Failure to commit a transaction into the CIL is not handled
correctly. This currently can only happen when racing with a
shutdown and requires an explicit shutdown check, so it rare and can
be avoided. Remove the shutdown check and make the CIL commit a void
function to indicate it will always succeed, thereby removing the
incorrectly handled failure case.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Reviewed-by: NAlex Elder <aelder@sgi.com>

The extent size hint can be set to larger than an AG. This means
that the alignment process can push the range to be allocated
outside the bounds of the AG, resulting in assert failures or
corrupted bmbt records. Similarly, if the extsize is larger than the
maximum extent size supported, the alignment process will produce
extents that are too large to fit into the bmbt records, resulting
in a different type of assert/corruption failure.

Fix this by limiting extsize at the time іt is set firstly to be
less than MAXEXTLEN, then to be a maximum of half the size of the
AGs in the filesystem for non-realtime inodes. Realtime inodes do
not allocate out of AGs, so don't have to be restricted by the size
of AGs.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Reviewed-by: NAlex Elder <aelder@sgi.com>

When doing delayed allocation, if the allocation size is for a
maximally sized extent, extent size alignment can push it over this
limit. This results in an assert failure in xfs_bmbt_set_allf() as
the extent length is too large to find in the extent record.

Fix this by ensuring that we allow for space that extent size
alignment requires (up to 2 * (extsize -1) blocks as we have to
handle both head and tail alignment) when limiting the maximum size
of the extent.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Reviewed-by: NAlex Elder <aelder@sgi.com>

Delayed allocation extents can be larger than AGs, so when trying to
convert a large range we may scan every AG inside
xfs_bmap_alloc_nullfb() trying to find an AG with a size larger than
an AG. We should stop when we find the first AG with a maximum
possible allocation size. This causes excessive CPU usage when there
are lots of AGs.

The same problem occurs when doing preallocation of a range larger
than an AG.

Fix the problem by limiting real allocation lengths to the maximum
that an AG can support. This means if we have empty AGs, we'll stop
the search at the first of them. If there are no empty AGs, we'll
still scan them all, but that is a different problem....
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Reviewed-by: NAlex Elder <aelder@sgi.com>

rounddown_power_of_2() returns an undefined result when passed a
value of zero. The specualtive delayed allocation code is doing this
when the inode is zero length. Hence occasionally the preallocation
is much, much larger than is necessary (e.g. 8GB for a 270 _byte_
file). Ensure we don't even pass a zero value to this function so
the result of preallocation is always the desired size.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Reviewed-by: NAlex Elder <aelder@sgi.com>

After test 139, kmemleak shows:

unreferenced object 0xffff880078b405d8 (size 400):
  comm "xfs_io", pid 4904, jiffies 4294909383 (age 1186.728s)
  hex dump (first 32 bytes):
    60 c1 17 79 00 88 ff ff 60 c1 17 79 00 88 ff ff  `..y....`..y....
    00 00 00 00 00 00 00 00 00 00 00 00 00 00 00 00  ................
  backtrace:
    [<ffffffff81afb04d>] kmemleak_alloc+0x2d/0x60
    [<ffffffff8115c6cf>] kmem_cache_alloc+0x13f/0x2b0
    [<ffffffff814aaa97>] kmem_zone_alloc+0x77/0xf0
    [<ffffffff814aab2e>] kmem_zone_zalloc+0x1e/0x50
    [<ffffffff8147cd6b>] xfs_efi_init+0x4b/0xb0
    [<ffffffff814a4ee8>] xfs_trans_get_efi+0x58/0x90
    [<ffffffff81455fab>] xfs_bmap_finish+0x8b/0x1d0
    [<ffffffff814851b4>] xfs_itruncate_finish+0x2c4/0x5d0
    [<ffffffff814a970f>] xfs_setattr+0x8df/0xa70
    [<ffffffff814b5c7b>] xfs_vn_setattr+0x1b/0x20
    [<ffffffff8117dc00>] notify_change+0x170/0x2e0
    [<ffffffff81163bf6>] do_truncate+0x66/0xa0
    [<ffffffff81163d0b>] sys_ftruncate+0xdb/0xe0
    [<ffffffff8103a002>] system_call_fastpath+0x16/0x1b
    [<ffffffffffffffff>] 0xffffffffffffffff

The cause of the leak is that the "remove" parameter of IOP_UNPIN()
is never set when a CIL push is aborted. This means that the EFI
item is never freed if it was in the push being cancelled. The
problem is specific to delayed logging, but has uncovered a couple
of problems with the handling of IOP_UNPIN(remove).

