| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| In the Linux kernel, the following vulnerability has been resolved:
signal: Prevent exec() race
Hyunwoo debugged the following KASAN UAF splat:
BUG: KASAN: slab-use-after-free in __send_signal_locked+0xb27/0xba0
Write of size 8 at addr ffff888007ed80c8 by task poc/79
...
Call Trace:
__send_signal_locked+0xb27/0xba0
do_send_sig_info+0xa7/0x160
do_send_specific+0x76/0xa0
__x64_sys_tgkill+0x193/0x270
...
Allocated by task 80:
do_timer_create+0x1a4/0x1030
__x64_sys_timer_create+0x145/0x190
...
Freed by task 12:
kmem_cache_free_bulk+0x1f8/0x4a0
kvfree_rcu_bulk+0x14f/0x1c0
kfree_rcu_work+0x128/0x1a0
...
Last potentially related work creation:
kvfree_call_rcu+0x39/0x390
__flush_itimer_signals+0x211/0x320
flush_itimer_signals+0x47/0x90
begin_new_exec+0xa6b/0x28c0
It turned out that this happens with a non-leader exec() as Hyunwoo
explained:
de_thread() calls exchange_tids() before release_task(leader), so the
struct pid held by a SIGEV_THREAD_ID timer created against the leader's tid
now points to the thread which called execve(). pid_task() returns that
thread and lock_task_sighand() on it succeeds.
If the timer signal is blocked, its sigqueue stays queued on the leader's
task::pending. The next expiry of that timer can then run while
release_task() flushes the queue.
posixtimer_send_sigqueue() checks whether the sigqueue is already queued
with a plain list_empty(), which only reads list_head::next.
list_del_init() is not atomic and INIT_LIST_HEAD() stores list_head::next
before list_head::prev, so the check can pass in between. list_add_tail()
queues the entry on the task::pending of the live thread, and the
list_head::prev store from the flush then overwrites the list_head::prev
link that list_add_tail() has just set.
__flush_itimer_signals() does not undo that either. With list_head::prev
pointing at the entry itself, its list_del_init() only stores the same
values again, so the entry is not removed from the list. It is still there
after the last reference is dropped and the timer is freed by RCU, and the
list_add_tail() of a later tgkill() follows that list_head::prev into the
freed timer.
This problem surfaced with the recent commit which moved the sigqueue flush
out of the sighand lock held region.
Hyonwoo proposed to fix this by using list_del_init_careful(), but that
just papers over the problem. After some disucssions and various attempts
to solve it, Eric pointed out that there is no reason to flush
task::pending late in release_task() and it should be done in
exit_signals() already.
As nothing can collect and deliver signals which are queued in a dying
task's pending queue, there is no reason to delay it further.
But it has to be ensured that no signals can be queued into it after that
point. exit_signals() sets PF_EXITING in task::flags, which can be used as
an indicator for this.
Cure it by:
- Preventing signal queueing for task private signals (PIDTYPE_PID) when
the task has PF_EXITING set in __send_signal_locked() and in
posixtimer_send_sigqueue().
- Protecting the unlocked setting of PF_EXITING in exit_signals() for the
task group empty and the group exit case with sighand lock
- Flushing task::pending signals right there.
Optimize that by moving the whole pending list to an on-stack list head
under sighand lock and free the signals without the lock held.
There has been quite some discussion about the lockless flush and the
non-leader exec case on weakly ordered systems. The problem is that a third
party which tries to send a posix timer signal relies on the PID lookup to
find the target task and that lookup might result in the new leader when
the signal was originaly directed to the old leader. In case that the
signal was queued on the old leader then the lockless flush raised a
concern over the following situation:
old_leader new_leader third party
A: flush_list() // list_del_in
---truncated--- |
| In the Linux kernel, the following vulnerability has been resolved:
swiotlb: use the adjusted address for the highmem page lookup
swiotlb_bounce() reads the page frame number from the slot's recorded
orig_addr, then advances orig_addr by tlb_offset to reach the address
the caller asked about. The highmem branch mixes the two: the offset
within the page comes from the adjusted address, the page from the value
before it.
Once the adjustment crosses a page boundary the pair no longer describes
one location, and the whole copy lands one page below the intended one
for a positive tlb_offset, one above for a negative one. DMA_FROM_DEVICE
writes the device data over the wrong page and leaves the intended one
stale, DMA_TO_DEVICE feeds the device from a page the mapping may not
cover. Partial syncs through dma_sync_single_range_for_*() are what make
tlb_offset non-zero.
