The serial layer is restructured to allow less code duplication (and
hence bug duplication) across various serial drivers.  Since ARM adds
six extra serial drivers, maintaining six copies of serial.c was not
my idea of fun.

Therefore, we've ended up with a core serial driver, which knows about
the interactions with the tty layer, and low-level hardware drivers,
which know all about the hardware.  The interface between the two is
described in "Documentation/serial/driver".

This patch completely removes the old serial.c driver and its associated
configuration options, as you requested at KS2002.  We keep a certain
amount of configuration compatibility with the per-architecture serial.h
file for the moment; this *will* be killed in the next round of patches.
The biggest user of this is x86, and since I don't have an x86 box to
test this stuff on, I think the changes are best kept separate.

This is a massive cleanup of the IRQ subsystem.  It's losely based on
Linus' original idea and DaveM's original implementation, to fold our
various irq, softirq and bh counters into the preemption counter.

with this approach it was possible:

 - to remove the 'big IRQ lock' on SMP - on which sti() and cli() relied.

 - to streamline/simplify arch/i386/kernel/irq.c significantly.

 - to simplify the softirq code.

 - to remove the preemption count increase/decrease code from the lowlevel
   IRQ assembly code.

 - to speed up schedule() a bit.

Global sti() and cli() is gone forever on SMP, there is no more globally
synchronizing irq-disabling capability.  All code that relied on sti()
and cli() and restore_flags() must use other locking mechanisms from now
on (spinlocks and __cli()/__sti()).

obviously this patch breaks massive amounts of code, so only limited
.configs are working at the moment (UP is expected to be unaffected, but
SMP will require various driver updates).

The patch was developed and tested on SMP systems, and while the code is
still a bit rough in places, the base IRQ code appears to be pretty
robust and clean.

while it boots already so the worst is over, there is lots of work left:
eg. to fix the serial layer to not use cli()/sti() and bhs ...

In the /proc/fs/jffs/* code we switch to passing number of mtd device
(as an integer) instead of messing with kdev_t (which would always be
mk_kdev(MTD_BLOCK_MAJOR, device_number) anyway).

Removed prototypes of several functions that do not exist...

->bios_param() switched from kdev_t to struct block_device *.

Caller and all instances updated.

somewhat related to the above - drivers/block/paride/* switched to
module_init()/module_exit(), pd.c taught to use LBA if disks support
it (needed for paride disks >8Gb; change is fairly trivial and I've
got 40Gb one ;-)

blk_ioctl() not exported anymore; calls moved from drivers to block_dev.c.

Horrors with open/reread_partition exclusion are starting to get fixed.

It's not the final variant, but at least we are getting the logics into
one place; switch to final variant will happen once we get per-disk
analog of gendisks.  New fields - ->bd_part_sem and ->bd_part_count.

The latter counts the amount of opened partitions.  The former protects
said count _and_ is held while we are rereading partition tables.
Helpers - dev_part_lock()/dev_part_unlock() (currently taking kdev_t; that
will change pretty soon).  No more ->open() and ->release() for partitions,
all that logics went to generic code.  Lock hierachy is currently messy:

  ->bd_sem for partitions -> ->bd_part_sem -> ->bd_sem for entire disks

Ugly, but that'll go away and to get the final variant of locking right
now would take _really_ big patch - with a lot of steps glued together.
The damn thing is large as it is...

for partitioned devices we use ->nr_sect to find the size; blk_size[] is
still used for things like floppy.c, etc.; that will go away later.

There was only one place (do_open()) that needed it - the rest uses
->bd_inode->i_size now.  So blkdev_size_in_bytes() is gone - it's
expanded in its only caller.  Same place (do_open()) finds the partition
offset and stores it in new field ->bd_offset.  As the result, call of
get_gendisk() is gone from the IO path - in blk_partition_remap() we
just add ->bd_offset.