Firstly, we cannot safely call xfs_trans_del_item() from IOP_UNPIN()
in the CIL commit failure path or the iclog write failure path
because for delayed loging we have no transaction context. Hence we
must only call xfs_trans_del_item() if the log item being unpinned
has an active log item descriptor.

Secondly, xfs_trans_uncommit() does not handle log item descriptor
freeing during the traversal of log items on a transaction. It can
reference a freed log item descriptor when unpinning an EFI item.
Hence it needs to use a safe list traversal method to allow items to
be removed from the transaction during IOP_UNPIN().
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NAlex Elder <aelder@sgi.com>

The kmemleak detector shows this after test 139:

unreferenced object 0xffff880079b88bb0 (size 264):
  comm "xfs_io", pid 4904, jiffies 4294909382 (age 276.824s)
  hex dump (first 32 bytes):
    00 00 00 00 ad 4e ad de ff ff ff ff 00 00 00 00  .....N..........
    ff ff ff ff ff ff ff ff 48 7b c9 82 ff ff ff ff  ........H{......
  backtrace:
    [<ffffffff81afb04d>] kmemleak_alloc+0x2d/0x60
    [<ffffffff8115c6cf>] kmem_cache_alloc+0x13f/0x2b0
    [<ffffffff814aaa97>] kmem_zone_alloc+0x77/0xf0
    [<ffffffff814aab2e>] kmem_zone_zalloc+0x1e/0x50
    [<ffffffff8148f394>] xlog_ticket_alloc+0x34/0x170
    [<ffffffff81494444>] xlog_cil_push+0xa4/0x3f0
    [<ffffffff81494eca>] xlog_cil_force_lsn+0x15a/0x160
    [<ffffffff814933a5>] _xfs_log_force_lsn+0x75/0x2d0
    [<ffffffff814a264d>] _xfs_trans_commit+0x2bd/0x2f0
    [<ffffffff8148bfdd>] xfs_iomap_write_allocate+0x1ad/0x350
    [<ffffffff814ac17f>] xfs_map_blocks+0x21f/0x370
    [<ffffffff814ad1b7>] xfs_vm_writepage+0x1c7/0x550
    [<ffffffff8112200a>] __writepage+0x1a/0x50
    [<ffffffff81122df2>] write_cache_pages+0x1c2/0x4c0
    [<ffffffff81123117>] generic_writepages+0x27/0x30
    [<ffffffff814aba5d>] xfs_vm_writepages+0x5d/0x80

By inspection, the leak occurs when xlog_write() returns and error
and we jump to the abort path without dropping the reference on the
active ticket.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Reviewed-by: NAlex Elder <aelder@sgi.com>

On platforms that call panic() inside their BUG() macro (m68k/sun3, and
all platforms that don't set HAVE_ARCH_BUG), compilation fails with:

| fs/xfs/support/debug.c: In function ‘xfs_cmn_err’:
| fs/xfs/support/debug.c:92: error: called object ‘panic’ is not a function

as the local variable "panic" conflicts with the "panic()" function.
Rename the local variable to resolve this.
Signed-off-by: NGeert Uytterhoeven <geert@linux-m68k.org>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Currently all filesystems except XFS implement fallocate asynchronously,
while XFS forced a commit.  Both of these are suboptimal - in case of O_SYNC
I/O we really want our allocation on disk, especially for the !KEEP_SIZE
case where we actually grow the file with user-visible zeroes.  On the
other hand always commiting the transaction is a bad idea for fast-path
uses of fallocate like for example in recent Samba versions.   Given
that block allocation is a data plane operation anyway change it from
an inode operation to a file operation so that we have the file structure
available that lets us check for O_SYNC.

This also includes moving the code around for a few of the filesystems,
and remove the already unnedded S_ISDIR checks given that we only wire
up fallocate for regular files.
Signed-off-by: NChristoph Hellwig <hch@lst.de>
Signed-off-by: NAl Viro <viro@zeniv.linux.org.uk>

Instead of various home grown checks that might need updates for new
flags just check for any bit outside the mask of the features supported
by the filesystem.  This makes the check future proof for any newly
added flag.
Signed-off-by: NChristoph Hellwig <hch@lst.de>
Signed-off-by: NAl Viro <viro@zeniv.linux.org.uk>