The branch test is picked the same way, so a slot recorded in lowmem can
be adjusted into highmem and the lowmem path then hands a highmem
address to phys_to_virt().
Take both from orig_addr once it is final and keep pfn in the branch
that uses it. PhysHighMem() asks the question straight from the address,
as dma-debug already does. |
| In the Linux kernel, the following vulnerability has been resolved:
RDMA/ucma: Serialize join and leave on copy_to_user failure
rdma_join_multicast() queues RoCE work that later reads the ucma_multicast
through event->param.ud.private_data, then list_add()s the CMA multicast
at the head of id_priv->mc_list. rdma_leave_multicast() matches only by
sockaddr and destroys the first hit.
ucma_process_join() used to drop ctx->mutex after a successful join and
retake it only if copy_to_user() failed. Two concurrent JOIN_MCAST calls
with the same address can therefore insert a second CMA entry before the
first thread's leave. leave then cancels the newer work and the older
worker still dereferences the ucma_multicast that the first thread frees.
Keep ctx->mutex held from rdma_join_multicast() through copy_to_user() and,
on -EFAULT, through rdma_leave_multicast() so leave cannot miss this join.
Do not leave if join itself failed: that path never published this address
on mc_list, and a leave-by-addr would destroy an earlier successful join. |
| In the Linux kernel, the following vulnerability has been resolved:
RDMA/core: fix refcount bug in iwpm_get_nlmsg_request()
iwpm_get_nlmsg_request() initializes refcount _after_ list_add_tail()
making it accessible to global list where another CPU can kref_get()
on nlmsg_request causing a refcount "addition on 0" bug. Fix this
by initializing kref _before_ list_add_tail() so refcount for
nlmsg_request can be incremented/decremented normally. In addition,
also initialize every field before list_add_tail(). |
| In the Linux kernel, the following vulnerability has been resolved:
openvswitch: avoid reallocating confirmed conntrack labels
ovs_ct_get_conn_labels() adds the labels extension when a conntrack
entry does not have one. Confirmed conntracks can be read locklessly,
so adding an extension may reallocate and free the extension block
while another CPU accesses it.
Only add the extension for unconfirmed conntracks. A confirmed
conntrack without labels now fails the caller's label operation instead
of reallocating its extension storage. |
| In the Linux kernel, the following vulnerability has been resolved:
dma-buf/dma-fence: fix checking signaling bit for timeline and driver name v3
The patch "dma-buf: dma-fence: Fix potential NULL pointer dereference"
changed the check to test for the ops pointer instead of the signaled
bit to avoid a potential NULL dereference when the ops pointer has been
cleared.
The problem is now that the ops pointer is cleared only when neither the
release nor the wait callback is implemented and this isn't true for a lot
of dma_fence implementations yet. So those implementations lost the RCU
protection after signaling of the returned string resulting in potential
use after free.
Add the signaling check additional to the ops pointer check so that we
have both the protection against NULL dereference as well as the RCU
protection after signaling for the returned string.
v2: improve comments to note RCU protection and explain why we check
both signaling state and ops pointer
v3: some comment improvements suggested by Philip |
| In the Linux kernel, the following vulnerability has been resolved:
ipv6: xfrm: use full sockets in local error paths
xfrm6_local_rxpmtu() and xfrm6_local_error() dereference skb->sk as if it
always pointed at a full IPv6 socket.
That is not guaranteed. TCP SYN-ACK skbs can be owned by a
TCP_NEW_SYN_RECV request_sock while the output path itself is driven by the
full listener. If rerouting selects an IPv6 XFRM tunnel route with a lower
MTU, the local PMTU/error handling path can reach these callbacks with that
mini-socket still attached to the skb.
The callbacks then miscast the request socket as a full inet/IPv6 socket and
can read beyond the request_sock allocation when they access inet_sock or
ipv6_pinfo state.
Resolve the owner with skb_to_full_sk() in both callbacks and bail out when
no full socket is attached. This matches the surrounding XFRM IPv6 PMTU/error
logic, which already reasons about full sockets with skb_to_full_sk(). |
| In the Linux kernel, the following vulnerability has been resolved:
net: lan743x: fix RX checksum use-after-free
lan743x_rx_process_buffer() adds each non-first receive buffer to the
head skb's frag_list. On the last descriptor, lan743x_rx_trim_skb()
linearizes the head and frees the fragment skb metadata.
