into penguin.transmeta.com:/home/penguin/torvalds/repositories/kernel/linux

into tp1.ruhr-uni-bochum.de:/home/kai/kernel/v2.5/linux-2.5.export

ips driver update ( version 5.10.13-BETA )

The lock.h header contained some hand-crafted lcoking routines from
the pre-SMP days.  In 2.5 only lock_super/unlock_super are left,
guarded by a number of completly unrelated (!) includes.

This patch moves lock_super/unlock_super to fs.h, which defined
struct super_block that is needed for those to operate it, removes
locks.h and updates all caller to not include it and add the missing,
previously nested includes where needed.

This patch implements ioctl() for getting space used by file.
I agree it's ioctl() abuse, it doesn't work on links and has
other ugly properties. Better would be to change 'struct stat'
but changing it just due to this is overkill and it will take
some time before there will be enough changes which will provoke
yet another struct stat :). So this is temporary solution...
If you don't like it, simply reject it. The function it provides
is not fundamental...

So that should be all patches. Any comments (or decision about
including/not including) welcome.
								Honza

This patch implements configurable backward compatible quota interface.
Maybe this isn't needed in 2.5 but as some people want to use patches
in 2.4 where it's necessary I have implemented it.

Implemented proper syncing of dquots - ie. also global information
about quota files are synced. We find info to sync by walking through
all superblocks...

Remove use of 'short' in parameters of functions. 'int' is used instead.

Implementation of new quota format. The code is almost the same
as in -ac versions of kernel. All the code for new format is in
quota_v2.c

Implementation of old quota format. All the code for old format is now in
quota_v1.c. Code mostly remained the same as in older kernels (just minor
changes were needed to bind it with quota interface).

This is probably the largest chunk in quota patches. It removes old quotactl interface
and implements new one. New interface should not need arch specific conversions so they
are removed. All quota interface stuff is moved to quota.c so we can
easily separate things which should be compiled even if quota is disabled (mainly
because XFS needs some interface even if standard VFS quota is disabled).
Callbacks to filesystem on quota_on() and quota_off() are implemented (needed by Ext3),
quota operations callbacks are now set in super.c on superblock initialization and
not on quota_on(). This way it starts to make sense to have callbacks on alloc_space(),
alloc_inode() etc. as filesystem can override them on read_super(). This will be used
later for implementing journalled quota.

This patch implements counting of used space in inodes in bytes.
New field i_bytes is added and used space modulo 512 is kept in
it (rest is still kept in i_blocks). Functions manipulating both
i_blocks and i_bytes are implemented (inode_add_bytes(), inode_sub_bytes()
and inode_set_bytes()). Filesystems allocating only in whole blocks
can safely ignore i_bytes field and continue using i_blocks...

This patch implements accounting of used space in bytes.

  This patch moves reporting of quota statistics from Q_GETSTATS call to
/proc/fs/quota. Also reporting of registered quota formats is added.

  This patch implements list 'quota_formats' with registered quota formats
and functions register_quota_format() and unregister_quota_format() for
manipulating the list.

This patch removes most format dependent code from dquot.c and quota.h
and puts calls of callback functions instead.

  This patch adds dq_dup_ref to struct dquot. Functions altering just usage of
quota take just this duplicated reference, inodes, quotactl() helpers take real
dq_count reference. dqput() blocks if there are some duplicated references and
put reference is last 'real one'. This way is assured that quota IO is not
done from functions altering quota usage (quota structure is written on last dqput()).

As of the last patch the inode_hashtable doesn't really need to be
indexed by i_ino anymore, the only reason we still have to keep the
hashvalue and i_ino identical is because of insert_inode_hash.

If at some point a FS specific getattr method is implemented it will be
possible to completely remove any use of i_ino by the VFS.

This patch starts taking i_ino dependencies out of the VFS. The FS
provided test and set callbacks become responsible for testing and
setting inode->i_ino.

Because most filesystems are based on 32-bit unique inode numbers
several functions are duplicated to keep iget_locked as a fast path. We
can avoid unnecessary pointer dereferences and function calls for this
specific case.

Now that we have no more users of iget4 we can kill the function and the
associated read_inode2 callback (i.e. the 'reiserfs specific hack').

Document iget5_locked as the replacement for iget4 in filesystems/porting.

Convert existing filesystems (Coda/NFS/ReiserFS) that currently use
iget4 to iget5_locked.

Now we introduce iget_locked and iget5_locked. These are similar to
iget, but return a locked inode and read_inode has not been called. So
the FS has to call read_inode to initialize the inode and then unlock
it with unlock_new_inode().

This patch is based on the icreate patch from the XFS group, i.e.
it is pretty much identical except for function naming.

Fix a race in iget4. The fs specific data that is used to find an inode
should be initialized while still holding the inode lock.

