Allows reading of SK_MEMINFO_VARS via socket option. This way an
application can get all meminfo related information in single socket
option call instead of multiple calls.

Adds helper function, sk_get_meminfo(), and uses that for both
getsockopt and sock_diag_put_meminfo().

Suggested by Eric Dumazet.

Signed-off-by: Josh Hunt <johunt@akamai.com>
Reviewed-by: Jason Baron <jbaron@akamai.com>
Acked-by: Eric Dumazet <edumazet@google.com>
Signed-off-by: David S. Miller <davem@davemloft.net>

Patch series "userfaultfd tmpfs/hugetlbfs/non-cooperative", v2

These userfaultfd features are finished and are ready for larger
exposure in -mm and upstream merging.

1) tmpfs non present userfault
2) hugetlbfs non present userfault
3) non cooperative userfault for fork/madvise/mremap

qemu development code is already exercising 2) and container postcopy
live migration needs 3).

1) is not currently used but there's a self test and we know some qemu
user for various reasons uses tmpfs as backing for KVM so it'll need it
too to use postcopy live migration with tmpfs memory.

All review feedback from the previous submit has been handled and the
fixes are included.  There's no outstanding issue AFIK.

Upstream code just did a s/fe/vmf/ conversion in the page faults and
this has been converted as well incrementally.

In addition to the previous submits, this also wakes up stuck userfaults
during UFFDIO_UNREGISTER.  The non cooperative testcase actually
reproduced this problem by getting stuck instead of quitting clean in
some rare case as it could call UFFDIO_UNREGISTER while some userfault
could be still in flight.  The other option would have been to keep
leaving it up to userland to serialize itself and to patch the testcase
instead but the wakeup during unregister I think is preferable.

David also asked the UFFD_FEATURE_MISSING_HUGETLBFS and
UFFD_FEATURE_MISSING_SHMEM feature flags to be added so QEMU can avoid
to probe if the hugetlbfs/shmem missing support is available by calling
UFFDIO_REGISTER.  QEMU already checks HUGETLBFS_MAGIC with fstatfs so if
UFFD_FEATURE_MISSING_HUGETLBFS is also set, it knows UFFDIO_REGISTER
will succeed (or if it fails, it's for some other more concerning
reason).  There's no reason to worry about adding too many feature
flags.  There are 64 available and worst case we've to bump the API if
someday we're really going to run out of them.

The round-trip network latency of hugetlbfs userfaults during postcopy
live migration is still of the order of dozen milliseconds on 10GBit if
at 2MB hugepage granularity so it's working perfectly and it should
provide for higher bandwidth or lower CPU usage (which makes it
interesting to add an option in the future to support THP granularity
too for anonymous memory, UFFDIO_COPY would then have to create THP if
alignment/len allows for it).  1GB hugetlbfs granularity will require
big changes in hugetlbfs to work so it's deferred for later.

This patch (of 42):

This adds proper documentation (inline) to avoid the risk of further
misunderstandings about the semantics of _IOW/_IOR and it also reminds
whoever will bump the UFFDIO_API in the future, to change the two ioctl
to _IOW.

This was found while implementing strace support for those ioctl,
otherwise we could have never found it by just reviewing kernel code and
testing it.

_IOC_READ or _IOC_WRITE alters nothing but the ioctl number itself, so
it's only worth fixing if the UFFDIO_API is bumped someday.

Link: http://lkml.kernel.org/r/20161216144821.5183-2-aarcange@redhat.com


Signed-off-by: Andrea Arcangeli <aarcange@redhat.com>
Reported-by: "Dmitry V. Levin" <ldv@altlinux.org>
Cc: Michael Rapoport <RAPOPORT@il.ibm.com>
Cc: "Dr. David Alan Gilbert" <dgilbert@redhat.com>
Cc: Mike Kravetz <mike.kravetz@oracle.com>
Cc: Pavel Emelyanov <xemul@parallels.com>
Cc: Hillf Danton <hillf.zj@alibaba-inc.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

This patch exports the sender chronograph stats via the socket
SO_TIMESTAMPING channel. Currently we can instrument how long a
particular application unit of data was queued in TCP by tracking
SOF_TIMESTAMPING_TX_SOFTWARE and SOF_TIMESTAMPING_TX_SCHED. Having
these sender chronograph stats exported simultaneously along with
these timestamps allow further breaking down the various sender
limitation.  For example, a video server can tell if a particular
chunk of video on a connection takes a long time to deliver because
TCP was experiencing small receive window. It is not possible to
tell before this patch without packet traces.

To prepare these stats, the user needs to set
SOF_TIMESTAMPING_OPT_STATS and SOF_TIMESTAMPING_OPT_TSONLY flags
while requesting other SOF_TIMESTAMPING TX timestamps. When the
timestamps are available in the error queue, the stats are returned
in a separate control message of type SCM_TIMESTAMPING_OPT_STATS,
in a list of TLVs (struct nlattr) of types: TCP_NLA_BUSY_TIME,
TCP_NLA_RWND_LIMITED, TCP_NLA_SNDBUF_LIMITED. Unit is microsecond.

Signed-off-by: Francis Yan <francisyyan@gmail.com>
Signed-off-by: Yuchung Cheng <ycheng@google.com>
Signed-off-by: Soheil Hassas Yeganeh <soheil@google.com>
Acked-by: Neal Cardwell <ncardwell@google.com>
Signed-off-by: David S. Miller <davem@davemloft.net>

pkey_set() and pkey_get() were syscalls present in older versions
of the protection keys patches.  They were fully excised from the
x86 code, but some cruft was left in the generic syscall code.  The
C++ comments were intended to help to make it more glaring to me to
fix them before actually submitting them.  That technique worked,
but later than I would have liked.

I test-compiled this for arm64.

Fixes: a60f7b69 ("generic syscalls: Wire up memory protection keys syscalls")
Signed-off-by: Dave Hansen <dave.hansen@linux.intel.com>
Acked-by: Arnd Bergmann <arnd@arndb.de>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: x86@kernel.org
Cc: linux-arch@vger.kernel.org
Cc: mgorman@techsingularity.net
Cc: linux-api@vger.kernel.org
Cc: linux-mm@kvack.org
Cc: luto@kernel.org
Cc: akpm@linux-foundation.org
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

These new syscalls are implemented as generic code, so enable them for
architectures like arm64 which use the generic syscall table.

According to Arnd:

  Even if the support is x86 specific for the forseeable future, it may be
  good to reserve the number just in case.  The other architecture specific
  syscall lists are usually left to the individual arch maintainers, most a
  lot of the newer architectures share this table.

Signed-off-by: Dave Hansen <dave.hansen@linux.intel.com>
Acked-by: Arnd Bergmann <arnd@arndb.de>
Cc: linux-arch@vger.kernel.org
Cc: Dave Hansen <dave@sr71.net>
Cc: mgorman@techsingularity.net
Cc: linux-api@vger.kernel.org
Cc: linux-mm@kvack.org
Cc: luto@kernel.org
Cc: akpm@linux-foundation.org
Cc: torvalds@linux-foundation.org
Link: http://lkml.kernel.org/r/20160729163018.505A6875@viggo.jf.intel.com


Signed-off-by: Thomas Gleixner <tglx@linutronix.de>

This patch adds two new system calls:

	int pkey_alloc(unsigned long flags, unsigned long init_access_rights)
	int pkey_free(int pkey);

These implement an "allocator" for the protection keys
themselves, which can be thought of as analogous to the allocator
that the kernel has for file descriptors.  The kernel tracks
which numbers are in use, and only allows operations on keys that
are valid.  A key which was not obtained by pkey_alloc() may not,
for instance, be passed to pkey_mprotect().

These system calls are also very important given the kernel's use
of pkeys to implement execute-only support.  These help ensure
that userspace can never assume that it has control of a key
unless it first asks the kernel.  The kernel does not promise to
preserve PKRU (right register) contents except for allocated
pkeys.