Additionally, we take driver probing (get_blkfops()) outside of ->bd_sem
(again, do_open()) - that will allow to kill ad-hackery in check_partitions()
(opening bdev by hand).

struct gendisk and partition parsers divorced; all these parsers (IBM style,
disklabel, etc.) just fill the structure they get from check_partitions().

Actual setting the things up (filling hd_struct arrays, telling RAID that
we had found partitions worth a look, etc.) is taken into check_partitions()
and done only when we are done with parsing.  Parsers don't know (or care)
what majors/minors they are dealing with; that knowledge also went to
check_partitions().

a bunch of places doing invalidate_device() either didn't need it at all
or actually wanted wipe_partitions().  Switched.

unrelated to the rest, replaces home-grown (racy) semaphores in fs/hfs
with the real thing.

We should _not_ update the current LDT if it's not the current
MM that we are tearing down.

it unconditional for now.

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

This can be overridden by editing the .config file if you really want it.

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

This includes the security_* functions, and the default and capability
modules.

This is needed due to the next header file changes.

This patch changes cont_prepare_write(), in order to support a 4G-1
file for FAT32.

 int cont_prepare_write(struct page *page, unsigned offset,
-		unsigned to, get_block_t *get_block, unsigned long *bytes)
+		unsigned to, get_block_t *get_block, loff_t *bytes)

And it fixes broken adfs/affs/fat/hfs/hpfs/qnx4 by this
cont_prepare_write() change.

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

Been looking at a workload which involves several processes which seek
around and read from a large file.  There are a few problems:
generic_file_lseek is bouncing i_sem around like mad, and readahead is
doing lots of pointless pagecache probing.

This patch addresses readahead.

Presumably the change will be larger on machines which have higher
bandwidth memory than my test box, of which there are many.

This patch teaches readahead to detect the situation where no IO is
actually being performed as a result of its actions.  Now, we don't
want to sacrifice IO efficiency to save a bit of CPU, so the code is
very cautious.  But eventually, after some tens of consecutive
readahead attempts were found to perform no I/O at all, readahead will
turn itself off.

readahead will be turned on again when either generic_file_read() or
filemap_nopage() get a pagecache miss.  The function
handle_ra_thrashing() has been renamed to handle_ra_miss() to reflect
its widened role.

A performance bug in page_cache_readround() was fixed - if
ra->next_size is zero, that function needs to leave it well alone,
because next_size==0 is a magic value meaning that the file has just
been opened and that readahead needs to get aggressive.  This change
makes a `make dep' run at the same speed as in the 2.4 kernel.  It used
to take 4x as long...

`make dep' is an interesting test because it uses mmap to read the files.

The kernel has a number of problems wrt heavy write traffic to multiple
spindles.  What keeps on happening is that all processes which are
responsible for writeback get blocked on one of the queues and all the
others fall idle.

This happens in the balance_dirty_pages() path (balance_dirty() in 2.4)
and in the page reclaim code, when a dirty page is found on the LRU.

The latter is particularly bad because it causes "innocent" processes
to be suspended for long periods due to the activity of heavy writers.

The general idea is: the primary resource for writeback should be the
process which is dirtying memory.  The secondary resource is the
pdflush pool (although this is mainly for providing async writeback in
the presence of light-moderate loads).  Add the final
oh-gee-we-screwed-up resource for writeback is a caller to
shrink_cache().

This patch addresses the balance_dirty_pages() path.  This code was
initially modelled on the 2.4 writeback scheme: throttled processes
writeback all data regardless of its queue.  Instead, the patch changes
it so that the balance_dirty_pages() caller only writes back pages
which are dirty against the queue which that caller just dirtied.

So the effect is a better allocation of writeback resources across the
queues and increased parallelism.

The per-queue writeback is implemented by using
mapping->backing_dev_info as a search key during the walk across the
superblocks and inodes.

The patch also fixes an initialisation problem in
block_dev.c:do_open(): it was setting up the blockdev's
mapping->backing_dev_info too early, before the queue has been
identified.

