Signed-off-by: NDavid S. Miller <davem@davemloft.net>

The kernel's implementation of notifier chains is unsafe.  There is no
protection against entries being added to or removed from a chain while the
chain is in use.  The issues were discussed in this thread:

    http://marc.theaimsgroup.com/?l=linux-kernel&m=113018709002036&w=2

We noticed that notifier chains in the kernel fall into two basic usage
classes:

	"Blocking" chains are always called from a process context
	and the callout routines are allowed to sleep;

	"Atomic" chains can be called from an atomic context and
	the callout routines are not allowed to sleep.

We decided to codify this distinction and make it part of the API.  Therefore
this set of patches introduces three new, parallel APIs: one for blocking
notifiers, one for atomic notifiers, and one for "raw" notifiers (which is
really just the old API under a new name).  New kinds of data structures are
used for the heads of the chains, and new routines are defined for
registration, unregistration, and calling a chain.  The three APIs are
explained in include/linux/notifier.h and their implementation is in
kernel/sys.c.

With atomic and blocking chains, the implementation guarantees that the chain
links will not be corrupted and that chain callers will not get messed up by
entries being added or removed.  For raw chains the implementation provides no
guarantees at all; users of this API must provide their own protections.  (The
idea was that situations may come up where the assumptions of the atomic and
blocking APIs are not appropriate, so it should be possible for users to
handle these things in their own way.)

There are some limitations, which should not be too hard to live with.  For
atomic/blocking chains, registration and unregistration must always be done in
a process context since the chain is protected by a mutex/rwsem.  Also, a
callout routine for a non-raw chain must not try to register or unregister
entries on its own chain.  (This did happen in a couple of places and the code
had to be changed to avoid it.)

Since atomic chains may be called from within an NMI handler, they cannot use
spinlocks for synchronization.  Instead we use RCU.  The overhead falls almost
entirely in the unregister routine, which is okay since unregistration is much
less frequent that calling a chain.

Here is the list of chains that we adjusted and their classifications.  None
of them use the raw API, so for the moment it is only a placeholder.

  ATOMIC CHAINS
  -------------
arch/i386/kernel/traps.c:		i386die_chain
arch/ia64/kernel/traps.c:		ia64die_chain
arch/powerpc/kernel/traps.c:		powerpc_die_chain
arch/sparc64/kernel/traps.c:		sparc64die_chain
arch/x86_64/kernel/traps.c:		die_chain
drivers/char/ipmi/ipmi_si_intf.c:	xaction_notifier_list
kernel/panic.c:				panic_notifier_list
kernel/profile.c:			task_free_notifier
net/bluetooth/hci_core.c:		hci_notifier
net/ipv4/netfilter/ip_conntrack_core.c:	ip_conntrack_chain
net/ipv4/netfilter/ip_conntrack_core.c:	ip_conntrack_expect_chain
net/ipv6/addrconf.c:			inet6addr_chain
net/netfilter/nf_conntrack_core.c:	nf_conntrack_chain
net/netfilter/nf_conntrack_core.c:	nf_conntrack_expect_chain
net/netlink/af_netlink.c:		netlink_chain

  BLOCKING CHAINS
  ---------------
arch/powerpc/platforms/pseries/reconfig.c:	pSeries_reconfig_chain
arch/s390/kernel/process.c:		idle_chain
arch/x86_64/kernel/process.c		idle_notifier
drivers/base/memory.c:			memory_chain
drivers/cpufreq/cpufreq.c		cpufreq_policy_notifier_list
drivers/cpufreq/cpufreq.c		cpufreq_transition_notifier_list
drivers/macintosh/adb.c:		adb_client_list
drivers/macintosh/via-pmu.c		sleep_notifier_list
drivers/macintosh/via-pmu68k.c		sleep_notifier_list
drivers/macintosh/windfarm_core.c	wf_client_list
drivers/usb/core/notify.c		usb_notifier_list
drivers/video/fbmem.c			fb_notifier_list
kernel/cpu.c				cpu_chain
kernel/module.c				module_notify_list
kernel/profile.c			munmap_notifier
kernel/profile.c			task_exit_notifier
kernel/sys.c				reboot_notifier_list
net/core/dev.c				netdev_chain
net/decnet/dn_dev.c:			dnaddr_chain
net/ipv4/devinet.c:			inetaddr_chain

It's possible that some of these classifications are wrong.  If they are,
please let us know or submit a patch to fix them.  Note that any chain that
gets called very frequently should be atomic, because the rwsem read-locking
used for blocking chains is very likely to incur cache misses on SMP systems.
(However, if the chain's callout routines may sleep then the chain cannot be
atomic.)