The 'init_access_rights' argument to pkey_alloc() specifies the
rights that will be established for the returned pkey.  For
instance:

	pkey = pkey_alloc(flags, PKEY_DENY_WRITE);

will allocate 'pkey', but also sets the bits in PKRU[1] such that
writing to 'pkey' is already denied.

The kernel does not prevent pkey_free() from successfully freeing
in-use pkeys (those still assigned to a memory range by
pkey_mprotect()).  It would be expensive to implement the checks
for this, so we instead say, "Just don't do it" since sane
software will never do it anyway.

Any piece of userspace calling pkey_alloc() needs to be prepared
for it to fail.  Why?  pkey_alloc() returns the same error code
(ENOSPC) when there are no pkeys and when pkeys are unsupported.
They can be unsupported for a whole host of reasons, so apps must
be prepared for this.  Also, libraries or LD_PRELOADs might steal
keys before an application gets access to them.

This allocation mechanism could be implemented in userspace.
Even if we did it in userspace, we would still need additional
user/kernel interfaces to tell userspace which keys are being
used by the kernel internally (such as for execute-only
mappings).  Having the kernel provide this facility completely
removes the need for these additional interfaces, or having an
implementation of this in userspace at all.

Note that we have to make changes to all of the architectures
that do not use mman-common.h because we use the new
PKEY_DENY_ACCESS/WRITE macros in arch-independent code.

1. PKRU is the Protection Key Rights User register.  It is a
   usermode-accessible register that controls whether writes
   and/or access to each individual pkey is allowed or denied.

Signed-off-by: Dave Hansen <dave.hansen@linux.intel.com>
Acked-by: Mel Gorman <mgorman@techsingularity.net>
Cc: linux-arch@vger.kernel.org
Cc: Dave Hansen <dave@sr71.net>
Cc: arnd@arndb.de
Cc: linux-api@vger.kernel.org
Cc: linux-mm@kvack.org
Cc: luto@kernel.org
Cc: akpm@linux-foundation.org
Cc: torvalds@linux-foundation.org
Link: http://lkml.kernel.org/r/20160729163015.444FE75F@viggo.jf.intel.com


Signed-off-by: Thomas Gleixner <tglx@linutronix.de>

The newer renameat2 syscall provides all the functionality provided by
the renameat syscall and adds flags, so future architectures won't need
to include renameat.

Therefore drop the renameat syscall from the generic syscall list unless
__ARCH_WANT_RENAMEAT is defined by the architecture's unistd.h prior to
including asm-generic/unistd.h, and adjust all architectures using the
generic syscall list to define it so that no in-tree architectures are
affected.

Signed-off-by: James Hogan <james.hogan@imgtec.com>
Acked-by: Vineet Gupta <vgupta@synopsys.com>
Cc: linux-arch@vger.kernel.org
Cc: linux-snps-arc@lists.infradead.org
Cc: Catalin Marinas <catalin.marinas@arm.com>
Cc: Will Deacon <will.deacon@arm.com>
Cc: linux-arm-kernel@lists.infradead.org
Cc: Mark Salter <msalter@redhat.com>
Cc: Aurelien Jacquiot <a-jacquiot@ti.com>
Cc: linux-c6x-dev@linux-c6x.org
Cc: Richard Kuo <rkuo@codeaurora.org>
Cc: linux-hexagon@vger.kernel.org
Cc: linux-metag@vger.kernel.org
Cc: Jonas Bonn <jonas@southpole.se>
Cc: linux@lists.openrisc.net
Cc: Chen Liqin <liqin.linux@gmail.com>
Cc: Lennox Wu <lennox.wu@gmail.com>
Cc: Chris Metcalf <cmetcalf@mellanox.com>
Cc: Guan Xuetao <gxt@mprc.pku.edu.cn>
Cc: Ley Foon Tan <lftan@altera.com>
Cc: nios2-dev@lists.rocketboards.org
Cc: Yoshinori Sato <ysato@users.sourceforge.jp>
Cc: uclinux-h8-devel@lists.sourceforge.jp
Signed-off-by: Arnd Bergmann <arnd@arndb.de>

Compat architectures that does not use generic unistd (mips, s390),
declare compat version in their syscall tables for preadv2 and
pwritev2. Generic unistd syscall table should do it as well.

[arnd: this initially slipped through the review and an
 incorrect patch got merged. arch/tile/ is the only architecture
 that could be affected for their 32-bit compat mode, every
 other architecture we support today is fine.]

Signed-off-by: Yury Norov <ynorov@caviumnetworks.com>
Signed-off-by: Arnd Bergmann <arnd@arndb.de>

These new syscalls are implemented as generic code, so enable them for
architectures like arm64 which use the generic syscall table.

Signed-off-by: Andre Przywara <andre.przywara@arm.com>
Signed-off-by: Arnd Bergmann <arnd@arndb.de>

Stephen Rothwell reported this linux-next build failure:

	http://lkml.kernel.org/r/20160226164406.065a1ffc@canb.auug.org.au



... caused by the Memory Protection Keys patches from the tip tree triggering
a newly introduced build-time sanity check on an ARM build, because they changed
the ABI of siginfo in an unexpected way.

If u64 has a natural alignment of 8 bytes (which is the case on most mainstream
platforms, with the notable exception of x86-32), then the leadup to the
_sifields union matters:

typedef struct siginfo {
        int si_signo;
        int si_errno;
        int si_code;

        union {
	...
        } _sifields;
} __ARCH_SI_ATTRIBUTES siginfo_t;

Note how the first 3 fields give us 12 bytes, so _sifields is not 8
naturally bytes aligned.

Before the _pkey field addition the largest element of _sifields (on
32-bit platforms) was 32 bits. With the u64 added, the minimum alignment
requirement increased to 8 bytes on those (rare) 32-bit platforms. Thus
GCC padded the space after si_code with 4 extra bytes, and shifted all
_sifields offsets by 4 bytes - breaking the ABI of all of those
remaining fields.

On 64-bit platforms this problem was hidden due to _sifields already
having numerous fields with natural 8 bytes alignment (pointers).

To fix this, we replace the u64 with an '__u32'.  The __u32 does not
increase the minimum alignment requirement of the union, and it is
also large enough to store the 16-bit pkey we have today on x86.

Reported-by: Stehen Rothwell <sfr@canb.auug.org.au>
Signed-off-by: Dave Hansen <dave.hansen@linux.intel.com>
Acked-by: Stehen Rothwell <sfr@canb.auug.org.au>
Cc: Andrew Morton <akpm@linux-foundation.org>
Cc: Dave Hansen <dave@sr71.net>
Cc: Helge Deller <deller@gmx.de>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: linux-next@vger.kernel.org
Fixes: cd0ea35f ("signals, pkeys: Notify userspace about protection key faults")
Link: http://lkml.kernel.org/r/20160301125451.02C7426D@viggo.jf.intel.com


Signed-off-by: Ingo Molnar <mingo@kernel.org>

This patch add the SO_CNX_ADVICE socket option (setsockopt only). The
purpose is to allow an application to give feedback to the kernel about
the quality of the network path for a connected socket. The value
argument indicates the type of quality report. For this initial patch
the only supported advice is a value of 1 which indicates "bad path,
please reroute"-- the action taken by the kernel is to call
dst_negative_advice which will attempt to choose a different ECMP route,
reset the TX hash for flow label and UDP source port in encapsulation,
etc.

This facility should be useful for connected UDP sockets where only the
application can provide any feedback about path quality. It could also
be useful for TCP applications that have additional knowledge about the
path outside of the normal TCP control loop.

Signed-off-by: Tom Herbert <tom@herbertland.com>
Signed-off-by: David S. Miller <davem@davemloft.net>

A protection key fault is very similar to any other access error.
There must be a VMA, etc...  We even want to take the same action
(SIGSEGV) that we do with a normal access fault.

However, we do need to let userspace know that something is
different.  We do this the same way what we did with SEGV_BNDERR
with Memory Protection eXtensions (MPX): define a new SEGV code:
SEGV_PKUERR.

