| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| In the Linux kernel, the following vulnerability has been resolved:
mm/damon/ops-common: putback folios on invalid migrate nid
damon_pa_migrate() and damos_va_migrate() isolate folios into a local list
and then call damon_migrate_pages(). When target_nid is invalid
(including the scheme default NUMA_NO_NODE / -1), damon_migrate_pages()
returns early without putting the folios back to the LRU.
Callers then discard the list head while those folios remain isolated with
an extra reference taken by folio_isolate_lru(). The pages stay off the
LRU for as long as the mapping exists (anon active+inactive counts drop
while RSS does not), and the leftover references can pin the pages after
the mapping is gone.
Put the folios back on the invalid-nid path so ignored migration requests
still return them to the LRU. |
| In the Linux kernel, the following vulnerability has been resolved:
ALSA: usx2y: bound the hwdep mmap fault offset
snd_us428ctls_vm_fault() turns the faulting page offset into a kernel
address with no bound of any kind:
offset = vmf->pgoff << PAGE_SHIFT;
vaddr = (char *)(...)->us428ctls_sharedmem + offset;
page = virt_to_page(vaddr);
get_page(page);
vmf->page = page;
return 0;
snd_us428ctls_mmap() checks only the length of the mapping, never the
offset, and us428ctls_sharedmem is a single page from
alloc_pages_exact(). For a character device file_mmap_size_max()
returns ULONG_MAX, so the mm layer imposes no ceiling either. Every page
offset above zero resolves to a struct page outside the object, and the
handler installs it into the caller's address space read-write; the vma
is not marked read-only.
The caller picks the page frame with a single mmap() argument and gets
read-write access to a page of kernel memory it does not own; an offset
that lands in an unpopulated vmemmap region oopses instead.
A process that can open the hwdep node of an attached US-X2Y reaches
this after loading the FPGA image through the same node; no capability
check is involved.
On 7.2.0-rc5 (arm64), mmap() with a large offset:
Unable to handle kernel paging request at virtual address fffffdffc45d5ac8
pc : snd_us428ctls_vm_fault+0x68/0x140 [snd_usb_usx2y]
Call trace:
snd_us428ctls_vm_fault+0x68/0x140 [snd_usb_usx2y]
__do_fault
__handle_mm_fault
handle_mm_fault
el0_da
Reject any offset outside the shared region. The pcm hwdep handler in
usx2yhwdeppcm.c computes its address the same way and needs the same
bound.
Discovered by XBOW, triaged by Baul Lee <[email protected]> |
| In the Linux kernel, the following vulnerability has been resolved:
perf/core: Fix group leader use-after-free after sibling detach
perf_group_detach() handles leader and sibling detach differently. When the
group leader is detached, all siblings are promoted to singleton events and
their group_leader pointer is reset to themselves. When a sibling is
detached, it is removed from the leader's sibling_list, but its
group_leader pointer is left pointing at the old leader.
That is harmless when the sibling is being closed and freed immediately, as
in the DETACH_DEAD path. It is not safe when the sibling is detached but
kept alive, such as during CPU hotplug with DETACH_GROUP. In that case the
sibling is removed from the context, while its file descriptor can still
keep it alive.
A typical failing sequence is:
- A group contains leader L and sibling S.
- CPU hot-unplug detaches S with DETACH_GROUP, removing it from
L->sibling_list but leaving S->group_leader == L.
- L is later closed and freed.
- A PERF_IOC_FLAG_GROUP ioctl on S follows S->group_leader and
dereferences the freed leader.
This was reproduced by running the perf event fuzzer, CPU hotplug, and a
stress workload concurrently:
Unable to handle kernel paging request at virtual address 006b6b6b6b6b6cdb
CPU: 2 PID: 12489 Comm: perf_fuzzer 6.18.7 PREEMPT
pc : perf_ioctl+0x34c/0xc68
x20: ffffff89a3fa2c70 x8 : 6b6b6b6b6b6b6b6b
Code: 943c4a0e 340047a0 f9404a94 f9411e88 (f940b908)
Call trace:
perf_ioctl+0x34c/0xc68 (P)
__arm64_sys_ioctl+0xa0/0xf4
invoke_syscall+0x58/0xe4
el0_svc_common+0xa8/0xdc
do_el0_svc+0x1c/0x28
el0_svc+0x40/0xc0
el0t_64_sync_handler+0x68/0xdc
el0t_64_sync+0x1c4/0x1c8
The fault happened in perf_ioctl(), where perf_event_for_each() follows
the stale group_leader pointer and perf_event_for_each_child() then
dereferences the freed leader's context.