The patch set was written by Alan Stern and Chandra Seetharaman, incorporating
material written by Keith Owens and suggestions from Paul McKenney and Andrew
Morton.

[jes@sgi.com: restructure the notifier chain initialization macros]
Signed-off-by: NAlan Stern <stern@rowland.harvard.edu>
Signed-off-by: NChandra Seetharaman <sekharan@us.ibm.com>
Signed-off-by: NJes Sorensen <jes@sgi.com>
Signed-off-by: NAndrew Morton <akpm@osdl.org>
Signed-off-by: NLinus Torvalds <torvalds@osdl.org>

This patchset provides a new (written from scratch) implementation of robust
futexes, called "lightweight robust futexes".  We believe this new
implementation is faster and simpler than the vma-based robust futex solutions
presented before, and we'd like this patchset to be adopted in the upstream
kernel.  This is version 1 of the patchset.

  Background
  ----------

What are robust futexes?  To answer that, we first need to understand what
futexes are: normal futexes are special types of locks that in the
noncontended case can be acquired/released from userspace without having to
enter the kernel.

A futex is in essence a user-space address, e.g.  a 32-bit lock variable
field.  If userspace notices contention (the lock is already owned and someone
else wants to grab it too) then the lock is marked with a value that says
"there's a waiter pending", and the sys_futex(FUTEX_WAIT) syscall is used to
wait for the other guy to release it.  The kernel creates a 'futex queue'
internally, so that it can later on match up the waiter with the waker -
without them having to know about each other.  When the owner thread releases
the futex, it notices (via the variable value) that there were waiter(s)
pending, and does the sys_futex(FUTEX_WAKE) syscall to wake them up.  Once all
waiters have taken and released the lock, the futex is again back to
'uncontended' state, and there's no in-kernel state associated with it.  The
kernel completely forgets that there ever was a futex at that address.  This
method makes futexes very lightweight and scalable.

"Robustness" is about dealing with crashes while holding a lock: if a process
exits prematurely while holding a pthread_mutex_t lock that is also shared
with some other process (e.g.  yum segfaults while holding a pthread_mutex_t,
or yum is kill -9-ed), then waiters for that lock need to be notified that the
last owner of the lock exited in some irregular way.

To solve such types of problems, "robust mutex" userspace APIs were created:
pthread_mutex_lock() returns an error value if the owner exits prematurely -
and the new owner can decide whether the data protected by the lock can be
recovered safely.

There is a big conceptual problem with futex based mutexes though: it is the
kernel that destroys the owner task (e.g.  due to a SEGFAULT), but the kernel
cannot help with the cleanup: if there is no 'futex queue' (and in most cases
there is none, futexes being fast lightweight locks) then the kernel has no
information to clean up after the held lock!  Userspace has no chance to clean
up after the lock either - userspace is the one that crashes, so it has no
opportunity to clean up.  Catch-22.

In practice, when e.g.  yum is kill -9-ed (or segfaults), a system reboot is
needed to release that futex based lock.  This is one of the leading
bugreports against yum.

To solve this problem, 'Robust Futex' patches were created and presented on
lkml: the one written by Todd Kneisel and David Singleton is the most advanced
at the moment.  These patches all tried to extend the futex abstraction by
registering futex-based locks in the kernel - and thus give the kernel a
chance to clean up.

E.g.  in David Singleton's robust-futex-6.patch, there are 3 new syscall
variants to sys_futex(): FUTEX_REGISTER, FUTEX_DEREGISTER and FUTEX_RECOVER.
The kernel attaches such robust futexes to vmas (via
vma->vm_file->f_mapping->robust_head), and at do_exit() time, all vmas are
searched to see whether they have a robust_head set.