We add a siginfo field: si_pkey that reveals to userspace which
protection key was set on the PTE that we faulted on.  There is
no other easy way for userspace to figure this out.  They could
parse smaps but that would be a bit cruel.

We share space with in siginfo with _addr_bnd.  #BR faults from
MPX are completely separate from page faults (#PF) that trigger
from protection key violations, so we never need both at the same
time.

Note that _pkey is a 64-bit value.  The current hardware only
supports 4-bit protection keys.  We do this because there is
_plenty_ of space in _sigfault and it is possible that future
processors would support more than 4 bits of protection keys.

The x86 code to actually fill in the siginfo is in the next
patch.

Signed-off-by: Dave Hansen <dave.hansen@linux.intel.com>
Reviewed-by: Thomas Gleixner <tglx@linutronix.de>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Amanieu d'Antras <amanieu@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Cc: Andy Lutomirski <luto@amacapital.net>
Cc: Arnd Bergmann <arnd@arndb.de>
Cc: Borislav Petkov <bp@alien8.de>
Cc: Brian Gerst <brgerst@gmail.com>
Cc: Dave Hansen <dave@sr71.net>
Cc: Denys Vlasenko <dvlasenk@redhat.com>
Cc: H. Peter Anvin <hpa@zytor.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Palmer Dabbelt <palmer@dabbelt.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Richard Weinberger <richard@nod.at>
Cc: Rik van Riel <riel@redhat.com>
Cc: Sasha Levin <sasha.levin@oracle.com>
Cc: Vegard Nossum <vegard.nossum@oracle.com>
Cc: Vladimir Davydov <vdavydov@parallels.com>
Cc: linux-arch@vger.kernel.org
Cc: linux-kernel@vger.kernel.org
Cc: linux-mm@kvack.org
Link: http://lkml.kernel.org/r/20160212210212.3A9B83AC@viggo.jf.intel.com


Signed-off-by: Ingo Molnar <mingo@kernel.org>

For uapi, need try to let all macros have same value, and MADV_FREE is
added into main branch recently, so need redefine MADV_FREE for it.

At present, '8' can be shared with all architectures, so redefine it to
'8'.

[sudipm.mukherjee@gmail.com: correct uniform value of MADV_FREE]
Signed-off-by: Chen Gang <gang.chen.5i5j@gmail.com>
Signed-off-by: Minchan Kim <minchan@kernel.org>
Acked-by: Hugh Dickins <hughd@google.com>
Cc: Ralf Baechle <ralf@linux-mips.org>
Cc: Arnd Bergmann <arnd@arndb.de>
Cc: Richard Henderson <rth@twiddle.net>
Cc: Ivan Kokshaysky <ink@jurassic.park.msu.ru>
Cc: Matt Turner <mattst88@gmail.com>
Cc: "James E.J. Bottomley" <jejb@parisc-linux.org>
Cc: Helge Deller <deller@gmx.de>
Cc: Chris Zankel <chris@zankel.net>
Cc: Max Filippov <jcmvbkbc@gmail.com>
Cc: Roland Dreier <roland@kernel.org>
Cc: Darrick J. Wong <darrick.wong@oracle.com>
Cc: David S. Miller <davem@davemloft.net>
Cc: "Kirill A. Shutemov" <kirill@shutemov.name>
Cc: Shaohua Li <shli@kernel.org>
Cc: <yalin.wang2010@gmail.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Andy Lutomirski <luto@amacapital.net>
Cc: Arnd Bergmann <arnd@arndb.de>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: Catalin Marinas <catalin.marinas@arm.com>
Cc: Daniel Micay <danielmicay@gmail.com>
Cc: Jason Evans <je@fb.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Kirill A. Shutemov <kirill@shutemov.name>
Cc: Mel Gorman <mgorman@suse.de>
Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: Michal Hocko <mhocko@suse.com>
Cc: Mika Penttil <mika.penttila@nextfour.com>
Cc: Rik van Riel <riel@redhat.com>
Cc: Russell King <rmk@arm.linux.org.uk>
Cc: Shaohua Li <shli@kernel.org>
Cc: Will Deacon <will.deacon@arm.com>
Cc: Wu Fengguang <fengguang.wu@intel.com>
Signed-off-by: Sudip Mukherjee <sudip@vectorindia.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

Linux doesn't have an ability to free pages lazy while other OS already
have been supported that named by madvise(MADV_FREE).

The gain is clear that kernel can discard freed pages rather than
swapping out or OOM if memory pressure happens.

Without memory pressure, freed pages would be reused by userspace
without another additional overhead(ex, page fault + allocation +
zeroing).

Jason Evans said:

: Facebook has been using MAP_UNINITIALIZED
: (https://lkml.org/lkml/2012/1/18/308

) in some of its applications for
: several years, but there are operational costs to maintaining this
: out-of-tree in our kernel and in jemalloc, and we are anxious to retire it
: in favor of MADV_FREE.  When we first enabled MAP_UNINITIALIZED it
: increased throughput for much of our workload by ~5%, and although the
: benefit has decreased using newer hardware and kernels, there is still
: enough benefit that we cannot reasonably retire it without a replacement.
:
: Aside from Facebook operations, there are numerous broadly used
: applications that would benefit from MADV_FREE.  The ones that immediately
: come to mind are redis, varnish, and MariaDB.  I don't have much insight
: into Android internals and development process, but I would hope to see
: MADV_FREE support eventually end up there as well to benefit applications
: linked with the integrated jemalloc.
:
: jemalloc will use MADV_FREE once it becomes available in the Linux kernel.
: In fact, jemalloc already uses MADV_FREE or equivalent everywhere it's
: available: *BSD, OS X, Windows, and Solaris -- every platform except Linux
: (and AIX, but I'm not sure it even compiles on AIX).  The lack of
: MADV_FREE on Linux forced me down a long series of increasingly
: sophisticated heuristics for madvise() volume reduction, and even so this
: remains a common performance issue for people using jemalloc on Linux.
: Please integrate MADV_FREE; many people will benefit substantially.

How it works:

When madvise syscall is called, VM clears dirty bit of ptes of the
range.  If memory pressure happens, VM checks dirty bit of page table
and if it found still "clean", it means it's a "lazyfree pages" so VM
could discard the page instead of swapping out.  Once there was store
operation for the page before VM peek a page to reclaim, dirty bit is
set so VM can swap out the page instead of discarding.

One thing we should notice is that basically, MADV_FREE relies on dirty
bit in page table entry to decide whether VM allows to discard the page
or not.  IOW, if page table entry includes marked dirty bit, VM
shouldn't discard the page.

However, as a example, if swap-in by read fault happens, page table
entry doesn't have dirty bit so MADV_FREE could discard the page
wrongly.

For avoiding the problem, MADV_FREE did more checks with PageDirty and
PageSwapCache.  It worked out because swapped-in page lives on swap
cache and since it is evicted from the swap cache, the page has PG_dirty
flag.  So both page flags check effectively prevent wrong discarding by
MADV_FREE.

However, a problem in above logic is that swapped-in page has PG_dirty
still after they are removed from swap cache so VM cannot consider the
page as freeable any more even if madvise_free is called in future.

Look at below example for detail.

    ptr = malloc();
    memset(ptr);
    ..
    ..
    .. heavy memory pressure so all of pages are swapped out
    ..
    ..
    var = *ptr; -> a page swapped-in and could be removed from
                   swapcache. Then, page table doesn't mark
                   dirty bit and page descriptor includes PG_dirty
    ..
    ..
    madvise_free(ptr); -> It doesn't clear PG_dirty of the page.
    ..
    ..
    ..
    .. heavy memory pressure again.
    .. In this time, VM cannot discard the page because the page
    .. has *PG_dirty*

To solve the problem, this patch clears PG_dirty if only the page is
owned exclusively by current process when madvise is called because
PG_dirty represents ptes's dirtiness in several processes so we could
clear it only if we own it exclusively.