Fix the use-after-free by promoting the detached sibling to a singleton.
Also fix __event_disable() cgroup accounting and event state change. |
| In the Linux kernel, the following vulnerability has been resolved:
mm/huge_memory: fix huge_zero_pfn race
Patch series "mm/huge_memory: fix huge_zero_pfn race", v2.
There is a subtle race in the reference-counted huge_zero_folio
implementation.
The fast path atomic logic fails to account for the fact that the shrinker
(which drops the final huge_zero_refcount pin) can overwrite huge_zero_pfn
with the ~0UL sentinel value in shrink_huge_zero_folio_scan() after a
racing get_huge_zero_folio() installed a valid value there.
This results in huge_zero_folio being correctly set but huge_zero_pfn
being set incorrectly and thus is_huge_zero_pfn() and consequently
is_huge_zero_pmd() will misidentify the huge zero folio as being an
ordinary THP folio.
This can result in the huge zero folio being split and otherwise treated
incorrectly.
The solution to this is very subtle as there is an atomic fast path, and
thus ordering in weakly ordered architectures has to be treated very
carefully.
The first commit fixes the issue by introducing a spinlock around
huge_zero_[pfn, folio, refcount] write, with careful consideration paid to
load/store ordering in the fast path. It is placed first and kept as
small as possible so that it can be backported on its own.
The second commit is a pure cleanup which reworks the
CONFIG_PERSISTENT_HUGE_ZERO_FOLIO logic to better separate the persistent
logic from the dynamically allocated one.
This patch (of 2):
If !CONFIG_PERSISTENT_HUGE_ZERO_FOLIO, the huge_zero_folio is refcounted
by huge_zero_refcount and returned by mm_get_huge_zero_folio().
When the caller is done with the huge zero page, its reference count is
decremented. Only a shrinker can set the reference count to zero.
A race can unfortunately occur between a shrinker decrementing the
reference count to zero and a concurrent page fault.
This is because shrink_huge_zero_folio_scan() might, if very unlucky, be
preempted between setting huge_zero_refcount to zero and writing an
invalid value.
During this time get_huge_zero_folio() could write to huge_zero_pfn before
shrink_huge_zero_folio_scan() resumes.
In this event the huge zero folio will be persistently misidentified
causing the THP code path to be entered inappropriately for the huge zero
folio:
CPU 0 CPU 1
=======================================|=================================
shrink_huge_zero_folio_scan() |
atomic_cmpxchg() sets refcount to 0 |
xchg() sets huge_zero_folio to NULL | get_huge_zero_folio()
| | atomic_inc_not_zero() -> zero
preempted for a long time | Allocate new huge zero folio
| | Write valid huge_zero_folio
v | Write valid huge_zero_pfn
Overwrite huge_zero_pfn with ~0UL <--- Invalid overwrite!
This results in is_huge_zero_pfn() and is_huge_zero_pmd() incorrectly
returning false for a huge zero page which could result in issues like the
huge zero folio being incorrectly split.
Note that the issue is with huge_zero_pfn not huge_zero_folio, as
get_huge_zero_folio() uses cmpxchg() gated on huge_zero_folio being NULL
with a retry loop and shrink_huge_zero_folio_scan() uses xchg() to set
huge_zero_folio.
Fix the issue by introducing a spinlock, huge_zero_lock, to prevent
concurrent write of huge_zero_folio, huge_zero_pfn and huge_zero_refcount.