Generally, this patch doesn't help much, because of the stalls in the
page allocator.  I have a patch which mostly fixes that up, and taken
together the kernel is achieving almost platter speed against six
spindles, but only when the system has a small amount of memory.  More
work is needed there.

A tasty patch from Hugh Dickens.  radix_tree_insert() fails if something
was already present at the target index, so that error can be
propagated back through add_to_page_cache().  Hence
add_to_page_cache_unique() is obsolete.

Hugh's patch removes add_to_page_cache_unique() and cleans up a bunch of
stuff.

Some cleanup from the surprise direct-to-bio for O_DIRECT merge.

- Remove bits and pieces from the kiobuf implementation

- Replace the waitqueue in struct dio with just a task_struct pointer
  and use wake_up_process.  (Ben).

- Only take mmap_sem around the individual calls to get_user_pages().
   (It pins the vmas, yes?)

- Remove some debug code.

- Fix JFS.

Cleanup patch from Martin Bligh: convert some loops which want to be
`for' loops into that, and add some commentary.

generic_writepages() is just a wrapper around mpage_writepages(), so
inline it.

Put the CHECK_EMERGENCY_SYNC back into the kupdate function.  I seem to
keep removing it.

Updated forward-port of Aodrea's O_DIRECT open() checks.  If the user
asked for O_DIRECT and the inode has no mapping or no a_ops then fail
the open up-front.

A patch from Rik which adds some operational statitics to the VM.

In /proc/meminfo:

PageTables:	Amount of memory used for process pagetables
PteChainTot:	Amount of memory allocated for pte_chain objects
PteChainUsed:	Amount of memory currently in use for pte chains.

In /proc/stat:

pageallocs:	Number of pages allocated in the page allocator
pagefrees:	Number of pages returned to the page allocator

		(These can be used to measure the allocation rate)

pageactiv:	Number of pages activated (moved to the active list)
pagedeact:	Number of pages deactivated (moved to the inactive list)
pagefault:	Total pagefaults
majorfault:	Major pagefaults
pagescan:	Number of pages which shrink_cache looked at
pagesteal:	Number of pages which shrink_cache freed
pageoutrun:	Number of calls to try_to_free_pages()
allocstall:	Number of calls to balance_classzone()


Rik will be writing a userspace app which interprets these things.

The /proc/meminfo stats are efficient, but the /proc/stat accumulators
will cause undesirable cacheline bouncing.  We need to break the disk
statistics out of struct kernel_stat and make everything else in there
per-cpu.  If that doesn't happen in time for 2.6 then we disable
KERNEL_STAT_INC().

Patch from David McCracken.  It is an optimisation to the rmap
pte_chains.

In the common case where a page is mapped by only a single pte, we
don't need to allocate a pte_chain structure.  Just make the page's
pte_chain pointer point straight at that pte and flag this with
PG_direct.

Fix to the page reclaim code from Rik.

Anonymous pages which have buffers arise when
truncate_complete_page()'s call to ->releasepage() failed.  Those pages
may still be mapped into process address spaces.

We should not remove them from the LRU, because that makes them
unswappable and they hang around until process exit.

This is the "minimal rmap" patch, writen by Rik, ported to 2.5 by Craig
Kulsea.

Basically,

before: When the page reclaim code decides that is has scanned too many
unreclaimable pages on the LRU it does a scan of process virtual
address spaces for pages to add to swapcache.  ptes pointing at the
page are unmapped as the scan proceeds.  When all ptes referring to a
page have been unmapped and it has been written to swap the page is
reclaimable.

after: When an anonymous page is encountered on the tail of the LRU we
use the rmap to see if it hasn't been referenced lately.  If so then
add it to swapcache.  When the page is again encountered on the LRU, if
it is still unreferenced then try to unmap all ptes which refer to it
in one hit, and if it is clean (ie: on swap) then free it.

The rest of the VM - list management, the classzone concept, etc
remains unchanged.

There are a number of things which the per-page pte chain could be
used for.  Bill Irwin has identified the following.