Firstly, heavy users would be general allocators(ex, jemalloc, tcmalloc
and hope glibc supports it) and jemalloc/tcmalloc already have supported
the feature for other OS(ex, FreeBSD)

  barrios@blaptop:~/benchmark/ebizzy$ lscpu
  Architecture:          x86_64
  CPU op-mode(s):        32-bit, 64-bit
  Byte Order:            Little Endian
  CPU(s):                12
  On-line CPU(s) list:   0-11
  Thread(s) per core:    1
  Core(s) per socket:    1
  Socket(s):             12
  NUMA node(s):          1
  Vendor ID:             GenuineIntel
  CPU family:            6
  Model:                 2
  Stepping:              3
  CPU MHz:               3200.185
  BogoMIPS:              6400.53
  Virtualization:        VT-x
  Hypervisor vendor:     KVM
  Virtualization type:   full
  L1d cache:             32K
  L1i cache:             32K
  L2 cache:              4096K
  NUMA node0 CPU(s):     0-11
  ebizzy benchmark(./ebizzy -S 10 -n 512)

  Higher avg is better.

   vanilla-jemalloc             MADV_free-jemalloc

  1 thread
  records: 10                   records: 10
  avg:   2961.90                avg:  12069.70
  std:     71.96(2.43%)         std:    186.68(1.55%)
  max:   3070.00                max:  12385.00
  min:   2796.00                min:  11746.00

  2 thread
  records: 10                   records: 10
  avg:   5020.00                avg:  17827.00
  std:    264.87(5.28%)         std:    358.52(2.01%)
  max:   5244.00                max:  18760.00
  min:   4251.00                min:  17382.00

  4 thread
  records: 10                   records: 10
  avg:   8988.80                avg:  27930.80
  std:   1175.33(13.08%)        std:   3317.33(11.88%)
  max:   9508.00                max:  30879.00
  min:   5477.00                min:  21024.00

  8 thread
  records: 10                   records: 10
  avg:  13036.50                avg:  33739.40
  std:    170.67(1.31%)         std:   5146.22(15.25%)
  max:  13371.00                max:  40572.00
  min:  12785.00                min:  24088.00

  16 thread
  records: 10                   records: 10
  avg:  11092.40                avg:  31424.20
  std:    710.60(6.41%)         std:   3763.89(11.98%)
  max:  12446.00                max:  36635.00
  min:   9949.00                min:  25669.00

  32 thread
  records: 10                   records: 10
  avg:  11067.00                avg:  34495.80
  std:    971.06(8.77%)         std:   2721.36(7.89%)
  max:  12010.00                max:  38598.00
  min:   9002.00                min:  30636.00

In summary, MADV_FREE is about much faster than MADV_DONTNEED.

This patch (of 12):

Add core MADV_FREE implementation.

[akpm@linux-foundation.org: small cleanups]
Signed-off-by: Minchan Kim <minchan@kernel.org>
Acked-by: Michal Hocko <mhocko@suse.com>
Acked-by: Hugh Dickins <hughd@google.com>
Cc: Mika Penttil <mika.penttila@nextfour.com>
Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Rik van Riel <riel@redhat.com>
Cc: Mel Gorman <mgorman@suse.de>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Jason Evans <je@fb.com>
Cc: Daniel Micay <danielmicay@gmail.com>
Cc: "Kirill A. Shutemov" <kirill@shutemov.name>
Cc: Shaohua Li <shli@kernel.org>
Cc: <yalin.wang2010@gmail.com>
Cc: Andy Lutomirski <luto@amacapital.net>
Cc: "James E.J. Bottomley" <jejb@parisc-linux.org>
Cc: "Kirill A. Shutemov" <kirill@shutemov.name>
Cc: "Shaohua Li" <shli@kernel.org>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Arnd Bergmann <arnd@arndb.de>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: Catalin Marinas <catalin.marinas@arm.com>
Cc: Chen Gang <gang.chen.5i5j@gmail.com>
Cc: Chris Zankel <chris@zankel.net>
Cc: Darrick J. Wong <darrick.wong@oracle.com>
Cc: David S. Miller <davem@davemloft.net>
Cc: Helge Deller <deller@gmx.de>
Cc: Ivan Kokshaysky <ink@jurassic.park.msu.ru>
Cc: Matt Turner <mattst88@gmail.com>
Cc: Max Filippov <jcmvbkbc@gmail.com>
Cc: Ralf Baechle <ralf@linux-mips.org>
Cc: Richard Henderson <rth@twiddle.net>
Cc: Roland Dreier <roland@kernel.org>
Cc: Russell King <rmk@arm.linux.org.uk>
Cc: Shaohua Li <shli@kernel.org>
Cc: Will Deacon <will.deacon@arm.com>
Cc: Wu Fengguang <fengguang.wu@intel.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

Expose socket options for setting a classic or extended BPF program
for use when selecting sockets in an SO_REUSEPORT group.  These options
can be used on the first socket to belong to a group before bind or
on any socket in the group after bind.

This change includes refactoring of the existing sk_filter code to
allow reuse of the existing BPF filter validation checks.

Signed-off-by: Craig Gallek <kraig@google.com>
Acked-by: Alexei Starovoitov <ast@kernel.org>
Signed-off-by: David S. Miller <davem@davemloft.net>

Add a copy_file_range() system call for offloading copies between
regular files.

This gives an interface to underlying layers of the storage stack which
can copy without reading and writing all the data.  There are a few
candidates that should support copy offloading in the nearer term:

- btrfs shares extent references with its clone ioctl
- NFS has patches to add a COPY command which copies on the server
- SCSI has a family of XCOPY commands which copy in the device

This system call avoids the complexity of also accelerating the creation
of the destination file by operating on an existing destination file
descriptor, not a path.

Currently the high level vfs entry point limits copy offloading to files
on the same mount and super (and not in the same file).  This can be
relaxed if we get implementations which can copy between file systems
safely.

Signed-off-by: Zach Brown <zab@redhat.com>
[Anna Schumaker: Change -EINVAL to -EBADF during file verification,
                 Change flags parameter from int to unsigned int,
                 Add function to include/linux/syscalls.h,
                 Check copy len after file open mode,
                 Don't forbid ranges inside the same file,
                 Use rw_verify_area() to veriy ranges,
                 Use file_out rather than file_in,
                 Add COPY_FR_REFLINK flag]
Signed-off-by: Anna Schumaker <Anna.Schumaker@Netapp.com>
Reviewed-by: Christoph Hellwig <hch@lst.de>
Signed-off-by: Al Viro <viro@zeniv.linux.org.uk>

The previous patch introduced a flag that specified pages in a VMA should
be placed on the unevictable LRU, but they should not be made present when
the area is created.  This patch adds the ability to set this state via
the new mlock system calls.

We add MLOCK_ONFAULT for mlock2 and MCL_ONFAULT for mlockall.
MLOCK_ONFAULT will set the VM_LOCKONFAULT modifier for VM_LOCKED.
MCL_ONFAULT should be used as a modifier to the two other mlockall flags.
When used with MCL_CURRENT, all current mappings will be marked with
VM_LOCKED | VM_LOCKONFAULT.  When used with MCL_FUTURE, the mm->def_flags
will be marked with VM_LOCKED | VM_LOCKONFAULT.  When used with both
MCL_CURRENT and MCL_FUTURE, all current mappings and mm->def_flags will be
marked with VM_LOCKED | VM_LOCKONFAULT.

Prior to this patch, mlockall() will unconditionally clear the
mm->def_flags any time it is called without MCL_FUTURE.  This behavior is
maintained after adding MCL_ONFAULT.  If a call to mlockall(MCL_FUTURE) is
followed by mlockall(MCL_CURRENT), the mm->def_flags will be cleared and
new VMAs will be unlocked.  This remains true with or without MCL_ONFAULT
in either mlockall() invocation.

munlock() will unconditionally clear both vma flags.  munlockall()
unconditionally clears for VMA flags on all VMAs and in the mm->def_flags
field.