There needs to be significant care taken here to ensure correctness:
The fast path in get_huge_zero_folio() uses atomic_inc_not_zero(), which
is outside of the critical section, and means huge zero allocation is
gated on zero huge_zero_refcount.
The fast path doesn't use huge_zero_lock, so the critical section is
irrelevant to it.
So invariants are required - huge_zero_refcount MUST:
* Only be set in the huge_zero_lock critical section to ensure
serialisation of huge_zero_pfn, huge_zero_folio and
---truncated--- |
| In the Linux kernel, the following vulnerability has been resolved:
net: smc: fix splice entry lifetime imbalance in smc_rx_splice
smc_rx_splice() passes pages to splice_to_pipe() before taking the
references that cover the lifetime of each splice entry. In the
VM-backed RMB path, splice_to_pipe() may drop unqueued entries through
smc_rx_spd_release(), while queued entries are released later via the
pipe buffer callback.
The old post-splice accounting also derives the number of queued VM pages
from an offset mutated while building the descriptor, and a multi-page
splice pairs one sock_hold() with multiple sock_put() calls.
Take the page and socket references for every candidate entry before
splice_to_pipe(), and drop the matching private state, page reference,
and socket reference from smc_rx_spd_release() for entries that never
get queued. This fixes a refcount imbalance that can underflow page
refcounts and trigger a use-after-free. |
| In the Linux kernel, the following vulnerability has been resolved:
net/x25: fix use-after-free of the socket by its timers
The x25 timers are armed with mod_timer() and cancelled with
timer_delete(), so a pending timer holds no reference on the socket and a
cancel does not wait for a callback already running on another CPU.
x25_heartbeat_expiry() also rearms unconditionally, so it can reinstall
sk->sk_timer after __x25_destroy_socket() has passed its cancel point.
The following __sock_put() frees the socket while the timer is still
queued, and the next expiry uses freed memory. KASAN reports a
slab-use-after-free on the kmalloc-2k object freed by close().
timer_delete_sync() cannot be used here: x25_heartbeat_expiry() and
x25_timer_expiry() both reach the cancels from inside the timer they
would wait on, through __x25_destroy_socket() and x25_disconnect().
Arm the timers with sk_reset_timer() and cancel them with sk_stop_timer()
so that an armed timer owns a reference, and release it in both expiry
handlers. Rearm the heartbeat only while sk_hashed(sk) is still true,
since __x25_destroy_socket() unlinks the socket before dropping it. Arm
the deferred destroy timer the same way and drop its reference in
x25_destroy_timer().
Reproduced on net with KASAN, with the heartbeat period shortened so the
window recurs. With this patch the reproducer no longer triggers a
report and /proc/net/x25 drains.
Discovered by XBOW, triaged by Baul Lee <[email protected]> |
| In the Linux kernel, the following vulnerability has been resolved:
NTB: ntb_netdev: Preserve RX queue depth on allocation failure
ntb_netdev_rx_handler() hands the received skb to the network stack
before allocating its replacement. If the allocation fails, nothing is
reposted. Every failure therefore takes one buffer out of the RX queue
while the interface remains up, and enough failures eventually stall
reception.
A retry path could refill the queue later, but ntb_netdev has none.
Allocate the replacement first instead. If that fails, drop the packet
and repost the same skb. This keeps the queue full and lets packet
delivery resume as soon as memory is available again. |
| In the Linux kernel, the following vulnerability has been resolved:
netfilter: nf_conntrack: defer invalid log until after unlock
TCP and SCTP conntrack paths can emit invalid-packet logs while ct->lock
is still held.
When invalid logging is routed to nfnetlink_log and conntrack export is
enabled, the log path can re-enter conntrack netlink glue and dump the
same conntrack again. Protocol attribute dumping may take ct->lock, so
logging while holding that lock can deadlock.
Defer the TCP invalid logs by storing only the minimal log context while
ct->lock is held and emitting the log after unlocking. Also make the TCP
timeout-lowering invalid path return whether a log is needed, then emit
that log after unlocking.
Do the same for the SCTP invalid state-transition log that can be reached
while ct->lock is held.