The checksum-success path then writes ip_summed through the local skb
pointer, which still points to the final fragment. This causes a
use-after-free write when a packet spans more than one receive buffer.
Set ip_summed on the surviving head skb instead. Multi-buffer receive
can occur after a live MTU increase because existing ring entries keep
their old buffer size until they are replenished.
A KUnit test invoking lan743x_rx_process_buffer() with a two-buffer
packet produced a one-byte KASAN use-after-free write before this change.
The same test passed after the change. The driver object also builds
with W=1. This was not tested on physical LAN743x hardware. |
| In the Linux kernel, the following vulnerability has been resolved:
xfrm: use hlist_del_init_rcu for state_cache and state_cache_input
Commit 14acf9652e56 ("xfrm: defensively unhash xfrm_state lists in
__xfrm_state_delete") converted bydst/bysrc/byseq/byspi from
hlist_del_rcu() to hlist_del_init_rcu() so that a second
__xfrm_state_delete() on the same object becomes a no-op rather than a
write through LIST_POISON pprev. It missed state_cache and
state_cache_input, which kept hlist_del_rcu():
- hlist_del_rcu() leaves pprev = LIST_POISON2 (non-NULL), so
hlist_unhashed() returns false.
- hlist_del_init_rcu() leaves pprev = NULL, so hlist_unhashed()
returns true.
A second __xfrm_state_delete() therefore enters __hlist_del() on the
already-deleted state_cache/state_cache_input nodes and does
WRITE_ONCE(*pprev, next) through LIST_POISON2 — a write use-after-free
once the slab is reused. The corruption can in turn cause a subsequent
hlist_for_each_entry_rcu traversal to follow a dangling next pointer,
producing the read use-after-free reported in xfrm_input_state_lookup().
Switch state_cache and state_cache_input to hlist_del_init_rcu() to
match the other four lists, closing the write use-after-free and, with
it, the read use-after-free it spawns. |
| In the Linux kernel, the following vulnerability has been resolved:
xfrm: save input state data before secpath resets
xfrm_input() stores the current xfrm_state in the skb secpath while it
continues receive-side processing. Some input paths can reset that secpath
before xfrm_input() has finished dereferencing the state.
Receive callback users such as VTI and XFRM interfaces can reset the
secpath. The VTI receive path does so before checking whether the packet
crosses network namespaces, while the XFRM interface path does so only for
cross-network-namespace packets. The XFRM_MAX_DEPTH error path can also
reset the secpath before the final drop callback reports the current
state's protocol.
If secpath_reset() drops the last state reference while the state is
concurrently deleted, xfrm_input() can still dereference the freed state
when selecting transport_finish() or reporting the drop callback protocol.
Save the state protocol on the stack while the state is still valid,
and use the already saved address family for transport_finish(). A larval
XFRM_STATE_ACQ state has no type, so retain nexthdr as its protocol. This
preserves the existing drop-path fallback while avoiding the post-reset
state dereferences without adding an extra state reference to every
received packet. |
| In the Linux kernel, the following vulnerability has been resolved:
mips: select CONFIG_WEAK_REORDERING_BEYOND_LLSC from CONFIG_EYEQ
On I6500 CPU cores, lld and scd give no ordering guarantees (same as all
other instructions). To respect the assumption that arch_cmpxchg() is
fully ordered, we must inject sync instructions above and below our
lld/scd loops using the already in place WEAK_REORDERING_BEYOND_LLSC
infrastructure.