Signed-off-by: Eric B Munson <emunson@akamai.com>
Acked-by: Michal Hocko <mhocko@suse.com>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Cc: Jonathan Corbet <corbet@lwn.net>
Cc: Catalin Marinas <catalin.marinas@arm.com>
Cc: Geert Uytterhoeven <geert@linux-m68k.org>
Cc: Guenter Roeck <linux@roeck-us.net>
Cc: Heiko Carstens <heiko.carstens@de.ibm.com>
Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com>
Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: Ralf Baechle <ralf@linux-mips.org>
Cc: Shuah Khan <shuahkh@osg.samsung.com>
Cc: Stephen Rothwell <sfr@canb.auug.org.au>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

With the refactored mlock code, introduce a new system call for mlock.
The new call will allow the user to specify what lock states are being
added.  mlock2 is trivial at the moment, but a follow on patch will add a
new mlock state making it useful.

Signed-off-by: Eric B Munson <emunson@akamai.com>
Acked-by: Michal Hocko <mhocko@suse.com>
Acked-by: Vlastimil Babka <vbabka@suse.cz>
Cc: Heiko Carstens <heiko.carstens@de.ibm.com>
Cc: Geert Uytterhoeven <geert@linux-m68k.org>
Cc: Catalin Marinas <catalin.marinas@arm.com>
Cc: Stephen Rothwell <sfr@canb.auug.org.au>
Cc: Guenter Roeck <linux@roeck-us.net>
Cc: Jonathan Corbet <corbet@lwn.net>
Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com>
Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: Ralf Baechle <ralf@linux-mips.org>
Cc: Shuah Khan <shuahkh@osg.samsung.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

MINSIGSTKSZ and SIGSTKSZ for ARM64 are not correctly set in latest kernel.
This patch fixes this issue.

This issue is reported in LTP (testcase: sigaltstack02.c).
Testcase failed when sigaltstack() called with stack size "MINSIGSTKSZ - 1"
Since in Glibc-2.22, MINSIGSTKSZ is set to 5120 but in kernel
it is set to 2048 so testcase gets failed.

Testcase Output:
sigaltstack02 1  TPASS  :  stgaltstack() fails, Invalid Flag value,errno:22
sigaltstack02 2  TFAIL  :  sigaltstack() returned 0, expected -1,errno:12

Reported Issue in Glibc Bugzilla:
Bugfix in Glibc-2.22: [Bug 16850]
https://sourceware.org/bugzilla/show_bug.cgi?id=16850



Acked-by: Arnd Bergmann <arnd@arndb.de>
Signed-off-by: Akhilesh Kumar <akhilesh.k@samsung.com>
Signed-off-by: Manjeet Pawar <manjeet.p@samsung.com>
Signed-off-by: Rohit Thapliyal <r.thapliyal@samsung.com>
Signed-off-by: Will Deacon <will.deacon@arm.com>

Add the userfaultfd syscalls to uapi asm-generic, it was tested with
postcopy live migration on aarch64 with both 4k and 64k pagesize
kernels.

Signed-off-by: Dr. David Alan Gilbert <dgilbert@redhat.com>
Signed-off-by: Andrea Arcangeli <aarcange@redhat.com>
Cc: Michael Ellerman <mpe@ellerman.id.au>
Cc: Shuah Khan <shuahkh@osg.samsung.com>
Cc: Thierry Reding <treding@nvidia.com>
Cc: Mathieu Desnoyers <mathieu.desnoyers@efficios.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

Here is an implementation of a new system call, sys_membarrier(), which
executes a memory barrier on all threads running on the system.  It is
implemented by calling synchronize_sched().  It can be used to
distribute the cost of user-space memory barriers asymmetrically by
transforming pairs of memory barriers into pairs consisting of
sys_membarrier() and a compiler barrier.  For synchronization primitives
that distinguish between read-side and write-side (e.g.  userspace RCU
[1], rwlocks), the read-side can be accelerated significantly by moving
the bulk of the memory barrier overhead to the write-side.

The existing applications of which I am aware that would be improved by
this system call are as follows:

* Through Userspace RCU library (http://urcu.so)
  - DNS server (Knot DNS) https://www.knot-dns.cz/
  - Network sniffer (http://netsniff-ng.org/)
  - Distributed object storage (https://sheepdog.github.io/sheepdog/)
  - User-space tracing (http://lttng.org)
  - Network storage system (https://www.gluster.org/)
  - Virtual routers (https://events.linuxfoundation.org/sites/events/files/slides/DPDK_RCU_0MQ.pdf)
  - Financial software (https://lkml.org/lkml/2015/3/23/189)

Those projects use RCU in userspace to increase read-side speed and
scalability compared to locking.  Especially in the case of RCU used by
libraries, sys_membarrier can speed up the read-side by moving the bulk of
the memory barrier cost to synchronize_rcu().

* Direct users of sys_membarrier
  - core dotnet garbage collector (https://github.com/dotnet/coreclr/issues/198)

Microsoft core dotnet GC developers are planning to use the mprotect()
side-effect of issuing memory barriers through IPIs as a way to implement
Windows FlushProcessWriteBuffers() on Linux.  They are referring to
sys_membarrier in their github thread, specifically stating that
sys_membarrier() is what they are looking for.

To explain the benefit of this scheme, let's introduce two example threads:

Thread A (non-frequent, e.g. executing liburcu synchronize_rcu())
Thread B (frequent, e.g. executing liburcu
rcu_read_lock()/rcu_read_unlock())

In a scheme where all smp_mb() in thread A are ordering memory accesses
with respect to smp_mb() present in Thread B, we can change each
smp_mb() within Thread A into calls to sys_membarrier() and each
smp_mb() within Thread B into compiler barriers "barrier()".

Before the change, we had, for each smp_mb() pairs:

Thread A                    Thread B
previous mem accesses       previous mem accesses
smp_mb()                    smp_mb()
following mem accesses      following mem accesses

After the change, these pairs become:

Thread A                    Thread B
prev mem accesses           prev mem accesses
sys_membarrier()            barrier()
follow mem accesses         follow mem accesses

As we can see, there are two possible scenarios: either Thread B memory
accesses do not happen concurrently with Thread A accesses (1), or they
do (2).

1) Non-concurrent Thread A vs Thread B accesses:

Thread A                    Thread B
prev mem accesses
sys_membarrier()
follow mem accesses
                            prev mem accesses
                            barrier()
                            follow mem accesses

In this case, thread B accesses will be weakly ordered. This is OK,
because at that point, thread A is not particularly interested in
ordering them with respect to its own accesses.

2) Concurrent Thread A vs Thread B accesses

Thread A                    Thread B
prev mem accesses           prev mem accesses
sys_membarrier()            barrier()
follow mem accesses         follow mem accesses

In this case, thread B accesses, which are ensured to be in program
order thanks to the compiler barrier, will be "upgraded" to full
smp_mb() by synchronize_sched().

* Benchmarks

On Intel Xeon E5405 (8 cores)
(one thread is calling sys_membarrier, the other 7 threads are busy
looping)

1000 non-expedited sys_membarrier calls in 33s =3D 33 milliseconds/call.

* User-space user of this system call: Userspace RCU library

Both the signal-based and the sys_membarrier userspace RCU schemes
permit us to remove the memory barrier from the userspace RCU
rcu_read_lock() and rcu_read_unlock() primitives, thus significantly
accelerating them. These memory barriers are replaced by compiler
barriers on the read-side, and all matching memory barriers on the
write-side are turned into an invocation of a memory barrier on all
active threads in the process. By letting the kernel perform this
synchronization rather than dumbly sending a signal to every process
threads (as we currently do), we diminish the number of unnecessary wake
ups and only issue the memory barriers on active threads. Non-running
threads do not need to execute such barrier anyway, because these are
implied by the scheduler context switches.

Results in liburcu:

Operations in 10s, 6 readers, 2 writers:

memory barriers in reader:    1701557485 reads, 2202847 writes
signal-based scheme:          9830061167 reads,    6700 writes
sys_membarrier:               9952759104 reads,     425 writes
sys_membarrier (dyn. check):  7970328887 reads,     425 writes

The dynamic sys_membarrier availability check adds some overhead to
the read-side compared to the signal-based scheme, but besides that,
sys_membarrier slightly outperforms the signal-based scheme. However,
this non-expedited sys_membarrier implementation has a much slower grace
period than signal and memory barrier schemes.