Lots of work went into the vma-based robust-futex patch, and recently it has
improved significantly, but unfortunately it still has two fundamental
problems left:

 - they have quite complex locking and race scenarios.  The vma-based
   patches had been pending for years, but they are still not completely
   reliable.

 - they have to scan _every_ vma at sys_exit() time, per thread!

The second disadvantage is a real killer: pthread_exit() takes around 1
microsecond on Linux, but with thousands (or tens of thousands) of vmas every
pthread_exit() takes a millisecond or more, also totally destroying the CPU's
L1 and L2 caches!

This is very much noticeable even for normal process sys_exit_group() calls:
the kernel has to do the vma scanning unconditionally!  (this is because the
kernel has no knowledge about how many robust futexes there are to be cleaned
up, because a robust futex might have been registered in another task, and the
futex variable might have been simply mmap()-ed into this process's address
space).

This huge overhead forced the creation of CONFIG_FUTEX_ROBUST, but worse than
that: the overhead makes robust futexes impractical for any type of generic
Linux distribution.

So it became clear to us, something had to be done.  Last week, when Thomas
Gleixner tried to fix up the vma-based robust futex patch in the -rt tree, he
found a handful of new races and we were talking about it and were analyzing
the situation.  At that point a fundamentally different solution occured to
me.  This patchset (written in the past couple of days) implements that new
solution.  Be warned though - the patchset does things we normally dont do in
Linux, so some might find the approach disturbing.  Parental advice
recommended ;-)

  New approach to robust futexes
  ------------------------------

At the heart of this new approach there is a per-thread private list of robust
locks that userspace is holding (maintained by glibc) - which userspace list
is registered with the kernel via a new syscall [this registration happens at
most once per thread lifetime].  At do_exit() time, the kernel checks this
user-space list: are there any robust futex locks to be cleaned up?

In the common case, at do_exit() time, there is no list registered, so the
cost of robust futexes is just a simple current->robust_list != NULL
comparison.  If the thread has registered a list, then normally the list is
empty.  If the thread/process crashed or terminated in some incorrect way then
the list might be non-empty: in this case the kernel carefully walks the list
[not trusting it], and marks all locks that are owned by this thread with the
FUTEX_OWNER_DEAD bit, and wakes up one waiter (if any).

The list is guaranteed to be private and per-thread, so it's lockless.  There
is one race possible though: since adding to and removing from the list is
done after the futex is acquired by glibc, there is a few instructions window
for the thread (or process) to die there, leaving the futex hung.  To protect
against this possibility, userspace (glibc) also maintains a simple per-thread
'list_op_pending' field, to allow the kernel to clean up if the thread dies
after acquiring the lock, but just before it could have added itself to the
list.  Glibc sets this list_op_pending field before it tries to acquire the
futex, and clears it after the list-add (or list-remove) has finished.

That's all that is needed - all the rest of robust-futex cleanup is done in
userspace [just like with the previous patches].

Ulrich Drepper has implemented the necessary glibc support for this new
mechanism, which fully enables robust mutexes.  (Ulrich plans to commit these
changes to glibc-HEAD later today.)

Key differences of this userspace-list based approach, compared to the vma
based method:

 - it's much, much faster: at thread exit time, there's no need to loop
   over every vma (!), which the VM-based method has to do.  Only a very
   simple 'is the list empty' op is done.

 - no VM changes are needed - 'struct address_space' is left alone.

 - no registration of individual locks is needed: robust mutexes dont need
   any extra per-lock syscalls.  Robust mutexes thus become a very lightweight
   primitive - so they dont force the application designer to do a hard choice
   between performance and robustness - robust mutexes are just as fast.

 - no per-lock kernel allocation happens.

 - no resource limits are needed.

 - no kernel-space recovery call (FUTEX_RECOVER) is needed.

 - the implementation and the locking is "obvious", and there are no
   interactions with the VM.

  Performance
  -----------

I have benchmarked the time needed for the kernel to process a list of 1
million (!) held locks, using the new method [on a 2GHz CPU]:

 - with FUTEX_WAIT set [contended mutex]: 130 msecs
 - without FUTEX_WAIT set [uncontended mutex]: 30 msecs

I have also measured an approach where glibc does the lock notification [which
it currently does for !pshared robust mutexes], and that took 256 msecs -
clearly slower, due to the 1 million FUTEX_WAKE syscalls userspace had to do.