This patch simply allows XFS to handle the hole punching flag in fallocate
properly.  I've tested this with a little program that does a bunch of random
hole punching with FL_KEEP_SIZE and without it to make sure it does the right
thing.  Thanks,
Signed-off-by: NJosef Bacik <josef@redhat.com>
Signed-off-by: NAl Viro <viro@zeniv.linux.org.uk>

We currently have a global error message buffer in cmn_err that is
protected by a spin lock that disables interrupts.  Recently there
have been reports of NMI timeouts occurring when the console is
being flooded by SCSI error reports due to cmn_err() getting stuck
trying to print to the console while holding this lock (i.e. with
interrupts disabled). The NMI watchdog is seeing this CPU as
non-responding and so is triggering a panic.  While the trigger for
the reported case is SCSI errors, pretty much anything that spams
the kernel log could cause this to occur.

Realistically the only reason that we have the intemediate message
buffer is to prepend the correct kernel log level prefix to the log
message. The only reason we have the lock is to protect the global
message buffer and the only reason the message buffer is global is
to keep it off the stack. Hence if we can avoid needing a global
message buffer we avoid needing the lock, and we can do this with a
small amount of cleanup and some preprocessor tricks:

	1. clean up xfs_cmn_err() panic mask functionality to avoid
	   needing debug code in xfs_cmn_err()
	2. remove the couple of "!" message prefixes that still exist that
	   the existing cmn_err() code steps over.
	3. redefine CE_* levels directly to KERN_*
	4. redefine cmn_err() and friends to use printk() directly
	   via variable argument length macros.

By doing this, we can completely remove the cmn_err() code and the
lock that is causing the problems, and rely solely on printk()
serialisation to ensure that we don't get garbled messages.

A series of followup patches is really needed to clean up all the
cmn_err() calls and related messages properly, but that results in a
series that is not easily back portable to enterprise kernels. Hence
this initial fix is only to address the direct problem in the lowest
impact way possible.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Signed-off-by: NAlex Elder <aelder@sgi.com>

I received a ppc64 bug report involving xfs but the assertion was
filtered out by the console log level. Use KERN_CRIT to ensure it
makes it out.
Signed-off-by: NAnton Blanchard <anton@samba.org>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Signed-off-by: NAlex Elder <aelder@sgi.com>

In fs/xfs/xfs_trans.c::xfs_trans_unreserve_and_mod_sb() at the out:
label we have this:
	ASSERT(error = 0);
I believe a comparison was intended, not an assignment. If I'm
right, the patch below fixes that up.
Signed-off-by: NJesper Juhl <jj@chaosbits.net>
Signed-off-by: NAlex Elder <aelder@sgi.com>

If we get an IO error on a synchronous superblock write, we attach an
error release function to it so that when the last reference goes away
the release function is called and the buffer is invalidated and
unlocked. The buffer is left locked until the release function is
called so that other concurrent users of the buffer will be locked out
until the buffer error is fully processed.

Unfortunately, for the superblock buffer the filesyetm itself holds a
reference to the buffer which prevents the reference count from
dropping to zero and the release function being called. As a result,
once an IO error occurs on a sync write, the buffer will never be
unlocked and all future attempts to lock the buffer will hang.

To make matters worse, this problems is not unique to such buffers;
if there is a concurrent _xfs_buf_find() running, the lookup will grab
a reference to the buffer and then wait on the buffer lock, preventing
the reference count from ever falling to zero and hence unlocking the
buffer.

As such, the whole b_relse function implementation is broken because it
cannot rely on the buffer reference count falling to zero to unlock the
errored buffer. The synchronous write error path is the only path that
uses this callback - it is used to ensure that the synchronous waiter
gets the buffer error before the error state is cleared from the buffer
by the release function.

Given that the only sychronous buffer writes now go through xfs_bwrite
and the error path in question can only occur for a write of a dirty,
logged buffer, we can move most of the b_relse processing to happen
inline in xfs_buf_iodone_callbacks, just like a normal I/O completion.
In addition to that we make sure the error is not cleared in
xfs_buf_iodone_callbacks, so that xfs_bwrite can reliably check it.
Given that xfs_bwrite keeps the buffer locked until it has waited for
it and checked the error this allows to reliably propagate the error
to the caller, and make sure that the buffer is reliably unlocked.

Given that xfs_buf_iodone_callbacks was the only instance of the
b_relse callback we can remove it entirely.