Besides diminishing the number of wake-ups, one major advantage of the
membarrier system call over the signal-based scheme is that it does not
need to reserve a signal. This plays much more nicely with libraries,
and with processes injected into for tracing purposes, for which we
cannot expect that signals will be unused by the application.

An expedited version of this system call can be added later on to speed
up the grace period. Its implementation will likely depend on reading
the cpu_curr()->mm without holding each CPU's rq lock.

This patch adds the system call to x86 and to asm-generic.

[1] http://urcu.so



membarrier(2) man page:

MEMBARRIER(2)              Linux Programmer's Manual             MEMBARRIER(2)

NAME
       membarrier - issue memory barriers on a set of threads

SYNOPSIS
       #include <linux/membarrier.h>

       int membarrier(int cmd, int flags);

DESCRIPTION
       The cmd argument is one of the following:

       MEMBARRIER_CMD_QUERY
              Query  the  set  of  supported commands. It returns a bitmask of
              supported commands.

       MEMBARRIER_CMD_SHARED
              Execute a memory barrier on all threads running on  the  system.
              Upon  return from system call, the caller thread is ensured that
              all running threads have passed through a state where all memory
              accesses  to  user-space  addresses  match program order between
              entry to and return from the system  call  (non-running  threads
              are de facto in such a state). This covers threads from all pro=E2=80=90
              cesses running on the system.  This command returns 0.

       The flags argument needs to be 0. For future extensions.

       All memory accesses performed  in  program  order  from  each  targeted
       thread is guaranteed to be ordered with respect to sys_membarrier(). If
       we use the semantic "barrier()" to represent a compiler barrier forcing
       memory  accesses  to  be performed in program order across the barrier,
       and smp_mb() to represent explicit memory barriers forcing full  memory
       ordering  across  the barrier, we have the following ordering table for
       each pair of barrier(), sys_membarrier() and smp_mb():

       The pair ordering is detailed as (O: ordered, X: not ordered):

                              barrier()   smp_mb() sys_membarrier()
              barrier()          X           X            O
              smp_mb()           X           O            O
              sys_membarrier()   O           O            O

RETURN VALUE
       On success, these system calls return zero.  On error, -1 is  returned,
       and errno is set appropriately. For a given command, with flags
       argument set to 0, this system call is guaranteed to always return the
       same value until reboot.

ERRORS
       ENOSYS System call is not implemented.

       EINVAL Invalid arguments.

Linux                             2015-04-15                     MEMBARRIER(2)

Signed-off-by: Mathieu Desnoyers <mathieu.desnoyers@efficios.com>
Reviewed-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Reviewed-by: Josh Triplett <josh@joshtriplett.org>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Nicholas Miell <nmiell@comcast.net>
Cc: Ingo Molnar <mingo@redhat.com>
Cc: Alan Cox <gnomes@lxorguk.ukuu.org.uk>
Cc: Lai Jiangshan <laijs@cn.fujitsu.com>
Cc: Stephen Hemminger <stephen@networkplumber.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: David Howells <dhowells@redhat.com>
Cc: Pranith Kumar <bobby.prani@gmail.com>
Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: Shuah Khan <shuahkh@osg.samsung.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

ENOSYS is the mechanism used by user code to detect whether the running
kernel implements a given system call.  It should not be returned by
anything except an unimplemented system call.

Unfortunately, it is rather frequently used in the kernel to indicate that
various new functions of existing system calls are not implemented.  This
should be discouraged.

Improve the comment in errno.h to help clarify ENOSYS's purpose.

Signed-off-by: Andy Lutomirski <luto@amacapital.net>
Cc: Pavel Machek <pavel@ucw.cz>
Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: Joe Perches <joe@perches.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

Fix clashing values for O_PATH and FMODE_NONOTIFY on sparc.  The
clashing O_PATH value was added in commit 5229645b ("vfs: add
nonconflicting values for O_PATH") but this can't be changed as it is
user-visible.

FMODE_NONOTIFY is only used internally in the kernel, but it is in the
same numbering space as the other O_* flags, as indicated by the comment
at the top of include/uapi/asm-generic/fcntl.h (and its use in
fs/notify/fanotify/fanotify_user.c).  So renumber it to avoid the clash.

All of this has happened before (commit 12ed2e36: "fanotify:
FMODE_NONOTIFY and __O_SYNC in sparc conflict"), and all of this will
happen again -- so update the uniqueness check in fcntl_init() to
include __FMODE_NONOTIFY.

Signed-off-by: David Drysdale <drysdale@google.com>
Acked-by: David S. Miller <davem@davemloft.net>
Acked-by: Jan Kara <jack@suse.cz>
Cc: Heinrich Schuchardt <xypron.glpk@gmx.de>
Cc: Alexander Viro <viro@zeniv.linux.org.uk>
Cc: Arnd Bergmann <arnd@arndb.de>
Cc: Stephen Rothwell <sfr@canb.auug.org.au>
Cc: Eric Paris <eparis@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

This patchset adds execveat(2) for x86, and is derived from Meredydd
Luff's patch from Sept 2012 (https://lkml.org/lkml/2012/9/11/528).

The primary aim of adding an execveat syscall is to allow an
implementation of fexecve(3) that does not rely on the /proc filesystem,
at least for executables (rather than scripts).  The current glibc version
of fexecve(3) is implemented via /proc, which causes problems in sandboxed
or otherwise restricted environments.

Given the desire for a /proc-free fexecve() implementation, HPA suggested
(https://lkml.org/lkml/2006/7/11/556) that an execveat(2) syscall would be
an appropriate generalization.

Also, having a new syscall means that it can take a flags argument without
back-compatibility concerns.  The current implementation just defines the
AT_EMPTY_PATH and AT_SYMLINK_NOFOLLOW flags, but other flags could be
added in future -- for example, flags for new namespaces (as suggested at
https://lkml.org/lkml/2006/7/11/474).

Rela...

introduce new setsockopt() command:

setsockopt(sock, SOL_SOCKET, SO_ATTACH_BPF, &prog_fd, sizeof(prog_fd))

where prog_fd was received from syscall bpf(BPF_PROG_LOAD, attr, ...)
and attr->prog_type == BPF_PROG_TYPE_SOCKET_FILTER

setsockopt() calls bpf_prog_get() which increments refcnt of the program,
so it doesn't get unloaded while socket is using the program.

The same eBPF program can be attached to multiple sockets.

User task exit automatically closes socket which calls sk_filter_uncharge()
which decrements refcnt of eBPF program

Signed-off-by: Alexei Starovoitov <ast@plumgrid.com>
Signed-off-by: David S. Miller <davem@davemloft.net>

This patch adds new fields about bound violation into siginfo
structure. si_lower and si_upper are respectively lower bound
and upper bound when bound violation is caused.

Signed-off-by: Qiaowei Ren <qiaowei.ren@intel.com>
Signed-off-by: Dave Hansen <dave.hansen@linux.intel.com>
Cc: linux-mm@kvack.org
Cc: linux-mips@linux-mips.org
Cc: Dave Hansen <dave@sr71.net>
Link: http://lkml.kernel.org/r/20141114151819.1908C900@viggo.jf.intel.com


Signed-off-by: Thomas Gleixner <tglx@linutronix.de>

Alternative to RPS/RFS is to use hardware support for multiple
queues.

Then split a set of million of sockets into worker threads, each
one using epoll() to manage events on its own socket pool.

Ideally, we want one thread per RX/TX queue/cpu, but we have no way to
know after accept() or connect() on which queue/cpu a socket is managed.

We normally use one cpu per RX queue (IRQ smp_affinity being properly
set), so remembering on socket structure which cpu delivered last packet
is enough to solve the problem.