(1 million held locks are unheard of - we expect at most a handful of locks to
be held at a time.  Nevertheless it's nice to know that this approach scales
nicely.)

  Implementation details
  ----------------------

The patch adds two new syscalls: one to register the userspace list, and one
to query the registered list pointer:

 asmlinkage long
 sys_set_robust_list(struct robust_list_head __user *head,
                     size_t len);

 asmlinkage long
 sys_get_robust_list(int pid, struct robust_list_head __user **head_ptr,
                     size_t __user *len_ptr);

List registration is very fast: the pointer is simply stored in
current->robust_list.  [Note that in the future, if robust futexes become
widespread, we could extend sys_clone() to register a robust-list head for new
threads, without the need of another syscall.]

So there is virtually zero overhead for tasks not using robust futexes, and
even for robust futex users, there is only one extra syscall per thread
lifetime, and the cleanup operation, if it happens, is fast and
straightforward.  The kernel doesnt have any internal distinction between
robust and normal futexes.

If a futex is found to be held at exit time, the kernel sets the highest bit
of the futex word:

	#define FUTEX_OWNER_DIED        0x40000000

and wakes up the next futex waiter (if any). User-space does the rest of
the cleanup.

Otherwise, robust futexes are acquired by glibc by putting the TID into the
futex field atomically.  Waiters set the FUTEX_WAITERS bit:

	#define FUTEX_WAITERS           0x80000000

and the remaining bits are for the TID.

  Testing, architecture support
  -----------------------------

I've tested the new syscalls on x86 and x86_64, and have made sure the parsing
of the userspace list is robust [ ;-) ] even if the list is deliberately
corrupted.

i386 and x86_64 syscalls are wired up at the moment, and Ulrich has tested the
new glibc code (on x86_64 and i386), and it works for his robust-mutex
testcases.

All other architectures should build just fine too - but they wont have the
new syscalls yet.

Architectures need to implement the new futex_atomic_cmpxchg_inuser() inline
function before writing up the syscalls (that function returns -ENOSYS right
now).

This patch:

Add placeholder futex_atomic_cmpxchg_inuser() implementations to every
architecture that supports futexes.  It returns -ENOSYS.
Signed-off-by: NIngo Molnar <mingo@elte.hu>
Signed-off-by: NThomas Gleixner <tglx@linutronix.de>
Signed-off-by: NArjan van de Ven <arjan@infradead.org>
Acked-by: NUlrich Drepper <drepper@redhat.com>
Signed-off-by: NAndrew Morton <akpm@osdl.org>
Signed-off-by: NLinus Torvalds <torvalds@osdl.org>

There are 3 memory models, FLATMEM, DISCONTIGMEM, SPARSEMEM.
Each arch has its own page_to_pfn(), pfn_to_page() for each models.
But most of them can use the same arithmetic.

This patch adds asm-generic/memory_model.h, which includes generic
page_to_pfn(), pfn_to_page() definitions for each memory model.

When CONFIG_OUT_OF_LINE_PFN_TO_PAGE=y, out-of-line functions are
used instead of macro. This is enabled by some archs and  reduces
text size.
Signed-off-by: NKAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Hugh Dickins <hugh@veritas.com>
Cc: Andi Kleen <ak@muc.de>
Cc: Paul Mackerras <paulus@samba.org>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: Richard Henderson <rth@twiddle.net>
Cc: Ivan Kokshaysky <ink@jurassic.park.msu.ru>
Cc: Russell King <rmk@arm.linux.org.uk>
Cc: Ian Molton <spyro@f2s.com>
Cc: Mikael Starvik <starvik@axis.com>
Cc: David Howells <dhowells@redhat.com>
Cc: Yoshinori Sato <ysato@users.sourceforge.jp>
Cc: Hirokazu Takata <takata.hirokazu@renesas.com>
Cc: Ralf Baechle <ralf@linux-mips.org>
Cc: Kyle McMartin <kyle@mcmartin.ca>
Cc: Heiko Carstens <heiko.carstens@de.ibm.com>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Cc: Paul Mundt <lethal@linux-sh.org>
Cc: Kazumoto Kojima <kkojima@rr.iij4u.or.jp>
Cc: Richard Curnow <rc@rc0.org.uk>
Cc: William Lee Irwin III <wli@holomorphy.com>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Jeff Dike <jdike@addtoit.com>
Cc: Paolo 'Blaisorblade' Giarrusso <blaisorblade@yahoo.it>
Cc: Miles Bader <uclinux-v850@lsi.nec.co.jp>
Cc: Chris Zankel <chris@zankel.net>
Cc: "Luck, Tony" <tony.luck@intel.com>
Signed-off-by: NAndrew Morton <akpm@osdl.org>
Signed-off-by: NLinus Torvalds <torvalds@osdl.org>

