| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| In the Linux kernel, the following vulnerability has been resolved:
bnxt_en: fix DMA mapping length for padded small packets
Stefan Fleischmann reported Intel IOMMU DMA Read faults on BCM57412
NetXtreme-E NICs when transmitting packets on VLAN/macvlan interfaces:
DMAR: [DMA Read NO_PASID] Request device [18:00.0] fault addr 0xfc499000
[fault reason 0x06] PTE Read access is not set
bnxt_en 0000:18:00.0 eno1np0: Abandoning msg {0xb4 0x41a} len: 0 due to firmware status: 0x2000001
...
NETDEV WATCHDOG: eno1np0 (bnxt_en): transmit queue 0 timed out
The fault address (0xfc499000) is on an exact 4KB page boundary,
pointing to a DMA read buffer overrun.
In bnxt_start_xmit(), packets smaller than BNXT_MIN_PKT_SIZE (52 bytes),
such as 42-byte untagged ARP frames, are padded:
if (length < BNXT_MIN_PKT_SIZE) {
pad = BNXT_MIN_PKT_SIZE - length;
if (skb_pad(skb, pad))
goto tx_kick_pending;
length = BNXT_MIN_PKT_SIZE;
}
mapping = dma_map_single(&pdev->dev, skb->data, len, DMA_TO_DEVICE);
...
dma_unmap_len_set(tx_buf, len, len);
However, 'len' was initialized earlier to skb_headlen(skb) (e.g. 42 bytes)
and is left unadjusted after padding. Consequently, dma_map_single() and
dma_unmap_len_set() map and track only 42 bytes.
Later, the hardware TX buffer descriptor is programmed with the padded length:
txbd->tx_bd_len_flags_type =
cpu_to_le32(((len + pad) << TX_BD_LEN_SHIFT) | flags |
TX_BD_FLAGS_PACKET_END);
The NIC DMA engine is thus instructed to read 52 bytes from a region where
only 42 bytes were DMA-mapped. If skb->data ends near the boundary of a 4KB
page (within 'pad' bytes of the next page), the hardware DMA read overruns
into the unmapped adjacent page, triggering an IOMMU fault.
This issue was exposed after commit 447cbe95ebb9 ("vlan: fix skb_under_panic
and races when toggling HW VLAN offload") because reserving extra VLAN
headroom rounded LL_RESERVED_SPACE from 48 up to 64 bytes, shifting skb->data
offsets and potentially causing small frames to land right against page
boundaries.
Fix this by using skb_put_padto(skb, BNXT_MIN_PKT_SIZE) in
bnxt_start_xmit(), before skb_shinfo(skb)->nr_frags is sampled,
because padding might linearize the skb.
This ensures skb->len and skb_headlen(skb) reflect the padded size so
that dma_map_single() maps the full buffer and the descriptor length is
consistent. This also removes the temporary 'pad' variable and masking logic.
The padding is done after the SW USO branch, otherwise
bnxt_sw_udp_gso_xmit() would account the padding as UDP payload
and send it in extra segments.
Note that SW USO segments are still not padded: with IPv4, a segment
carrying less than 10 bytes of UDP payload (small gso_size, or a short
last segment) is sent as a frame shorter than BNXT_MIN_PKT_SIZE.
This is a separate issue, left for a followup patch. |
| In the Linux kernel, the following vulnerability has been resolved:
net/mlx5e: Prevent stale XSK buffer release on refill retry
When an XDP redirect to an AF_XDP socket fails because its RX ring is
full, the XSK core frees the buffer. During the subsequent batched refill
of a legacy cyclic RQ, mlx5e also releases the WQE's XSK buffer before
allocating a replacement. If that refill succeeds only partially, a WQE
left without a replacement retains its old buffer pointer.
The buffer can meanwhile be allocated to another WQE. A later refill
retry can then free the live buffer through the stale pointer and publish
the same UMEM frame twice.
Mark the WQE as released immediately after the driver-side free. The flag
is already cleared when a replacement buffer is assigned, so refill
retries no longer release stale pointers.
The failure is silent and produces no kernel warning or splat. A
standalone legacy cyclic-RQ zero-copy libxsk reproducer, using 64-byte UDP
traffic offered at 12 Mpps, detected it: stock stopped after 2,854,914
packets in 4.094 seconds, with 4,542 xdp_rx_ring_full events and 64
ownership/double-publication errors. With this change it processed
356,904,225 packets in 30 seconds despite 571,405 xdp_rx_ring_full events,
with no ownership or data errors. |
| In the Linux kernel, the following vulnerability has been resolved:
net/mlx5e: Prevent stale XSK buffer release on MPWQE refill retry
With AF_XDP on a striding RQ, mlx5e defers releasing XSK buffers until
an MPWQE is refilled. If XSK allocation then returns -ENOMEM,
actual_wq_head is not advanced and a later NAPI poll retries the same
WQE.
mlx5e_free_rx_mpwqe() leaves each released slot marked as releasable. On
retry it can therefore call xsk_buff_free() again through stale pointers
after the frames have returned to the XSK pool and been reallocated.