After accept(), connect(), or even file descriptor passing around
processes, applications can use :

 int cpu;
 socklen_t len = sizeof(cpu);

 getsockopt(fd, SOL_SOCKET, SO_INCOMING_CPU, &cpu, &len);

And use this information to put the socket into the right silo
for optimal performance, as all networking stack should run
on the appropriate cpu, without need to send IPI (RPS/RFS).

Signed-off-by: Eric Dumazet <edumazet@google.com>
Signed-off-by: David S. Miller <davem@davemloft.net>

done as separate commit to ease conflict resolution

Signed-off-by: Alexei Starovoitov <ast@plumgrid.com>
Signed-off-by: David S. Miller <davem@davemloft.net>

Commit 9183df25 ("shm: add memfd_create() syscall") added a new
system call (memfd_create) but didn't update the asm-generic unistd
header.

This patch adds the new system call to the asm-generic version of
unistd.h so that it can be used by architectures such as arm64.

Cc: Arnd Bergmann <arnd@arndb.de>
Reviewed-by: David Herrmann <dh.herrmann@gmail.com>
Signed-off-by: Will Deacon <will.deacon@arm.com>

The getrandom(2) system call was requested by the LibreSSL Portable
developers.  It is analoguous to the getentropy(2) system call in
OpenBSD.

The rationale of this system call is to provide resiliance against
file descriptor exhaustion attacks, where the attacker consumes all
available file descriptors, forcing the use of the fallback code where
/dev/[u]random is not available.  Since the fallback code is often not
well-tested, it is better to eliminate this potential failure mode
entirely.

The other feature provided by this new system call is the ability to
request randomness from the /dev/urandom entropy pool, but to block
until at least 128 bits of entropy has been accumulated in the
/dev/urandom entropy pool.  Historically, the emphasis in the
/dev/urandom development has been to ensure that urandom pool is
initialized as quickly as possible after system boot, and preferably
before the init scripts start execution.

This is because changing /dev/urandom reads to block represents an
interface change that could potentially break userspace which is not
acceptable.  In practice, on most x86 desktop and server systems, in
general the entropy pool can be initialized before it is needed (and
in modern kernels, we will printk a warning message if not).  However,
on an embedded system, this may not be the case.  And so with this new
interface, we can provide the functionality of blocking until the
urandom pool has been initialized.  Any userspace program which uses
this new functionality must take care to assure that if it is used
during the boot process, that it will not cause the init scripts or
other portions of the system startup to hang indefinitely.

SYNOPSIS
	#include <linux/random.h>

	int getrandom(void *buf, size_t buflen, unsigned int flags);

DESCRIPTION
	The system call getrandom() fills the buffer pointed to by buf
	with up to buflen random bytes which can be used to seed user
	space random number generators (i.e., DRBG's) or for other
	cryptographic uses.  It should not be used for Monte Carlo
	simulations or other programs/algorithms which are doing
	probabilistic sampling.

	If the GRND_RANDOM flags bit is set, then draw from the
	/dev/random pool instead of the /dev/urandom pool.  The
	/dev/random pool is limited based on the entropy that can be
	obtained from environmental noise, so if there is insufficient
	entropy, the requested number of bytes may not be returned.
	If there is no entropy available at all, getrandom(2) will
	either block, or return an error with errno set to EAGAIN if
	the GRND_NONBLOCK bit is set in flags.

	If the GRND_RANDOM bit is not set, then the /dev/urandom pool
	will be used.  Unlike using read(2) to fetch data from
	/dev/urandom, if the urandom pool has not been sufficiently
	initialized, getrandom(2) will block (or return -1 with the
	errno set to EAGAIN if the GRND_NONBLOCK bit is set in flags).

	The getentropy(2) system call in OpenBSD can be emulated using
	the following function:

            int getentropy(void *buf, size_t buflen)
            {
                    int     ret;

                    if (buflen > 256)
                            goto failure;
                    ret = getrandom(buf, buflen, 0);
                    if (ret < 0)
                            return ret;
                    if (ret == buflen)
                            return 0;
            failure:
                    errno = EIO;
                    return -1;
            }

RETURN VALUE
       On success, the number of bytes that was filled in the buf is
       returned.  This may not be all the bytes requested by the
       caller via buflen if insufficient entropy was present in the
       /dev/random pool, or if the system call was interrupted by a
       signal.

       On error, -1 is returned, and errno is set appropriately.

ERRORS
	EINVAL		An invalid flag was passed to getrandom(2)

	EFAULT		buf is outside the accessible address space.

	EAGAIN		The requested entropy was not available, and
			getentropy(2) would have blocked if the
			GRND_NONBLOCK flag was not set.

	EINTR		While blocked waiting for entropy, the call was
			interrupted by a signal handler; see the description
			of how interrupted read(2) calls on "slow" devices
			are handled with and without the SA_RESTART flag
			in the signal(7) man page.

NOTES
	For small requests (buflen <= 256) getrandom(2) will not
	return EINTR when reading from the urandom pool once the
	entropy pool has been initialized, and it will return all of
	the bytes that have been requested.  This is the recommended
	way to use getrandom(2), and is designed for compatibility
	with OpenBSD's getentropy() system call.

	However, if you are using GRND_RANDOM, then getrandom(2) may
	block until the entropy accounting determines that sufficient
	environmental noise has been gathered such that getrandom(2)
	will be operating as a NRBG instead of a DRBG for those people
	who are working in the NIST SP 800-90 regime.  Since it may
	block for a long time, these guarantees do *not* apply.  The
	user may want to interrupt a hanging process using a signal,
	so blocking until all of the requested bytes are returned
	would be unfriendly.

	For this reason, the user of getrandom(2) MUST always check
	the return value, in case it returns some error, or if fewer
	bytes than requested was returned.  In the case of
	!GRND_RANDOM and small request, the latter should never
	happen, but the careful userspace code (and all crypto code
	should be careful) should check for this anyway!

	Finally, unless you are doing long-term key generation (and
	perhaps not even then), you probably shouldn't be using
	GRND_RANDOM.  The cryptographic algorithms used for
	/dev/urandom are quite conservative, and so should be
	sufficient for all purposes.  The disadvantage of GRND_RANDOM
	is that it can block, and the increased complexity required to
	deal with partially fulfilled getrandom(2) requests.

Signed-off-by: Theodore Ts'o <tytso@mit.edu>
Reviewed-by: Zach Brown <zab@zabbo.net>

This adds the new "seccomp" syscall with both an "operation" and "flags"
parameter for future expansion. The third argument is a pointer value,
used with the SECCOMP_SET_MODE_FILTER operation. Currently, flags must
be 0. This is functionally equivalent to prctl(PR_SET_SECCOMP, ...).

In addition to the TSYNC flag later in this patch series, there is a
non-zero chance that this syscall could be used for configuring a fixed
argument area for seccomp-tracer-aware processes to pass syscall arguments
in the future. Hence, the use of "seccomp" not simply "seccomp_add_filter"
for this syscall. Additionally, this syscall uses operation, flags,
and user pointer for arguments because strictly passing arguments via
a user pointer would mean seccomp itself would be unable to trivially
filter the seccomp syscall itself.

Signed-off-by: Kees Cook <keescook@chromium.org>
Reviewed-by: Oleg Nesterov <oleg@redhat.com>
Reviewed-by: Andy Lutomirski <luto@amacapital.net>

Add the renameat2 syscall to the generic syscall list, which is used by the
following architectures: arc, arm64, c6x, hexagon, metag, openrisc, score,
tile, unicore32.