It adds a 'set' callback function that should be a non-blocking FS
provided function which initializes the private parts of the inode so
that the 'test' callback function can correctly match new inodes.

Touches all filesystems that use iget4 (Coda/NFS/ReiserFS).

into home.transmeta.com:/home/torvalds/v2.5/linux

	- fix copy_{to,from}_user error handling, thanks to Rusty to pointing this out on lkml

  - fix copy_{to,from}_user error handling, thanks to Rusty for
    pointing this out on lkml

Removal of PG_launder.

It's not obvious (to me) why this ever existed.  If it's to prevent
deadlocks then I'd like to know who was performing __GFP_FS allocations
while holding a page lock?

But in 2.5, the only memory allocations which are performed when the
caller holds PG_writeback against an unsubmitted page are those which
occur inside submit_bh().  There will be no __GFS_FS allocations in
that call chain.

Removing PG_launder means that memory allocators can block on any
PageWriteback() page at all, which reduces the risk of very long list
walks inside pagemap_lru_lock in shrink_cache().

The ext3-no-steal patch has exposed a long-standing race in ext3.  It
has been there all the time in 2.4, but never triggered until some
timing change in the ext3-no-steal patch exposed it.  The race was not
present in 2.2 because 2.2's bdflush runs inside lock_kernel().

The problem is that when ext3 is shuffling a buffer between journalling
lists there is a small window where the buffer is marked BH_dirty.
Aonther CPU can grab it, mark it clean and write it out.  Then ext3
puts the buffer onto a list of buffers which are expected to be dirty,
and gets confused later on when the buffer turns out to be clean.

The patch from Stephen records the expected dirtiness of the buffer in
a local variable, so BH_dirty is not transiently set while ext3
shuffles.

Patch from sct fixes a long-standing (I did it!) and rather complex
problem with ext3.

The problem is to do with buffers which are continually being dirtied
by an external agent.  I had code in there (for easily-triggerable
livelock avoidance) which steals the buffer from checkpoint mode and
reattaches it to the running transaction.  This violates ext3 ordering
requirements - it can permit journal space to be reclaimed before the
relevant data has really been written out.

Also, we do have to reliably get a lock on the buffer when moving it
between lists and inspecting its internal state.  Otherwise a competing
read from the underlying block device can trigger an assertion failure,
and a competing write to the underlying block device can confuse ext3
journalling state completely.

Fixes a performance problem with many-small-file writeout.

At present, files are written out via their mapping and their indirect
blocks are written out via the blockdev mapping.  As we know that
indirects are disk-adjacent to the data it is better to start I/O
against the indirects at the same time as the data.

The delalloc pathes have code in ext2_writepage() which recognises when
the target page->index was at an indirect boundary and does an explicit
hunt-and-write against the neighbouring indirect block.  Which is
ideal.  (Unless the file was dirtied seekily and the page which is next
to the indirect was not dirtied).

This patch does it the other way: when we start writeback against a
mapping, also start writeback against any dirty buffers which are
attached to mapping->private_list.  Let the elevator take care of the
rest.

The patch makes a number of tuning changes to the writeback path in
fs-writeback.c.  This is very fiddly code: getting the throughput
tuned, getting the data-integrity "sync" operations right, avoiding
most of the livelock opportunities, getting the `kupdate' function
working efficiently, keeping it all least somewhat comprehensible.

An important intent here is to ensure that metadata blocks for inodes
are marked dirty before writeback starts working the blockdev mapping,
so all the inode blocks are efficiently written back.

The patch removes try_to_writeback_unused_inodes(), which became
unreferenced in vm-writeback.patch.

The patch has a tweak in ext2_put_inode() to prevent ext2 from
incorrectly droppping its preallocation window in response to a random
iput().


Generally, many-small-file writeout is a lot faster than 2.5.7 (which
is linux-before-I-futzed-with-it).  The workload which was optimised was

	tar xfz /nfs/mountpoint/linux-2.4.18.tar.gz ; sync

on mem=128M and mem=2048M.

With these patches, 2.5.15 is completing in about 2/3 of the time of
2.5.7.  But it is only a shade faster than 2.4.19-pre7.  Why is 2.5.7
so much slower than 2.4.19?  Not sure yet.

Heavy dbench loads (dbench 32 on mem=128M) are slightly faster than
2.5.7 and significantly slower than 2.4.19.  It appears that the cause
is poor read throughput at the later stages of the run.  Because there
are background writeback threads operating at the same time.

The 2.4.19-pre8 write scheduling manages to stop writeback during the
latter stages of the dbench run in a way which I haven't been able to
sanely emulate yet.  It may not be desirable to do this anyway - it's
optimising for the case where the files are about to be deleted.  But
it would be good to find a way of "pausing" the writeback for a few
seconds to allow readers to get an interval of decent bandwidth.

tiobench throughput is basically the same across all recent kernels.
CPU load on writes is down maybe 30% in 2.5.15.

