.
The only nasty case is pre-existing preallocation which we need to skip.
We already do this for page discard during writeback, so make the delayed
allocation block punching a generic function and call it from the failed
write path as well as xfs_aops_discard_page. The callers are
responsible for ensuring that partial blocks are not truncated away,
and that they hold the ilock.
Based on a fix originally from Christoph Hellwig. This version used
filesystem blocks as the range unit.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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Before we introduce per-buftarg LRU lists, split the shrinker
implementation into per-buftarg shrinker callbacks. At the moment
we wake all the xfsbufds to run the delayed write queues to free
the dirty buffers and make their pages available for reclaim.
However, with an LRU, we want to be able to free clean, unused
buffers as well, so we need to separate the xfsbufd from the
shrinker callbacks.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Alex Elder <aelder@sgi.com>
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now that we are using RCU protection for the inode cache lookups,
the lock is only needed on the modification side. Hence it is not
necessary for the lock to be a rwlock as there are no read side
holders anymore. Convert it to a spin lock to reflect it's exclusive
nature.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Alex Elder <aelder@sgi.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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With delayed logging greatly increasing the sustained parallelism of inode
operations, the inode cache locking is showing significant read vs write
contention when inode reclaim runs at the same time as lookups. There is
also a lot more write lock acquistions than there are read locks (4:1 ratio)
so the read locking is not really buying us much in the way of parallelism.
To avoid the read vs write contention, change the cache to use RCU locking on
the read side. To avoid needing to RCU free every single inode, use the built
in slab RCU freeing mechanism. This requires us to be able to detect lookups of
freed inodes, so enѕure that ever freed inode has an inode number of zero and
the XFS_IRECLAIM flag set. We already check the XFS_IRECLAIM flag in cache hit
lookup path, but also add a check for a zero inode number as well.
We canthen convert all the read locking lockups to use RCU read side locking
and hence remove all read side locking.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Alex Elder <aelder@sgi.com>
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Introduce RCU freeing of XFS inodes so that we can convert lookup
traversals to use rcu_read_lock() protection. This patch only
introduces the RCU freeing to minimise the potential conflicts with
mainline if this is merged into mainline via a VFS patchset. It
abuses the i_dentry list for the RCU callback structure because the
VFS patches make this a union so it is safe to use like this and
simplifies and merge issues.
This patch uses basic RCU freeing rather than SLAB_DESTROY_BY_RCU.
The later lookup patches need the same "found free inode" protection
regardless of the RCU freeing method used, so once again the RCU
freeing method can be dealt with apprpriately at merge time without
affecting any other code.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
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A long standing problem for streaming writeѕ through the NFS server
has been that the NFS server opens and closes file descriptors on an
inode for every write. The result of this behaviour is that the
->release() function is called on every close and that results in
XFS truncating speculative preallocation beyond the EOF. This has
an adverse effect on file layout when multiple files are being
written at the same time - they interleave their extents and can
result in severe fragmentation.
To avoid this problem, keep track of ->release calls made on a dirty
inode. For most cases, an inode is only going to be opened once for
writing and then closed again during it's lifetime in cache. Hence
if there are multiple ->release calls when the inode is dirty, there
is a good chance that the inode is being accessed by the NFS server.
Hence set a flag the first time ->release is called while there are
delalloc blocks still outstanding on the inode.
If this flag is set when ->release is next called, then do no
truncate away the speculative preallocation - leave it there so that
subsequent writes do not need to reallocate the delalloc space. This
will prevent interleaving of extents of different inodes written
concurrently to the same AG.
If we get this wrong, it is not a big deal as we truncate
speculative allocation beyond EOF anyway in xfs_inactive() when the
inode is thrown out of the cache.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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Currently the size of the speculative preallocation during delayed
allocation is fixed by either the allocsize mount option of a
default size. We are seeing a lot of cases where we need to
recommend using the allocsize mount option to prevent fragmentation
when buffered writes land in the same AG.
Rather than using a fixed preallocation size by default (up to 64k),
make it dynamic by basing it on the current inode size. That way the
EOF preallocation will increase as the file size increases. Hence
for streaming writes we are much more likely to get large
preallocations exactly when we need it to reduce fragementation.