Set all skip_release_bitmap bits in the common error path of
mlx5e_xsk_alloc_rx_mpwqe(). This matches mlx5e_alloc_rx_mpwqe(). A
successful allocation already clears the bitmap after replacing every
buffer, so retries become idempotent without changing the success path.
Fault injection forced three consecutive failures for one selected MPWQE.
Both an early allocation failure and a partial 8-of-16-buffer unwind
released the original 16 XSK buffers only once. Each error left a full
bitmap, the following NAPI retry skipped the release, and a later
successful allocation cleared it. A 20-second AF_XDP zero-copy pressure
run exercised 1,575,262 buffer allocation failures without invalid
descriptors, WQE errors, or kernel warnings. |
| In the Linux kernel, the following vulnerability has been resolved:
ipv4: udp: Create exceptions before socket matching
Currently, when ICMP Fragmentation Needed and Redirect Message packets
are locally delivered and quote a UDP packet, a FIB nexthop exception
(FNHE) is only created if the kernel can match the UDP packet to an
existing socket.
This behavior allows off-path attackers to conduct a side-channel attack
on the FNHE cache in order to discover the ephemeral port used by a
connected UDP socket.
Commit 6457378fe796 ("ipv4: use siphash instead of Jenkins in
fnhe_hashfun()") and commit 67d6d681e15b ("ipv4: make exception cache
less predictible") tried to mitigate such attacks by making it harder
for attackers to discover hash collisions in the FNHE cache and by
randomizing the number of exceptions a hash bucket can hold,
respectively. Unfortunately, both of the mitigations can be bypassed.
Instead, mitigate such attacks by always creating a FNHE, even before
trying to find a matching socket. Do that by calling ipv4_update_pmtu()
and ipv4_redirect(), the helpers used when the quoted packet did not
originate from a socket.
This means that guesses (right or wrong) from an off-path attacker will
always result in a FNHE being created or updated in the cache that the
attacker can observe.
Pass an oif of 0, in a similar fashion to icmp_err(). This is also the
oif used by the socket path for sockets that are not bound to a device.
Note that this does not allow attackers to create FNHEs that they could
not create before, as both helpers can already be reached with little to
no validation. For example, by sending an ICMP error that quotes an ICMP
Echo Reply or one that quotes a UDP source port that matches a wildcard
socket.
Also note that in the good case (matched socket) the above scheme comes
at the cost of an extra route lookup, as the no socket helpers perform
their own lookup before the one performed by ipv4_sk_update_pmtu() /
ipv4_sk_redirect(). When the two resolve to different nexthops, it also
results in two exceptions being created for the same destination IP. One
in the FNHE cache of the nexthop resolved by the no socket helpers and
another in the FNHE cache of the nexthop used by the socket. |
| In the Linux kernel, the following vulnerability has been resolved:
cipso: adjust cached option offsets when removing CIPSO
cipso_v4_skbuff_delattr() removes the CIPSO bytes but does not adjust
cached offsets for options that follow them.
For example, with a valid 10-byte CIPSO option followed by a seven-byte
RR option, parsing records rr at offset 30. Removing CIPSO moves RR to
offset 20, while the cached offset remains 30. Consumers such as
ip_forward_options() and __ip_options_echo() then access the wrong bytes;
the latter may interpret packet data as the option length and copy it into
fixed-size option storage.
Mirror cipso_v4_delopt() and subtract cipso_len from the srr, rr, ts and
router_alert offsets when they follow CIPSO. cipso_len is the distance
the first memmove() shifts those options. The later header move and
network-header reset relocate the bytes and their offset base together. |
| In the Linux kernel, the following vulnerability has been resolved:
wireguard: noise: reject response consumption after intermediate initiation
Two threads begin processing the identical response message, received
twice. The first thread, A, runs. While it's running, the second one,
B, gets partway through, and during that slow calculation, or even while
blocking on down_write(), A completes and then also a handshake
initiation that's already been queued up runs in thread C, which itself
takes that same down_write(). The handshake initiation creation
succeeds, and sets the state back to waiting-for-response, and calls
up_write(), at which point thread B resumes, because either its finished
its calculations or was finally allowed to acquire down_write(). Thread
B then copies the state back to the peer, and begins a new session,
using that state, which is the same session as the one made in thread A.