- remove __{,test_and_}{set,clear,change}_bit() and test_bit()
- remove ffz()
- remove __ffs()
- remove generic_fls()
- remove generic_fls64()
- remove sched_find_first_bit()
- remove ffs()

- unless defined(ULTRA_HAS_POPULATION_COUNT)

  - remove generic_hweight{64,32,16,8}()

- remove find_{next,first}{,_zero}_bit()
- remove ext2_{set,clear,test,find_first_zero,find_next_zero}_bit()
- remove minix_{test,set,test_and_clear,test,find_first_zero}_bit()
Signed-off-by: NAkinobu Mita <mita@miraclelinux.com>
Cc: "David S. Miller" <davem@davemloft.net>
Signed-off-by: NAndrew Morton <akpm@osdl.org>
Signed-off-by: NLinus Torvalds <torvalds@osdl.org>

Bitmap functions for the minix filesystem and the ext2 filesystem except
ext2_set_bit_atomic() and ext2_clear_bit_atomic() do not require the atomic
guarantees.

But these are defined by using atomic bit operations on several architectures.
 (cris, frv, h8300, ia64, m32r, m68k, m68knommu, mips, s390, sh, sh64, sparc,
sparc64, v850, and xtensa)

This patch switches to non atomic bit operation.
Signed-off-by: NAkinobu Mita <mita@miraclelinux.com>
Signed-off-by: NAndrew Morton <akpm@osdl.org>
Signed-off-by: NLinus Torvalds <torvalds@osdl.org>

Implement the half-closed devices notifiation, by adding a new POLLRDHUP
(and its alias EPOLLRDHUP) bit to the existing poll/select sets.  Since the
existing POLLHUP handling, that does not report correctly half-closed
devices, was feared to be changed, this implementation leaves the current
POLLHUP reporting unchanged and simply add a new bit that is set in the few
places where it makes sense.  The same thing was discussed and conceptually
agreed quite some time ago:

http://lkml.org/lkml/2003/7/12/116

Since this new event bit is added to the existing Linux poll infrastruture,
even the existing poll/select system calls will be able to use it.  As far
as the existing POLLHUP handling, the patch leaves it as is.  The
pollrdhup-2.6.16.rc5-0.10.diff defines the POLLRDHUP for all the existing
archs and sets the bit in the six relevant files.  The other attached diff
is the simple change required to sys/epoll.h to add the EPOLLRDHUP
definition.

There is "a stupid program" to test POLLRDHUP delivery here:

 http://www.xmailserver.org/pollrdhup-test.c

It tests poll(2), but since the delivery is same epoll(2) will work equally.
Signed-off-by: NDavide Libenzi <davidel@xmailserver.org>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Michael Kerrisk <mtk-manpages@gmx.net>
Signed-off-by: NAndrew Morton <akpm@osdl.org>
Signed-off-by: NLinus Torvalds <torvalds@osdl.org>

When we stop allocating percpu memory for not-possible CPUs we must not touch
the percpu data for not-possible CPUs at all.  The correct way of doing this
is to test cpu_possible() or to use for_each_cpu().

This patch is a kernel-wide sweep of all instances of NR_CPUS.  I found very
few instances of this bug, if any.  But the patch converts lots of open-coded
test to use the preferred helper macros.