For default settings, the size of the initial extents is determined
by the number of parallel writers and the amount of memory in the
machine. For 4GB RAM and 4 concurrent 32GB file writes:
EXT: FILE-OFFSET BLOCK-RANGE AG AG-OFFSET TOTAL
0: [0..1048575]: 1048672..2097247 0 (1048672..2097247) 1048576
1: [1048576..2097151]: 5242976..6291551 0 (5242976..6291551) 1048576
2: [2097152..4194303]: 12583008..14680159 0 (12583008..14680159) 2097152
3: [4194304..8388607]: 25165920..29360223 0 (25165920..29360223) 4194304
4: [8388608..16777215]: 58720352..67108959 0 (58720352..67108959) 8388608
5: [16777216..33554423]: 117440584..134217791 0 (117440584..134217791) 16777208
6: [33554424..50331511]: 184549056..201326143 0 (184549056..201326143) 16777088
7: [50331512..67108599]: 251657408..268434495 0 (251657408..268434495) 16777088
and for 16 concurrent 16GB file writes:
EXT: FILE-OFFSET BLOCK-RANGE AG AG-OFFSET TOTAL
0: [0..262143]: 2490472..2752615 0 (2490472..2752615) 262144
1: [262144..524287]: 6291560..6553703 0 (6291560..6553703) 262144
2: [524288..1048575]: 13631592..14155879 0 (13631592..14155879) 524288
3: [1048576..2097151]: 30408808..31457383 0 (30408808..31457383) 1048576
4: [2097152..4194303]: 52428904..54526055 0 (52428904..54526055) 2097152
5: [4194304..8388607]: 104857704..109052007 0 (104857704..109052007) 4194304
6: [8388608..16777215]: 209715304..218103911 0 (209715304..218103911) 8388608
7: [16777216..33554423]: 452984848..469762055 0 (452984848..469762055) 16777208
Because it is hard to take back specualtive preallocation, cases
where there are large slow growing log files on a nearly full
filesystem may cause premature ENOSPC. Hence as the filesystem nears
full, the maximum dynamic prealloc size іs reduced according to this
table (based on 4k block size):
freespace max prealloc size
>5% full extent (8GB)
4-5% 2GB (8GB >> 2)
3-4% 1GB (8GB >> 3)
2-3% 512MB (8GB >> 4)
1-2% 256MB (8GB >> 5)
<1% 128MB (8GB >> 6)
This should reduce the amount of space held in speculative
preallocation for such cases.
The allocsize mount option turns off the dynamic behaviour and fixes
the prealloc size to whatever the mount option specifies. i.e. the
behaviour is unchanged.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
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When listing attributes, we are doiing memory allocations under the
inode ilock using only KM_SLEEP. This allows memory allocation to
recurse back into the filesystem and do writeback, which may the
ilock we already hold on the current inode. THis will deadlock.
Hence use KM_NOFS for such allocations outside of transaction
context to ensure that reclaim recursion does not occur.
Reported-by: Nick Piggin <npiggin@gmail.com>
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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The XFS iolock needs to be re-initialised to a new lock class before
it enters reclaim to prevent lockdep false positives. Unfortunately,
this is not sufficient protection as inodes in the XFS_IRECLAIMABLE
state can be recycled and not re-initialised before being reused.
We need to re-initialise the lock state when transfering out of
XFS_IRECLAIMABLE state to XFS_INEW, but we need to keep the same
class as if the inode was just allocated. Hence we need a specific
lockdep class variable for the iolock so that both initialisations
use the same class.
While there, add a specific class for inodes in the reclaim state so
that it is easy to tell from lockdep reports what state the inode
was in that generated the report.
Signed-off-by: Dave Chinner <dchinner@redhat.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
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Add a new xfs_alloc_find_best_extent that does a forward/backward
search in the allocation btree. That code previously was existed
two times in xfs_alloc_ag_vextent_near, once for each search
direction.
Based on an earlier patch from Dave Chinner.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Alex Elder <aelder@sgi.com>
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Use a goto label to consolidate all block not found cases, and add a
tracepoint for them. Also clean up a few whitespace issues.
Based on an earlier patch from Dave Chinner.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Alex Elder <aelder@sgi.com>
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Move the buffer locking into the callers as they need to do it
wether they call xfs_map_at_offset or not. Remove the b_bdev
assignment, which is already done by get_blocks. Remove the
duplicate extent type asserts in xfs_convert_page just before
calling xfs_map_at_offset.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Alex Elder <aelder@sgi.com>
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After the last patches the code for overwrites is the same as for
delayed and unwritten extents except that it doesn't need to call
xfs_map_at_offset. Take care of that fact to simplify
xfs_vm_writepage.
The buffer loop now first checks the type of buffer and checks/sets
the ioend type, or continues to the next buffer if it's not
interesting to us. Only after that we validate the iomap and
perform the block mapping if needed, all in common code for the
cases where we have to do work.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Alex Elder <aelder@sgi.com>
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The all_bh flag is always set when entering the page clustering
machinery with a regular written extent, which means the check for
it is superflous.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Alex Elder <aelder@sgi.com>
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xfs_map_blocks always calls xfs_bmapi with the XFS_BMAPI_ENTIRE
entire flag, which tells it to not cap the extent at the passed in
size, but just treat the size as an minimum to map. This means
xfs_probe_cluster is entirely useless as we'll always get the whole
extent back anyway.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Alex Elder <aelder@sgi.com>
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No need to lock the extent map exclusive when performing an
overwrite, we know the extent map must already have been loaded by
get_blocks. Apply the non-blocking inode semantics to all mapping
types instead of just delayed allocations. Remove the handling of
not yet allocated blocks for the IO_UNWRITTEN case - if an extent is
marked as unwritten allocated in the buffer it must already have an
extent on disk.
Add asserts to verify all the assumptions above in debug builds.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Alex Elder <aelder@sgi.com>
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Opencode the xfs_iomap code in it's two callers. The overlap of
passed flags already was minimal and will be further reduced in the
next patch.
As a side effect the BMAPI_* flags for xfs_bmapi and the IO_* flags
for I/O end processing are merged into a single set of flags, which
should be a bit more descriptive of the operation we perform.
Also improve the tracing by giving each caller it's own type set of
tracepoints.
Signed-off-by: Christoph Hellwig <hch@lst.de>
Reviewed-by: Dave Chinner <dchinner@redhat.com>
Signed-off-by: Alex Elder <aelder@sgi.com>
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