Based on earlier patches by Dave Chinner and Ajeet Yadav.
Signed-off-by: NChristoph Hellwig <hch@lst.de>
Reported-by: NAjeet Yadav <ajeet.yadav.77@gmail.com>
Reviewed-by: NDave Chinner <dchinner@redhat.com>
Signed-off-by: NAlex Elder <aelder@sgi.com>

Allow manual discards from userspace using the FITRIM ioctl.  This is not
intended to be run during normal workloads, as the freepsace btree walks
can cause large performance degradation.
Signed-off-by: NChristoph Hellwig <hch@lst.de>
Reviewed-by: NDave Chinner <dchinner@redhat.com>
Signed-off-by: NAlex Elder <aelder@sgi.com>

To ensure the log is covered and the filesystem idles correctly, we
need to ensure that dummy transactions hit the disk and do not stay
pinned in memory.  If the superblock is pinned in memory, it can't
be flushed so the log covering cannot make progress. The result is
dependent on timing - more oftent han not we continue to issues a
log covering transaction every 36s rather than idling after ~90s.

Fix this by making the log covering transaction synchronous. To
avoid additional log force from xfssyncd, make the log covering
transaction take the place of the existing log force in the xfssyncd
background sync process.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>
Signed-off-by: NAlex Elder <aelder@sgi.com>

When two concurrent unaligned, non-overlapping direct IOs are issued
to the same block, the direct Io layer will race to zero the block.
The result is that one of the concurrent IOs will overwrite data
written by the other IO with zeros. This is demonstrated by the
xfsqa test 240.

To avoid this problem, serialise all unaligned direct IOs to an
inode with a big hammer. We need a big hammer approach as we need to
serialise AIO as well, so we can't just block writes on locks.
Hence, the big hammer is calling xfs_ioend_wait() while holding out
other unaligned direct IOs from starting.

We don't bother trying to serialised aligned vs unaligned IOs as
they are overlapping IO and the result of concurrent overlapping IOs
is undefined - the result of either IO is a valid result so we let
them race. Hence we only penalise unaligned IO, which already has a
major overhead compared to aligned IO so this isn't a major problem.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NAlex Elder <aelder@sgi.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

The buffered IO and direct IO write paths share a common set of
checks and limiting code prior to issuing the write. Factor that
into a common helper function.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NAlex Elder <aelder@sgi.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

Complete the split of the different write IO paths by splitting the
buffered IO write path out of xfs_file_aio_write(). This makes the
different mechanisms of the write patchs easier to follow.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NAlex Elder <aelder@sgi.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

The current xfs_file_aio_write code is a mess of locking shenanigans
to handle the different locking requirements of buffered and direct
IO. Start to clean this up by disentangling the direct IO path from
the mess.

This also removes the failed direct IO fallback path to buffered IO.
XFS handles all direct IO cases without needing to fall back to
buffered IO, so we can safely remove this unused path. This greatly
simplifies the logic and locking needed in the write path.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

We need to obtain the i_mutex, i_iolock and i_ilock during the read
and write paths. Add a set of wrapper functions to neatly
encapsulate the lock ordering and shared/exclusive semantics to make
the locking easier to follow and get right.

Note that this changes some of the exclusive locking serialisation in
that serialisation will occur against the i_mutex instead of the
XFS_IOLOCK_EXCL. This does not change any behaviour, and it is
arguably more efficient to use the mutex for such serialisation than
the rw_sem.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NAlex Elder <aelder@sgi.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NAlex Elder <aelder@sgi.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

xfs_file_aio_write() only returns the error from synchronous
flushing of the data and inode if error == 0. At the point where
error is being checked, it is guaranteed to be > 0. Therefore any
errors returned by the data or fsync flush will never be returned.
Fix the checks so we overwrite the current error once and only if an
error really occurred.
Signed-off-by: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NAlex Elder <aelder@sgi.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

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: NNick Piggin <npiggin@kernel.dk>

Signed-off-by: NNick Piggin <npiggin@kernel.dk>

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: NNick Piggin <npiggin@kernel.dk>

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: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

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: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

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: NTejun Heo <tj@kernel.org>
Acked-by: N"David S. Miller" <davem@davemloft.net>
Acked-by: NGreg Kroah-Hartman <gregkh@suse.de>
Acked-by: NEvgeniy 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

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: NChristoph Hellwig <hch@lst.de>
Reviewed-by: NDave Chinner <dchinner@redhat.com>
Signed-off-by: NAlex Elder <aelder@sgi.com>

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: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

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: NDave Chinner <dchinner@redhat.com>

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: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

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: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

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: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

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: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

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: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>

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: NDave Chinner <dchinner@redhat.com>
Reviewed-by: NChristoph Hellwig <hch@lst.de>