When ext2 creates a new inode, perform an asynchronous preread against
its backing block.

Without this patch, many-file writeout gets stalled by having to read
many individual inode table blocks in the middle of writeback.

It's worth about a 20% gain in writeback bandwidth for the many-file
writeback case.

ext3 already reads the inode's backing block in
ext3_new_inode->ext3_mark_inode_dirty, so no change is needed there.

A backport to 2.4 would make sense.

Tune up the VM-based writeback a bit.

- Always use the multipage clustered-writeback function from within
  shrink_cache(), even if the page's mapping has a NULL ->vm_writeback().  So
  clustered writeback is turned on for all address_spaces, not just ext2.

  Subtle effect of this change: it is now the case that *all* writeback
  proceeds along the mapping->dirty_pages list.  The orderedness of the page
  LRUs no longer has an impact on disk scheduling.  So we only have one list
  to keep well-sorted rather than two, and churning pages around on the LRU
  will no longer damage write bandwidth - it's all up to the filesystem.

- Decrease the clustered writeback from 1024 pages(!) to 32 pages.

  (1024 was a leftover from when this code was always dispatching writeback
  to a pdflush thread).

- Fix wakeup_bdflush() so that it actually does write something (duh).

  do_wp_page() needs to call balance_dirty_pages_ratelimited(), so we
  throttle mmap page-dirtiers in the same way as write(2) page-dirtiers.
  This may make wakeup_bdflush() obsolete, but it doesn't hurt.

- Converts generic_vm_writeback() to directly call ->writeback_mapping(),
  rather that going through writeback_single_inode().  This prevents memory
  allocators from blocking on the inode's I_LOCK.  But it does mean that two
  processes can be writing pages from the same mapping at the same time.  If
  filesystems care about this (for layout reasons) then they should serialise
  in their ->writeback_mapping a_op.

  This means that memory-allocators will writeback only pages, not pages
  and inodes.  There are no locks in that writeback path (except for request
  queue exhaustion).  Reduces memory allocation latency.

- Implement new background_writeback function, which when kicked off
  will perform writeback until dirty memory falls below the background
  threshold.

- Put written-back pages onto the remote end of the page LRU.  It
  does this in the slow-and-stupid way at present.  pagemap_lru_lock
  stress-relief is planned...

- Remove the funny writeback_unused_inodes() stuff from prune_icache().
  Writeback from wakeup_bdflush() and the `kupdate' function now just
  naturally cleanses the oldest inodes so we don't need to do anything
  there.

- Dirty memory balancing is still using magic numbers: "after you
  dirtied your 1,000th page, go write 1,500".  Obviously, this needs
  more work.

Use the pdflush exclusion infrastructure to ensure that only one
pdlfush thread is ever performing writeback against a particular
request_queue.

This works rather well.  It requires a lot of activity against a lot of
disks to cause more pdflush threads to start up.  Possibly the
thread-creation logic is a little weak: it starts more threads when a
pdflush thread goes back to sleep.  It may be better to start new
threads within pdlfush_operation().

All non-request_queue-backed address_spaces share the global
default_backing_dev_info structure.  So at present only a single
pdflush instance will be available for background writeback of *all*
NFS filesystems (for example).

If there is benefit in concurrent background writeback for multiple NFS
mounts then NFS would need to create per-mount backing_dev_info
structures and install those into new inode's address_spaces in some
manner.

Collision avoidance for pdflush threads.

Turns the request_queue-based `unsigned long ra_pages' into a structure
which contains ra_pages as well as a longword.

That longword is used to record the fact that a pdflush thread is
currently writing something back against this request_queue.

Avoids the situation where several pdflush threads are sleeping on the
same request_queue.

This patch provides only the infrastructure for the pdflush exclusion.
This infrastructure gets used in pdflush-single.patch

Fix the "race with umount" in __sync_list().  __sync_list() no longer
puts inodes onto a local list while writing them out.

The super_block.sb_dirty list is kept time-ordered.  Mappings which
have the "oldest" ->dirtied_when are kept at sb->s_dirty.prev.

So the time-based writeback (kupdate) can just bale out when it
encounters a not-old-enough mapping, rather than walking the entire
list.

dirtied_when is set on the *first* dirtying of a mapping.  So once the
mapping is marked dirty it strictly retains its place on s_dirty until
it reaches the oldest end and is written out.  So frequently-dirtied
mappings don't stay dirty at the head of the list for all time.

That local inode list was there for livelock avoidance.  Livelock is
instead avoided by looking at each mapping's ->dirtied_when.  If we
encounter one which was dirtied after this invokation of __sync_list(),
then just bale out - the sync functions are only required to write out
data which was dirty at the time when they were called.

Keeping the s_dirty list in time-order is the right thing to do anyway
- so all the various writeback callers always work against the oldest
data.