Add a lockdep assertion to nf_ct_l4proto_log_invalid() so future callers
that log invalid conntracks while holding ct->lock are caught outside TCP
and SCTP as well. |
| In the Linux kernel, the following vulnerability has been resolved:
net: atlantic: free stranded TX buffers on ring deinit
aq_vec_deinit() drains the TX rings with a single aq_ring_tx_clean()
call, which frees at most AQ_CFG_TX_CLEAN_BUDGET (256) descriptors and
stops at hw_head, which no longer moves once aq_vec_stop() has stopped
the hardware and NAPI. Completed descriptors beyond the budget and
everything still posted in [hw_head, sw_tail) keep their skb or
xdp_frame when the interface goes down: aq_vec_ring_free() then frees
the buffer ring and the references are lost for good.
Today this is a silent memory leak on every interface down under
TX/XDP_TX load. With the conversion of the RX path to page_pool posted
for net-next it becomes much more visible: XDP_TX frames carry fragment
references on the RX ring's page_pool, so a single stranded frame keeps
the pool's inflight count above zero forever. page_pool_destroy() then
never completes, the pool is leaked together with its pages, and
"page_pool_release_retry() stalled pool shutdown" is warned every 60
seconds from that point on, on every ifdown, XDP detach or ring resize
under XDP_TX load.
Bring back aq_ring_tx_deinit() as it was before the removal and use it
for teardown again, with one extension: TX rings can hold xdp_frames
nowadays, so release those too. They are returned with
xdp_return_frame() since this runs in process context. |
| In the Linux kernel, the following vulnerability has been resolved:
net/sched: act_ct: fix sk_buff leak when the header checks reject a packet
tcf_ct_handle_fragments() runs its header sanity checks before handing
anything to the defragmentation engine:
if (family == NFPROTO_IPV4)
err = tcf_ct_ipv4_is_fragment(skb, &frag);
else
err = tcf_ct_ipv6_is_fragment(skb, &frag);
if (err || !frag)
return err;
tcf_ct_ipv4_is_fragment() returns -EINVAL or -ENOMEM;
tcf_ct_ipv6_is_fragment() adds -EPROTO when ipv6_find_hdr() fails. None of
them frees or queues the skb, so on that path the caller still owns it.
tcf_ct_act() however funnels every non-zero return into the
ownership-transfer exit:
err = tcf_ct_handle_fragments(net, skb, family, p->zone, &defrag);
if (err)
goto out_frag;
...
out_frag:
if (err != -EINPROGRESS)
tcf_action_inc_drop_qstats(&c->common);
return TC_ACT_CONSUMED;
TC_ACT_CONSUMED means the action took ownership of the skb, so no caller
frees it - sch_handle_ingress(), sch_handle_egress() and
tcf_qevent_handle() all deliberately skip the free for that verdict. The
skb is therefore orphaned: one sk_buff plus its data buffer is leaked per
malformed packet, unbounded. Note the drop counter is already incremented
for these errors, so the statistics claim a drop that never happens.
Three different ownership states reach out_frag: today - the skb may be
queued by the defrag engine (-EINPROGRESS), already freed by
nf_ct_handle_fragments(), or still owned by us. Tell the caller which of
those it is, and free the packet ourselves in the last case, which
restores the TC_ACT_SHOT behaviour that predated the Fixes: commit.
Reproduced on v7.2-rc6 with a 54-byte frame carrying a 40-byte IPv6
header with nexthdr = 0 (hop-by-hop) and nothing after it, on a
clsact ingress chain with "action ct". kmemleak reports one leaked
232-byte skbuff_head_cache object plus its 704-byte data buffer per
packet; with this patch it reports none. |
| In the Linux kernel, the following vulnerability has been resolved:
net/sched: act_gact, act_police: range check the fallback control action
tcf_action_check_ctrlact() range checks the primary control action:
if (!opcode)
ret = action > TC_ACT_VALUE_MAX ? -EINVAL : 0;
TC_ACT_VALUE_MAX is TC_ACT_TRAP, so kernel-internal verdicts above it
cannot be set that way. But act_gact and act_police each carry a second,
independent control action supplied by user space that never reaches that
helper - TCA_GACT_PROB.paction and TCA_POLICE_RESULT. Both only reject
TC_ACT_GOTO_CHAIN, so any other value is stored verbatim and returned
verbatim from the action.