Otherwise, bad things can happen:
[ 34.054496] CPU 3 Unable to handle kernel paging request at virtual address 0000000000000000, epc == a80000080838e01c, ra == a80000080838dfc4
[ 34.054559] Oops[#1]:
[ 34.069561] CPU: 3 UID: 0 PID: 170 Comm: pipe_race Not tainted 7.2.0-rc6-01553-gb73c35220968-dirty #103 VOLUNTARY
[ 34.079932] Hardware name: Mobile EyeQ5 MP5 Evaluation board
[ 34.085592] $ 0 : 0000000000000000 0000000000000001 0000000000000000 0000000000000000
[ 34.093616] $ 4 : a800000808ee2618 000000000b7a879d 0000000000001000 0000000000000000
[ 34.101638] $ 8 : 0000000000e3f2c9 0000000000000000 a800000808a2a9f8 0000000000000000
[ 34.109660] $12 : a8000008139ffcd8 ffffffff84080018 a80000080837fae0 7878787878787878
[ 34.117682] $16 : a800000807e82940 0000000000001000 0000000000000000 0000000000000000
[ 34.125704] $20 : a800000802920e00 a8000008139ffdf8 a800000802649400 0000000000e3f2c9
[ 34.133726] $24 : 0000000000000006 00000001200406e0
[ 34.141783] $28 : a8000008139fc000 a8000008139ffd10 0000000000e3f2c8 a80000080838dfc4
[ 34.149837] epc : a80000080838e01c anon_pipe_read+0xd4/0x428
[ 34.155697] ra : a80000080838dfc4 anon_pipe_read+0x7c/0x428
[ 34.161549] Status: 140000e3 KX SX UX KERNEL EXL IE
[ 34.166551] Cause : 40800408 (ExcCode 02)
[ 34.170574] BadVA : 0000000000000000
[ 34.174161] PrId : 0001b028 (MIPS I6500)
[ 34.178183] Process pipe_race (pid: 170, threadinfo=000000005ca35720, task=00000000e1013890, tls=000000014ebbb780)
[ 34.188568] Stack : a800000802649400 0000000000000000 0000000000000000 a8000008139ffdd0
[ 34.196623] 0000000000000fba a800000808ee0000 0000000000000001 a8000008130c3e80
[ 34.204676] a8000008080d1280 a8000008139ffd58 a8000008139ffd58 1dbd2b22ea1dd500
[ 34.212729] a800000802649400 a800000808ee0000 ffffffffffffffea 0000000000000001
[ 34.220783] 0000000000001000 0000000000000000 00000001200ae518 ffffffffffffffff
[ 34.228836] 000000fffbe0e530 a80000080837edf4 000000fffbe0e530 0000000000000000
[ 34.236890] 0000000000000000 0000000000000000 000000014ebb55a0 0000000000001000
[ 34.244943] 0000000000000001 a800000802649400 0000000000000000 0000000000000000
[ 34.252996] 0000000000000000 0000400400000000 0000000000000000 1dbd2b22ea1dd500
[ 34.261049] 00000000140000e3 a800000802649400 a800000802649400 a800000808ee0000
[ 34.269103] ...
[ 34.271568] Call Trace:
[ 34.274026] [<a80000080838e01c>] anon_pipe_read+0xd4/0x428
[ 34.279533] [<a80000080837edf4>] vfs_read+0x25c/0x318
[ 34.284607] [<a80000080837faac>] ksys_read+0x104/0x138
[ 34.289763] [<a80000080802b9cc>] syscall_common+0x44/0x68
[ 34.295187]
[ 34.296689] Code: f84000cf 02209825 de020010 <dc420000> d8400004 02002825 0040f809 02802025 f84000c3
[ 34.306504]
[ 34.308099] ---[ end trace 0000000000000000 ]---
My initial reproducer was the xdp-tools test suite. A standalone
reproducer would be an lld/scd loop that, when the read is reordered by
the CPU, triggers a fault. We can achieve this from userspace by
stressing an anonymous pipe, which uses a mutex. Program used:
// SPDX-License-Identifier: GPL-2.0
// pipe_race.c - reproducer for MIPS LL/SC reordering vs fs/pipe.c
//
// Two userspace processes on an anonymous pipe:
// parent = writer: tight write() loop
// child = reader: tight read() loop
#define _GNU_SOURCE
#include <assert.h>
#include <errno.h>
#include <sched.h>
#include <signal.h>
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
#include <sys/types.h>
#include
---truncated--- |
| In the Linux kernel, the following vulnerability has been resolved:
IB/hfi1: Fix the PIO_CRED credit-return mmap
hfi1_file_mmap()'s PIO_CRED case must hand user space the single
credit-return page that holds this context's entry. That page is the
second or third page of the per-node credit-return allocation once the
hardware send context index reaches 64 or 128, so the failure below is
intermittent: when the entry lands on the first page the offset is zero
and everything works.
Two things are wrong.