Signed-off-by: James Hogan <james.hogan@imgtec.com>
Acked-by: Arnd Bergmann <arnd@arndb.de>
Signed-off-by: Miklos Szeredi <mszeredi@suse.cz>
Cc: linux-arch@vger.kernel.org
Cc: Vineet Gupta <vgupta@synopsys.com>
Cc: Catalin Marinas <catalin.marinas@arm.com>
Cc: Will Deacon <will.deacon@arm.com>
Cc: linux-arm-kernel@lists.infradead.org
Cc: Mark Salter <msalter@redhat.com>
Cc: Aurelien Jacquiot <a-jacquiot@ti.com>
Cc: Richard Kuo <rkuo@codeaurora.org>
Cc: linux-hexagon@vger.kernel.org
Cc: linux-metag@vger.kernel.org
Cc: Jonas Bonn <jonas@southpole.se>
Cc: Chen Liqin <liqin.linux@gmail.com>
Cc: Lennox Wu <lennox.wu@gmail.com>
Cc: Chris Metcalf <cmetcalf@tilera.com>
Cc: Guan Xuetao <gxt@mprc.pku.edu.cn>

_STK_LIM_MAX could be used to override the RLIMIT_STACK hard limit from
an arch's include/uapi/asm-generic/resource.h file, but is no longer
used since both parisc and metag removed the override. Therefore remove
it entirely, setting the hard RLIMIT_STACK limit to RLIM_INFINITY
directly in include/asm-generic/resource.h.

Signed-off-by: James Hogan <james.hogan@imgtec.com>
Cc: Arnd Bergmann <arnd@arndb.de>
Cc: linux-arch@vger.kernel.org
Cc: Helge Deller <deller@gmx.de>
Cc: John David Anglin <dave.anglin@bell.net>

File-private locks have been merged into Linux for v3.15, and *now*
people are commenting that the name and macro definitions for the new
file-private locks suck.

...and I can't even disagree. The names and command macros do suck.

We're going to have to live with these for a long time, so it's
important that we be happy with the names before we're stuck with them.
The consensus on the lists so far is that they should be rechristened as
"open file description locks".

The name isn't a big deal for the kernel, but the command macros are not
visually distinct enough from the traditional POSIX lock macros. The
glibc and documentation folks are recommending that we change them to
look like F_OFD_{GETLK|SETLK|SETLKW}. That lessens the chance that a
programmer will typo one of the commands wrong, and also makes it easier
to spot this difference when reading code.

This patch makes the following changes that I think are necessary before
v3.15 ships:

1) rename the command macros to their new names. These end up in the uapi
   headers and so are part of the external-facing API. It turns out that
   glibc doesn't actually use the fcntl.h uapi header, but it's hard to
   be sure that something else won't. Changing it now is safest.

2) make the the /proc/locks output display these as type "OFDLCK"

Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: Christoph Hellwig <hch@infradead.org>
Cc: Carlos O'Donell <carlos@redhat.com>
Cc: Stefan Metzmacher <metze@samba.org>
Cc: Andy Lutomirski <luto@amacapital.net>
Cc: Frank Filz <ffilzlnx@mindspring.com>
Cc: Theodore Ts'o <tytso@mit.edu>
Signed-off-by: Jeff Layton <jlayton@redhat.com>

s/MADV_NODUMP/MADV_DONTDUMP/

Signed-off-by: Zhang Yanfei <zhangyanfei@cn.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>

Due to some unfortunate history, POSIX locks have very strange and
unhelpful semantics. The thing that usually catches people by surprise
is that they are dropped whenever the process closes any file descriptor
associated with the inode.

This is extremely problematic for people developing file servers that
need to implement byte-range locks. Developers often need a "lock
management" facility to ensure that file descriptors are not closed
until all of the locks associated with the inode are finished.

Additionally, "classic" POSIX locks are owned by the process. Locks
taken between threads within the same process won't conflict with one
another, which renders them useless for synchronization between threads.

This patchset adds a new type of lock that attempts to address these
issues. These locks conflict with classic POSIX read/write locks, but
have semantics that are more like BSD locks with respect to inheritance
and behavior on close.

This is implemented primarily by changing how fl_owner field is set for
these locks. Instead of having them owned by the files_struct of the
process, they are instead owned by the filp on which they were acquired.
Thus, they are inherited across fork() and are only released when the
last reference to a filp is put.

These new semantics prevent them from being merged with classic POSIX
locks, even if they are acquired by the same process. These locks will
also conflict with classic POSIX locks even if they are acquired by
the same process or on the same file descriptor.

The new locks are managed using a new set of cmd values to the fcntl()
syscall. The initial implementation of this converts these values to
"classic" cmd values at a fairly high level, and the details are not
exposed to the underlying filesystem. We may eventually want to push
this handing out to the lower filesystem code but for now I don't
see any need for it.

Also, note that with this implementation the new cmd values are only
available via fcntl64() on 32-bit arches. There's little need to
add support for legacy apps on a new interface like this.

Signed-off-by: Jeff Layton <jlayton@redhat.com>

In the 32-bit case fcntl assigns the 64-bit f_pos and i_size to a 32-bit
off_t.

The existing range checks also seem to depend on signed arithmetic
wrapping when it overflows.  In practice maybe that works, but we can be
more careful.  That also allows us to make a more reliable distinction
between -EINVAL and -EOVERFLOW.

Note that in the 32-bit case SEEK_CUR or SEEK_END might allow the caller
to set a lock with starting point no longer representable as a 32-bit
value.  We could return -EOVERFLOW in such cases, but the locks code is
capable of handling such ranges, so we choose to be lenient here.  The
only problem is that subsequent GETLK calls on such a lock will fail
with EOVERFLOW.

While we're here, do some cleanup including consolidating code for the
flock and flock64 cases.

Signed-off-by: J. Bruce Fields <bfields@fieldses.org>
Signed-off-by: Jeff Layton <jlayton@redhat.com>

For architecture dependent compat syscalls in common code an architecture
must define something like __ARCH_WANT_<WHATEVER> if it wants to use the
code.
This however is not true for compat_sys_getdents64 for which architectures
must define __ARCH_OMIT_COMPAT_SYS_GETDENTS64 if they do not want the code.

This leads to the situation where all architectures, except mips, get the
compat code but only x86_64, arm64 and the generic syscall architectures
actually use it.

So invert the logic, so that architectures actively must do something to
get the compat code.

This way a couple of architectures get rid of otherwise dead code.

Signed-off-by: Heiko Carstens <heiko.carstens@de.ibm.com>

Add the sched_setattr and sched_getattr syscalls to the generic syscall
list, which is used by the following architectures: arc, arm64, c6x,
hexagon, metag, openrisc, score, tile, unicore32.

Signed-off-by: James Hogan <james.hogan@imgtec.com>
Acked-by: Arnd Bergmann <arnd@arndb.de>
Acked-by: Catalin Marinas <catalin.marinas@arm.com>
Cc: linux-arch@vger.kernel.org
Cc: Vineet Gupta <vgupta@synopsys.com>
Cc: Will Deacon <will.deacon@arm.com>
Cc: linux-arm-kernel@lists.infradead.org
Cc: Mark Salter <msalter@redhat.com>
Cc: Aurelien Jacquiot <a-jacquiot@ti.com>
Cc: linux-c6x-dev@linux-c6x.org
Cc: Richard Kuo <rkuo@codeaurora.org>
Cc: linux-hexagon@vger.kernel.org
Cc: linux-metag@vger.kernel.org
Cc: Jonas Bonn <jonas@southpole.se>
Cc: linux@lists.openrisc.net
Cc: Chen Liqin <liqin.linux@gmail.com>
Cc: Lennox Wu <lennox.wu@gmail.com>
Cc: Chris Metcalf <cmetcalf@tilera.com>
Cc: Guan Xuetao <gxt@mprc.pku.edu.cn>

Now all 64-bit architectures have been converted to int-ll64.h, we can
remove int-l64.h in kernelspace.

For backwards compatibility, alpha, ia64, mips64, and powerpc64 still
use int-l64.h in userspace.

This is the (reworked for UAPI) non-documentation part of more than two
year old "asm/types.h: All architectures use int-ll64.h in kernelspace"
(https://lkml.org/lkml/2011/8/13/104

)

Since <asm/types.h> (from include/uapi/asm-generic/types.h) is used for
both kernel and user space, include/asm-generic/int-ll64.h cannot just
become include/asm-generic/types.h, as Arnd suggested.

Signed-off-by: Geert Uytterhoeven <geert@linux-m68k.org>
Acked-by: Arnd Bergmann <arnd@arndb.de>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Randy Dunlap <rdunlap@xenotime.net>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>