In particular user space can store TC_ACT_CONSUMED, which is
TC_ACT_VALUE_MAX + 1 and is deliberately not part of the UAPI value
range. That verdict tells every caller the action took ownership of the
skb, so nobody frees it: sch_handle_ingress(), sch_handle_egress() and
tcf_qevent_handle() all deliberately skip the free for it. The result is
one leaked sk_buff plus its data buffer per packet traversing the filter,
unbounded, for all traffic on the chain including kernel-generated
packets.
Both are trivially deterministic. act_gact clamps tcfg_pval to >= 1, so
with pval = 1 gact_determ() returns the fallback for every packet.
act_police has no mandatory rate, so rate = 0 leaves tcfp_mtu = ~0 and
tcf_police_mtu_check() always passes.
TC_ACT_CONSUMED was added by commit 720f22fed81b ("net: sched: refactor
reinsert action"), after both goto-chain guards were written:
commit 9469f375ab09 ("net/sched: act_gact: disallow 'goto chain' on
fallback control action") and
commit c08f5ed5d625 ("net/sched: act_police: disallow 'goto chain' on
fallback control action"). Neither guard was widened when the new
verdict appeared.
Factor the existing range test out of tcf_action_check_ctrlact() as
tcf_action_valid() and apply it to both fallbacks. The helper cannot call
tcf_action_check_ctrlact() directly because that also allocates a
goto_chain, which is exactly what these two sites must not do.
Reproduced on v7.2-rc6: kmemleak reports one leaked 232-byte
skbuff_head_cache object plus its 704-byte data buffer per packet. With
this patch both configurations are rejected with -EINVAL and kmemleak
reports none. |
| In the Linux kernel, the following vulnerability has been resolved:
veth: fix skb length accounting after XDP frag adjustment
veth exposes non-linear skb fragments through an xdp_buff. If an XDP
program adjusts the fragment area, veth_xdp_rcv_skb() copies
xdp_frags_size back to skb->data_len but leaves skb->len containing the
old fragment contribution.
After a fragment shrink, this makes skb_headlen() larger than the actual
linear area. In the reproduced UDP receive path, __skb_datagram_iter()
copied 1024 bytes past the actual linear tail to userspace, starting at
struct skb_shared_info. The copied bytes included the affected skb's
nr_frags, xdp_frags_size, and a kernel pointer from
skb_shinfo(skb)->frags[0]. Real packet data was displaced by the same
amount and truncated at the end.
Subtract the old data_len before replacing it and add the new data_len
afterwards, keeping skb->len and skb->data_len synchronized.
Additionally, bpf_xdp_pull_data() can advance data_end while leaving
frags present. The skb is then still non-linear, so the old
__skb_put(skb, off) triggers SKB_LINEAR_ASSERT().
Use skb_set_tail_pointer() and update skb->len explicitly instead,
following bpf_prog_run_generic_xdp(). Unlike __skb_put(),
skb_set_tail_pointer() does not require a linear skb.
A 60000-byte UDP datagram on a veth pair with MTU 64000 was shortened by
1024 bytes from its fragment area. Before the fix, all 10 runs produced
corrupted payloads. After the fix, all 10 runs matched the expected
payload exactly. A forced-tailroom reproducer also exercises
bpf_xdp_pull_data() with frags still present; the old code triggers
SKB_LINEAR_ASSERT(), while this fix passes 10/10 runs. |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: SVM: Serialize accesses to the owner and mirror list with separate lock
Interaction between KVM_CAP_VM_MOVE_ENC_CONTEXT_FROM and
KVM_CAP_VM_COPY_ENC_CONTEXT_FROM can cause two separate issues:
- in sev_migrate_from(), when the destination KVM is a mirror, the mirror
entry is moved from the source's list to the owner's mirror_vms list,
without holding the owner's lock unlike other writers of the owner's
mirror list (sev_vm_copy_enc_context_from(), sev_vm_destroy()).