First, cr_page_offset is a byte offset but .va is a struct
credit_return *, so adding it is pointer arithmetic and scales the offset
by sizeof(struct credit_return) == 64. memvirt then lands 256 KiB or
512 KiB past a 10240-byte allocation. With an IOMMU translating, that
address is inside the vmalloc range but in no vm_area, so
dma_mmap_coherent() -> iommu_dma_mmap() finds no pages, vmalloc_to_pfn()
returns page_to_pfn(NULL), and remap_pfn_range() installs a frame above
MAXPHYADDR. The first user read then takes:
psm2_ep_open_pr: Corrupted page table at address 7a14d007e000
PGD 800000013886a067 P4D 800000013886a067 PUD 13886b067 PMD 13886c067
PTE 800049168e911235
Oops: Bad pagetable: 000d [#1] SMP PTI
Second, and still wrong once the arithmetic is corrected,
dma_mmap_coherent() describes a whole coherent buffer and selects the
page within it with vma->vm_pgoff. Offsetting cpu_addr has no effect:
for a vmap'd allocation iommu_dma_mmap() uses cpu_addr only to locate the
vm_area and then maps pages[vm_pgoff], which hfi1_file_mmap() has just
set to 0. User space therefore always receives the first credit-return
page, every credit read is for the wrong context, and send PIO stalls
forever.
Use the DMA API as intended: pass the base of the allocation with its
full length and select the page with vm_pgoff. A separate length is
needed because memlen must keep describing the VMA for the existing size
check. The dma-direct path stays correct as well, since dma_direct_mmap()
adds the same vm_pgoff to the base pfn.
Tested on a Dell T7610 (Xeon E5-2650 v2, Intel IOMMU in DMA-FQ mode)
against a Threadripper PRO 3995WX peer, both Omni-Path 100. Before this
change psm2_ep_open() Oopses the kernel; with only the arithmetic
corrected psm2_ep_open() succeeds but any transfer that uses send PIO
hangs, PSM2_SDMA=2 (send PIO disabled) completing normally while
PSM2_SDMA=0 (send PIO only) hangs every time. With this change send PIO,
send DMA and the default mixed mode all work. |
| In the Linux kernel, the following vulnerability has been resolved:
hwmon: (w83791d) remove fan/pwm 4-5 sysfs group on remove
When the fan/pwm 4-5 pins are not used as GPIO, w83791d_probe()
creates the w83791d_group_fanpwm45 sysfs group on the I2C client
device.
The probe error path removes this group when a later initialization
step fails, but the normal remove path only removes w83791d_group.
As a result, the optional fan/pwm 4-5 sysfs files can remain after the
driver is unbound.
The callbacks associated with these files access the driver data,
which is devm allocated and released after driver unbind. Leaving the
sysfs files behind can therefore result in accesses to stale driver
data.
Remove w83791d_group_fanpwm45 during normal teardown as well.
This issue was found by manual code inspection. |
| In the Linux kernel, the following vulnerability has been resolved:
drm/msm: RCU-free the scheduler-containing ring and VM objects
Both struct msm_ringbuffer and struct msm_gem_vm embed a struct
drm_gpu_scheduler. msm_ringbuffer_destroy() and the VM free callback
msm_gem_vm_free() call drm_sched_fini() on the embedded scheduler and then
free the containing object with plain kfree().
drm_sched_fence_get_timeline_name() returns fence->sched->name, and the
scheduler fence keeps a .release callback so it is not ops-detached on
signalling. A finished fence exported to userspace (the submit out-fence, or
a VM_BIND fence, via sync_file / drm_syncobj) keeps pointing at the embedded
scheduler after the ring/VM is freed, so a later get_timeline_name() --
reachable unprivileged through SYNC_IOC_FILE_INFO -- dereferences freed slab
memory (KASAN slab-use-after-free read).
Per the dma-fence lifetime contract the exporter must keep the data backing a
signalled fence alive for an RCU grace period. Free the scheduler-containing
objects with kfree_rcu() instead of kfree().
Patchwork: https://patchwork.freedesktop.org/patch/750234/ |
| In the Linux kernel, the following vulnerability has been resolved:
smb: client: cancel reconnect work in clean_demultiplex_info()
clean_demultiplex_info() cancels server->echo delayed work but not
server->reconnect, which can cause a use-after-free when the
demultiplex thread exits while a reconnect work is still queued:
cifs_demultiplex_thread()
cifs_readv_from_socket()
cifs_reconnect()
__cifs_reconnect()
cifs_queue_server_reconn()
mod_delayed_work(cifsiod_wq, &server->reconnect, 0)
clean_demultiplex_info()
cancel_delayed_work_sync(&server->echo) // echo canceled
// reconnect NOT canceled
kfree_sensitive(server) // server freed
...later, on cifsiod_wq:
smb2_reconnect_server()
server->srv_count // UAF read of freed server
Fix this by canceling server->reconnect delayed work in
clean_demultiplex_info() before the server is freed, the same way
cifs_put_tcp_session() already does. |
| In the Linux kernel, the following vulnerability has been resolved:
smb: client: fix rlist race and missing initialization
TCP_Server_Info.rlist is allocated via kzalloc which zeros both ->next
and ->prev to NULL instead of pointing to itself, making list_empty()
always return false and list_add() dereference a NULL ->prev pointer.