Cc: Mikael Starvik <starvik@axis.com>
Cc: David Howells <dhowells@redhat.com>
Acked-by: NKyle McMartin <kyle@parisc-linux.org>
Cc: Anton Blanchard <anton@samba.org>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: Paul Mackerras <paulus@samba.org>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Cc: Heiko Carstens <heiko.carstens@de.ibm.com>
Cc: Paul Mundt <lethal@linux-sh.org>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: William Lee Irwin III <wli@holomorphy.com>
Cc: Andi Kleen <ak@muc.de>
Cc: Christian Zankel <chris@zankel.net>
Cc: Philippe Elie <phil.el@wanadoo.fr>
Cc: Nathan Scott <nathans@sgi.com>
Cc: Jens Axboe <axboe@suse.de>
Cc: Eric Dumazet <dada1@cosmosbay.com>
Signed-off-by: NAndrew Morton <akpm@osdl.org>
Signed-off-by: NLinus Torvalds <torvalds@osdl.org>

Without branch hints, the very unlikely chance of the loop repeating due to
cmpxchg failure is unrolled with gcc-4 that I have tested.

Improve this for architectures with a native cas/cmpxchg.  llsc archs
should try to implement this natively.
Signed-off-by: NNick Piggin <npiggin@suse.de>
Cc: Andi Kleen <ak@muc.de>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Cc: Heiko Carstens <heiko.carstens@de.ibm.com>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Roman Zippel <zippel@linux-m68k.org>
Signed-off-by: NAndrew Morton <akpm@osdl.org>
Signed-off-by: NLinus Torvalds <torvalds@osdl.org>

Seems like needless clutter having a bunch of #if defined(CONFIG_$ARCH) in
include/linux/cache.h.  Move the per architecture section definition to
asm/cache.h, and keep the if-not-defined dummy case in linux/cache.h to
catch architectures which don't implement the section.

Verified that symbols still go in .data.read_mostly on parisc,
and the compile doesn't break.
Signed-off-by: NKyle McMartin <kyle@parisc-linux.org>
Signed-off-by: NAndrew Morton <akpm@osdl.org>
Signed-off-by: NLinus Torvalds <torvalds@osdl.org>

Signed-off-by: NDavid S. Miller <davem@davemloft.net>

1) huge_pte_offset() did not check the page table hierarchy
   elements as being empty correctly, resulting in an OOPS

2) Need platform specific hugetlb_get_unmapped_area() to handle
   the top-down vs. bottom-up address space allocation strategies.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

We only need to write an invalid tag every 16 bytes,
so taking advantage of this can save many instructions
compared to the simple memset() call we make now.

A prefetching implementation is implemented for sun4u
and a block-init store version if implemented for Niagara.

The next trick is to be able to perform an init and
a copy_tsb() in parallel when growing a TSB table.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Put it one page below the top of the 32-bit address space.
This gives us ~16MB more address space to work with.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Currently allocations are very constrained for 32-bit processes.
It grows down-up from 0x70000000 to 0xf0000000 which gives about
2GB of stack + dynamic mmap() space.

So support the top-down method, and we need to override the
generic helper function in order to deal with D-cache coloring.

With these changes I was able to squeeze out a mmap() just over
3.6GB in size in a 32-bit process.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

This is good for up to %50 performance improvement of some test cases.
The problem has been the race conditions, and hopefully I've plugged
them all up here.

1) There was a serious race in switch_mm() wrt. lazy TLB
   switching to and from kernel threads.

   We could erroneously skip a tsb_context_switch() and thus
   use a stale TSB across a TSB grow event.

   There is a big comment now in that function describing
   exactly how it can happen.

2) All code paths that do something with the TSB need to be
   guarded with the mm->context.lock spinlock.  This makes
   page table flushing paths properly synchronize with both
   TSB growing and TLB context changes.

3) TSB growing events are moved to the end of successful fault
   processing.  Previously it was in update_mmu_cache() but
   that is deadlock prone.  At the end of do_sparc64_fault()
   we hold no spinlocks that could deadlock the TSB grow
   sequence.  We also have dropped the address space semaphore.

While we're here, add prefetching to the copy_tsb() routine
and put it in assembler into the tsb.S file.  This piece of
code is quite time critical.