A concurrent COPY or destroy can race with sev_migrate_from() and
corrupt the list.
- In sev_vm_destroy(), the *owner* is still active and could receive
concurrently a KVM_CAP_VM_MOVE_ENC_CONTEXT_FROM that causes
sev->enc_context_owner to change. In this case the incorrect VM
receives kvm_put_kvm().
The second issue needs particular care because the owner could disappear
altogether (even though the race window is impossibly small) between
reading it and locking it. There is thus no way to perform the checks
under the owner lock without putting struct kvm under SLAB_TYPESAFE_BY_RCU
(which would allow kvm_get_kvm_safe() under RCU critical section).
It is much simpler to just use a global lock, since the critical
sections are so small and the new lock is always a leaf lock. |
| In the Linux kernel, the following vulnerability has been resolved:
ring-buffer: Initialise reader page order in rb_allocate_cpu_buffer()
In rb_allocate_cpu_buffer(), bpage->order was omitted, leaving it as 0.
This is an issue for a ring-buffer with subbufs bigger than PAGE_SIZE if
when freed: free_buffer_page() relies on this value. Align the value
with the actual allocation size (buffer::subbuf_order). |
| In the Linux kernel, the following vulnerability has been resolved:
ring-buffer: Use current_context for safe per-CPU buffer swap
The ring_buffer_swap_cpu() function currently checks the per-CPU
committing counter to determine if a buffer is actively being written to
before performing the swap. However, there exists a race window where
this check can be bypassed:
ring_buffer_lock_reserve
cpu_buffer = buffer->buffers[cpu]; // cpu_buffer_a
rb_reserve_next_event
rb_start_commit // inc committing
if (unlikely(READ_ONCE(cpu_buffer->buffer) != buffer)) {...}
__rb_reserve_next
rb_move_tail
rb_end_commit(cpu_buffer); // dec committing => 0
/* interrupt hits here, successfully swaps! */
local_inc(&cpu_buffer->committing);
ring_buffer_unlock_commit
cpu_buffer = buffer->buffers[cpu]; // cpu_buffer_b
rb_commit
rb_end_commit
RB_WARN_ON(cpu_buffer, !local_read(&cpu_buffer->committing))
// triggers warning
The committing counter can temporarily drop to 0 during a single write
operation (within rb_move_tail), creating a window where swap can
succeed even though the write is still in progress. This leads to
inconsistent buffer state and triggers the RB_WARN_ON in rb_commit().
Replace the committing counter check with current_context checks, which
are set at the entry of ring_buffer_lock_reserve() and remain valid
throughout the entire write operation, providing a reliable indicator of
buffer busy state during swap. |
| In the Linux kernel, the following vulnerability has been resolved:
mm/page_table_check: skip special zero mappings
page_table_check_set() and page_table_check_clear() account mappings based
on PageAnon(). Shared zero-page PTEs and huge zero PMDs are special
mappings, but page_table_check can still account them as file-backed
pages.
An unprivileged process can populate enough zero mappings to overflow
file_map_count and hit the existing BUG_ON(). The PTE path can do this
with the shared zero page, and the PMD path can do the same with huge zero
mappings.
Skip special zero mappings in the user page-table accounting paths. Keep
the PTE-side pte_special() check, and identify huge zero PMDs from the
mapped folio instead of pmd_special(). That covers architectures where
pmd_special() is a no-op without adding huge_zero_pfn checks to the
generic counter helpers. |
| In the Linux kernel, the following vulnerability has been resolved:
mm/ptdump: always stabilise against page table freeing using init_mm
Previous commits have established the invariant that kernel page table
freeing is performed while an mmap read lock on init_mm is held, which
fixes races between ptdump and kernel page table freeing over init_mm.
However, x86 and arm64 can perform a ptdump over an mm other than init_mm
via ptdump_walk_pgd() and since kernel memory ranges are shared across
non-kernel mm's, this means that the race still exists for these cases.