Also, cifs_signal_cifsd_for_reconnect() can be called concurrently
from multiple cifsd threads, allowing the same server's rlist node to
be added twice into the local list, corrupting it. |
| In the Linux kernel, the following vulnerability has been resolved:
smb: client: fix use-after-free of iface in cifs_try_adding_channels()
cifs_try_adding_channels() iterates ses->iface_list with
list_for_each_entry_safe_from(), which captures the next entry
(niface) under iface_lock. The loop body then drops iface_lock for
the whole duration of cifs_ses_add_channel().
A concurrent interface refresh (SMB3_request_interfaces() ->
parse_server_interfaces()) marks all ifaces inactive and removes and
frees any that are not re-advertised via list_del() + kref_put(),
where release_iface() is a bare kfree(). Since niface typically has
no channel holding a reference, the list reference is its last and it
can be freed inside the unlocked window. On continue, the iterator
advance step then dereferences niface->iface_head.next, and the loop
body reads iface->rdma_capable/is_active, both on freed memory.
Fix this by never keeping an unreferenced list pointer across the
unlocked window. Each channel attempt now re-scans the list from the
head under iface_lock, takes a kref on the selected candidate, and
passes only that referenced candidate to cifs_ses_add_channel().
weight_fulfilled still tracks selection progress, so restarting the
scan preserves the original weighted distribution and the
weight_fulfilled-before-kref_put ordering on the failure path.
Add a per-pass attempts cap so a flapping interface refresh cannot
keep the inner loop spinning within a single tries increment. |
| In the Linux kernel, the following vulnerability has been resolved:
smb: client: fix next_buffer UAF and NextCommand bounds in compound PDUs
Fix several related bounds checking and pointer lifecycle issues in
receive_encrypted_standard()'s handling of compound encrypted frames:
- Clear next_buffer after assigning it to server->bigbuf. A stale
next_buffer pointer can lead to a use-after-free on subsequent
error paths.
- Update pdu_length to the decrypted plaintext size (buf_size). Using
the pre-decryption length allows NextCommand to point into stale
ciphertext residue.
- Reject next_cmd values smaller than MID_HEADER_SIZE(server).
- Fix an integer overflow in the upper bound check by verifying
pdu_length - next_cmd < MID_HEADER_SIZE(server), ensuring the
trailing slice is large enough for a header. |
| In the Linux kernel, the following vulnerability has been resolved:
smb: client: fix potential OOB read in smb3_enum_snapshots()
If snapshot_array_size is smaller than GMT_TOKEN_SIZE,
smb3_enum_snapshots() sets ret_data_len to
sizeof(struct smb_snapshot_array) without verifying the actual length
of the server's reply.
Because SMB2_ioctl() places no lower bound on the server-supplied
OutputCount and allocates retbuf to exactly that length, a short reply
results in ret_data_len exceeding the size of retbuf. The subsequent
copy_to_user() then reads past the end of retbuf, leaking adjacent slab
memory to userspace. The subsequent clamp check is ineffective as it
only reduces ret_data_len.
Fix this by rejecting replies shorter than
sizeof(struct smb_snapshot_array) with -EIO. Note that the bound is set
to the 12-byte struct size rather than the 16-byte
MIN_SNAPSHOT_ARRAY_SIZE defined in MS-SMB2 3.3.5.15.1, because 12 bytes
is exactly what copy_to_user() attempts to read. |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: x86/mmu: Check write tracking in all address spaces
kvm_gfn_is_write_tracked() checks only the supplied memslot, but page
tracking is per-address-space and shadow pages are shared across all
address spaces. With SMM, a GFN can therefore be write-tracked in one
address space and appear untracked through the other.
Check the supplied slot first, then the slot for the other address space.
This ensures all callers honor write tracking regardless of the active
address space. In particular, it prevents mmu_try_to_unsync_pages() from
marking an upper-level shadow page unsync and eventually triggering the
BUG in pte_list_remove().
[invert direction of the conditional. - Paolo] |