(1)  page replacement no longer goes around randomly unmapping things

(2)  referenced bits are more accurate because there aren't several ms
        or even seconds between find the multiple pte's mapping a page

(3)  reduces page replacement from O(total virtually mapped) to O(physical)

(4)  enables defragmentation of physical memory

(5)  enables cooperative offlining of memory for friendly guest instance
        behavior in UML and/or LPAR settings

(6)  demonstrable benefit in performance of swapping which is common in
        end-user interactive workstation workloads (I don't like the word
        "desktop"). c.f. Craig Kulesa's post wrt. swapping performance

(7)  evidence from 2.4-based rmap trees indicates approximate parity
        with mainline in kernel compiles with appropriate locking bits

(8)  partitioning of physical memory can reduce the complexity of page
        replacement searches by scanning only the "interesting" zones
        implemented and merged in 2.4-based rmap

(9)  partitioning of physical memory can increase the parallelism of page
        replacement searches by independently processing different zones
        implemented, but not merged in 2.4-based rmap

(10) the reverse mappings may be used for efficiently keeping pte cache
        attributes coherent

(11) they may be used for virtual cache invalidation (with changes)

(12) the reverse mappings enable proper RSS limit enforcement
        implemented and merged in 2.4-based rmap



The code adds a pointer to struct page, consumes additional storage for
the pte chains and adds computational expense to the page reclaim code
(I measured it at 3% additional load during streaming I/O).  The
benefits which we get back for all this are, I must say, theoretical
and unproven.  If it has real advantages (or, indeed, disadvantages)
then why has nobody demonstrated them?



There are a number of things remaining to be done:

1: Demonstrate the above advantages.

2: Make it work with pte-highmem  (Bill Irwin is signed up for this)

3: Don't add pte_chains to non-shared pages optimisation (Dave McCracken's
   patch does this)

4: Move the pte_chains into highmem too (Bill, I guess)

5: per-cpu pte_chain freelists (Rik?)

6: maybe GC the pte_chain backing pages. (Seems unavoidable.  Rik?)

7: multithread the page reclaim code.  (I have patches).

8: clustered add-to-swap.  Not sure if I buy this.  anon pages are
   often well-ordered-by-virtual-address on the LRU, so it "just
   works" for benchmarky loads.  But there may be some other loads...

9: Fix bad IO latency in page reclaim (I have lame patches)

10: Develop tuning tools, use them.

11: The nightly updatedb run is still evicting everything.

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

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

create_serial, get_free_serial and usb_serial_probe all do pretty much
the same thing. I'd like to reorg this into create_serial does all the
alloc and most of the setup, and get_free_serial just fills in the
MAGIC.

There's currently a memory leak: if create_serial is called at probe
time or calc_ports time, and then get_free_serial returns NULL because
the table has no entries left, that usb_serial struct is leaked.

get_free_serial doesn't check properly for free slots. The middle loop
doesn't terminate when the end of the table is reached, although the
assignment loop later does. The effect is that stuff past the end of
the table is allowed to decide if there's free space or not, and
occasionally it'll say "yes" and then the assignment loop will only
allocate slots up to the end of the table, preventing memory
scribbling.

I haven't fixed any of this just yet because I'm not sure what the
intended behaviour is. Should get_free_serial allocate as many slots
as possible, or just be all or nothing? Similarly, I don't see a
problem with calling create_serial early in usb_serial_probe, and
removing the alloc code from get_free_serial; this would fix the leak.

Ah heck, here's a patch. This is what I think things should look like.
get_free_serial is all or none, the leak is fixed and create_serial
does all the allocation.

Name: Designated initializers for drivers/usb
Author: Rusty Russell
Status: Trivial

D: The old form of designated initializers are obsolete: we need to
D: replace them with the ISO C forms before 2.6.  Gcc has always supported
D: both forms anyway.

  Trivia time:

  - C99 conforming initializations by Rusty.

  - ide__sti() -> local_irq_enable() and its friends.