Fix this by acquiring a nested mmap write lock for init_mm in
ptdump_walk_pgd().
This is safe as we take this after mmap write locking the mm, and nothing
acquires the init_mm lock first before locking an arbitrary mm, so no
deadlock is possible.
Also update walk_page_range_debug() to assert that init_mm is write
locked, add a comment explaining why and remove some redundant code, and
eliminate the unnecessary and confusing invocation of
walk_kernel_page_table_range().
We can safely remove the non-NULL check for walk.mm, as the mmap lock
asserts would NULL pointer deref if it was (and of course no callers do
this).
The first point at which ptdump can race kernel page table freeing is
commit b6bdb7517c3d ("mm/vmalloc: add interfaces to free unmapped page
table"), so we target this in the Fixes tag. |
| In the Linux kernel, the following vulnerability has been resolved:
fscrypt: use the mount idmap for the owner check in fscrypt_ioctl_set_policy()
fscrypt_ioctl_set_policy() calls inode_owner_or_capable() with
&nop_mnt_idmap before allowing an encryption policy to be set, instead
of the idmap of the mount the ioctl was issued on.
fscrypt is used by filesystems that support idmapped mounts (e.g. ext4,
f2fs), so on such a mount this compares the caller's fsuid against the
unmapped on-disk owner rather than the mapped owner: the actual owner
can be wrongly denied with -EACCES and an unrelated caller wrongly
allowed. Use file_mnt_idmap(filp) instead. |
| In the Linux kernel, the following vulnerability has been resolved:
sched/psi: Shut down rtpoll_timer in psi_cgroup_free()
psi_schedule_rtpoll_work() is called locklessly from the scheduler hotpath
and can race psi_trigger_destroy() taking down the last rtpoll trigger under
rtpoll_trigger_lock:
psi_schedule_rtpoll_work() psi_trigger_destroy()
rcu_read_lock();
task = rcu_dereference(rtpoll_task);
rcu_assign_pointer(rtpoll_task, NULL);
timer_delete(&rtpoll_timer);
mod_timer(&rtpoll_timer, ...);
rcu_read_unlock();
synchronize_rcu();
kthread_stop(task_to_destroy);
The group can then be freed with the re-armed timer still pending, and
poll_timer_fn() runs on freed memory.
461daba06bdc ("psi: eliminate kthread_worker from psi trigger scheduling
mechanism") deleted the timer synchronously after the synchronize_rcu(),
which prevented this but raced trigger creation instead: the deletion could
cancel the timer that a new trigger set armed during the grace period and,
as creation also reinitialized the timer at the time, corrupt it.
8f91efd870ea ("psi: Fix race between psi_trigger_create/destroy") moved the
initialization into group_init() and the deletion into the locked section,
trading the creation races for the window above.
Neither placement in the destruction path works. A pending timer firing
while the group is alive is harmless though. poll_timer_fn() just wakes the
rtpoll waitqueue and doesn't re-arm itself. Bind the timer to the group's
lifetime instead and shut it down in psi_cgroup_free(). Nothing can arm it
by then. timer_shutdown_sync() because the timer is never armed again. |
| In the Linux kernel, the following vulnerability has been resolved:
sched_ext: Take cgroup_lock() first in scx_cgroup_lock()
scx_cgroup_lock() write-locks scx_cgroup_ops_rwsem and then takes
cgroup_lock(), which can deadlock through kernfs:
scx enable/disable cgroup rmdir cpu.weight write
------------------ ------------ ----------------
cgroup_lock()
percpu_down_write(rwsem)
cgroup_lock()
kernfs_get_active()
percpu_down_read(rwsem)
kernfs_drain()
The enable path waits for the rmdir to release cgroup_mutex. The rmdir,
deactivating the cpu controller's files, waits in kernfs_drain() for the
write's active reference. The write, in scx_group_set_weight(), waits for
the rwsem behind the pending writer.
Take cgroup_lock() first. The set_* paths take no cgroup locks inside the
read side, so a pending write-lock then only waits for read sections that
always run to completion, and no dependency from the rwsem back to
cgroup_mutex remains. |