There are some small negative side effects to this code which
can be improved upon.  In particular we grab the mm->context.lock
even for the tsb insert done by update_mmu_cache() now and that's
a bit excessive.  We can get rid of that locking, and the same
lock taking in flush_tsb_user(), by disabling PSTATE_IE around
the whole operation including the capturing of the tsb pointer
and tsb_nentries value.  That would work because anyone growing
the TSB won't free up the old TSB until all cpus respond to the
TSB change cross call.

I'm not quite so confident in that optimization to put it in
right now, but eventually we might be able to and the description
is here for reference.

This code seems very solid now.  It passes several parallel GCC
bootstrap builds, and our favorite "nut cruncher" stress test which is
a full "make -j8192" build of a "make allmodconfig" kernel.  That puts
about 256 processes on each cpu's run queue, makes lots of process cpu
migrations occur, causes lots of page table and TLB flushing activity,
incurs many context version number changes, and it swaps the machine
real far out to disk even though there is 16GB of ram on this test
system. :-)
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Report 'sun4v' when appropriate in /proc/cpuinfo

Remove all the verifications of the OBP version string.  Just
make sure it's there, and report it raw in the bootup logs and
via /proc/cpuinfo.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

The mapping is a simple "(cpuid >> 2) == core" for now.
Later we'll add more sophisticated code that will walk
the sun4v machine description and figure this out from
there.

We should also add core mappings for jaguar and panther
processors.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

This has been pending for a long time, and the fact
that we waste a ton of ram on some configurations
kind of pushed things over the edge.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Don't piggy back the SMP receive signal code to do the
context version change handling.

Instead allocate another fixed PIL number for this
asynchronous cross-call.  We can't use smp_call_function()
because this thing is invoked with interrupts disabled
and a few spinlocks held.

Also, fix smp_call_function_mask() to count "cpus" correctly.
There is no guarentee that the local cpu is in the mask
yet that is exactly what this code was assuming.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

1) Always spin_lock_init() in init_context().  The caller essentially
   clears it out, or copies the mm info from the parent.  In both
   cases we need to explicitly initialize the spinlock.

2) Always do explicit IRQ disabling while taking mm->context.lock
   and ctx_alloc_lock.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

The UltraSPARC T1 manual recommends this because the chip
could instruction prefetch into the VA hole, and this would
also make decoding  certain kinds of memory access traps
more difficult (because the chip sign extends certain pieces
of trap state).
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

These are all implemented inline earlier in the file.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

There were several bugs in the SUN4V cpu mondo dispatch code.

In fact, if we ever got a EWOULDBLOCK or other error from
the hypervisor call, we'd potentially send a cpu mondo multiple
times to the same cpu and even worse we could loop until the
timeout resending the same mondo over and over to such cpus.

So let's bulletproof this thing as follows:

1) Implement cpu_mondo_send() and cpu_state() hypervisor calls
   in arch/sparc64/kernel/entry.S, add prototypes to asm/hypervisor.h

2) Don't build and update the cpulist using inline functions, this
   was causing the cpu mask to not get updated in the caller.

3) Disable interrupts during the entire mondo send, otherwise our
   cpu list and/or mondo block could get overwritten if we take
   an interrupt and do a cpu mondo send on the current cpu.

4) Check for all possible error return types from the cpu_mondo_send()
   hypervisor call.  In particular:

   HV_EOK) Our work is done, all cpus have received the mondo.
   HV_CPUERROR) One or more of the cpus in the cpu list we passed
                to the hypervisor are in error state.  Use cpu_state()
                calls over the entries in the cpu list to see which
		ones.  Record them in "error_mask" and report this
		after we are done sending the mondo to cpus which are
		not in error state.
   HV_EWOULDBLOCK) We need to keep trying.

   Any other error we consider fatal, we report the event and exit
   immediately.

5) We only timeout if forward progress is not made.  Forward progress
   is defined as having at least one cpu get the mondo successfully
   in a given cpu_mondo_send() call.  Otherwise we bump a counter
   and delay a little.  If the counter hits a limit, we signal an
   error and report the event.

Also, smp_call_function_mask() error handling reports the number
of cpus incorrectly.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

We no longer have the problems that require using the smaller
sizes.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Signed-off-by: NDavid S. Miller <davem@davemloft.net>

The context allocation scheme we use depends upon there being a 1<-->1
mapping from cpu to physical TLB for correctness.  Chips like Niagara
break this assumption.

So what we do is notify all cpus with a cross call when the context
version number changes, and if necessary this makes them allocate
a valid context for the address space they are running at the time.

Stress tested with make -j1024, make -j2048, and make -j4096 kernel
builds on a 32-strand, 8 core, T2000 with 16GB of ram.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Niagara helps us find a ancient bug in the sparc64 port :-)

The ASI_* values are plain constant defines, thus signed 32-bit
on sparc64.  To put shift this into the regs->tstate value we were
doing or'ing "(ASI_PNF << 24)" into there.

ASI_PNF is 0x82 and shifted left by 24 makes that topmost bit the
sign bit in a 32-bit value.  This would get sign extended to 64-bits
and thus corrupt the top-half of the reg->tstate value.

This never caused problems in pre-Niagara cpus because the only thing
up there were the condition code values.  But Niagara has the global
register level field, and this all 1's value is illegal there so
Niagara gives an illegal instruction trap due to this bug.

I'm pretty sure this bug is about as old as the sparc64 port itself.

This also points out that we weren't setting ASI_PNF for 32-bit tasks.
We should, so fix that while we're here.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

It can map all of the linear kernel mappings with zero TSB hash
conflicts for systems with 16GB or less ram.  In such cases, on
SUN4V, once we load up this TSB the first time with all the
mappings, we never take a linear kernel mapping TLB miss ever
again, the hypervisor handles them all.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Set, but never used.

We used to use this for dynamic IRQ retargetting, but that
code died a long time ago.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

For drivers/media/*, noticed by Fabbione.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Signed-off-by: NFabio M. Di Nitto <fabbione@ubuntu.com>
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

The SUN4V convention with non-shared TSBs is that the context
bit of the TAG is clear.  So we have to choose an "invalid"
bit and initialize new TSBs appropriately.  Otherwise a zero
TAG looks "valid".

Make sure, for the window fixup cases, that we use the right
global registers and that we don't potentially trample on
the live global registers in etrap/rtrap handling (%g2 and
%g6) and that we put the missing virtual address properly
in %g5.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

This gives more consistent bogomips and delay() semantics,
especially on sun4v.  It gives weird looking values though...
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Signed-off-by: NDavid S. Miller <davem@davemloft.net>

Signed-off-by: NDavid S. Miller <davem@davemloft.net>

We need to use the real hardware processor ID when
targetting interrupts, not the "define to 0" thing
the uniprocessor build gives us.

Also, fill in the Node-ID and Agent-ID fields properly
on sun4u/Safari.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

The sibling cpu bringup is extremely fragile.  We can only
perform the most basic calls until we take over the trap
table from the firmware/hypervisor on the new cpu.

This means no accesses to %g4, %g5, %g6 since those can't be
TLB translated without our trap handlers.

In order to achieve this:

1) Change sun4v_init_mondo_queues() so that it can operate in
   several modes.

   It can allocate the queues, or install them in the current
   processor, or both.

   The boot cpu does both in it's call early on.

   Later, the boot cpu allocates the sibling cpu queue, starts
   the sibling cpu, then the sibling cpu loads them in.

2) init_cur_cpu_trap() is changed to take the current_thread_info()
   as an argument instead of reading %g6 directly on the current
   cpu.

3) Create a trampoline stack for the sibling cpus.  We do our basic
   kernel calls using this stack, which is locked into the kernel
   image, then go to our proper thread stack after taking over the
   trap table.

4) While we are in this delicate startup state, we put 0xdeadbeef
   into %g4/%g5/%g6 in order to catch accidental accesses.

5) On the final prom_set_trap_table*() call, we put &init_thread_union
   into %g6.  This is a hack to make prom_world(0) work.  All that
   wants to do is restore the %asi register using
   get_thread_current_ds().

Longer term we should just do the OBP calls to set the trap table by
hand just like we do for everything else.  This would avoid that silly
prom_world(0) issue, then we can remove the init_thread_union hack.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>

So we can read the %gl register for debugging.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>