Thread A Thread B Thread C
down_read()
sA = handshake->state
memcpy(cA, handshake->crypto)
up_read()
if (sA != 1)
goto fail
slow_crypto(cA)
down_read()
sB = handshake->state
memcpy(cB, handshake->crypto)
up_read()
if (sB != 1)
goto fail
slow_crypto(cB)
down_write()
if (sA != handshake->state)
goto fail
memcpy(handshake->crypto, cA)
handshake->state = 2
up_write()
down_write()
if (handshake->state != 2)
goto fail
derive_session(handshake->crypto)
up_write()
down_write()
slow_crypto(handshake->crypto)
handshake->state = 1
up_write()
down_write()
if (sB != handshake->state)
goto fail
memcpy(handshake->crypto, cB)
handshake->state = 2
up_write()
down_write()
if (handshake->state != 2)
goto fail
derive_session(handshake->crypto)
up_write()
This seems basically impossible to hit in a meaningful way in practice,
but ensure that it absolutely cannot happen by comparing the ephemeral
private key that's on the stack with the latest one that the peer's
handshake state has. |
| In the Linux kernel, the following vulnerability has been resolved:
strparser: make sure __strp_recv isn't running before tearing down the parser
The comment above strp_done() claims that if strp is stopped,
strp_recv will no longer be called. That's only true if the caller has
a mechanism (eg locking) to guarantee that strp_recv() calls that
started before strp_stop() have completed by the time we call
strp_done().
This adds a dummy lock_sock/release_sock pair to guarantee that any
in-flight strp_recv() (which runs under either bh_lock_sock or
lock_sock, depending if it's called from ->sk_data_ready or strp_work)
has completed.
Only espintcp can be affected by this race condition, but this patch
makes sure no future user of strp can have the bug.
This could crash on strp->sk ("general mode" of strp), but this mode
has been here for 9 years and has never been used. It'll be gone soon,
no point worrying about it.
Switch the stopped/paused/etc bits to u8's to avoid races between
strp_stop() and strp_pause/unpause().
A reproducer has been published and turns the possible UAF into an
exploit [1]. |
| In the Linux kernel, the following vulnerability has been resolved:
tipc: destroy topsrv workqueues before closing connections
tipc_topsrv_stop() closed subscriber connections while the topology
server's send and receive workqueues were still running. Socket
callbacks and subscription events could then queue more work, and
in-flight send/recv work could drop the last connection reference
during the conn_idr walk.
That race produced several teardown failures: refcount_t addition on
0 from conn_get() on a connection whose release was blocked on
idr_lock, a subsequent use-after-free in tipc_conn_close(),
queue_work() on an already destroyed workqueue from the listener
data-ready callback, and an RCU stall in tipc_topsrv_exit_net()
while the walk spun under idr_lock.
Clear srv->listener under idr_lock so it acts as a shutdown flag,
skip queue_work() once it is NULL, destroy the workqueues to flush
in-flight work, and only then close the remaining connections.
Refuse tipc_conn_lookup() after that flag is cleared, and drop
idr_lock when the teardown walk finds no connection, so an
in-flight subscription event cannot pin idr_in_use while the
walk holds the lock.
v3 supersedes the narrower in-thread diff Tung Quang Nguyen
posted on 2026-09-23. It keeps his teardown order and adds the
lookup refusal and the empty-idr unlock. |
| In the Linux kernel, the following vulnerability has been resolved:
tipc: hold a reference to nodes found by link name
tipc_node_find_by_name() returns a node after dropping its RCU read lock
without taking a reference. The LINK_SET, LINK_GET and LINK_RESET_STATS
handlers then lock and access the node, racing with timer-driven cleanup of
a down peer. Generic netlink serialization does not exclude the node
timer.
The following interleaving can leave a handler using a freed node:
CPU 0: find the node under RCU and release the node read lock
CPU 1: tipc_node_timeout() clears the links and unlinks the down node
CPU 1: drop the list and timer references, queuing tipc_node_free()
CPU 0: leave the RCU read-side critical section
CPU 1: complete the grace period and free the node
CPU 0: acquire the node lock through the stale pointer
LINK_SET also uses the node's media address after releasing the node lock,
when passing queued packets to tipc_bearer_xmit().
KASAN on v7.3-rc5 reported:
BUG: KASAN: slab-use-after-free in _raw_read_lock_bh
Write of size 4 at addr ffff88807e06f008 by task poc/92
Call Trace:
_raw_read_lock_bh kernel/locking/spinlock.c:287
tipc_nl_node_set_link net/tipc/node.c:2475
genl_family_rcv_msg_doit net/netlink/genetlink.c:1114
genl_rcv_msg net/netlink/genetlink.c:1209
netlink_rcv_skb net/netlink/af_netlink.c:2575
Allocated by task 0:
tipc_node_create net/tipc/node.c:539
tipc_node_check_dest net/tipc/node.c:1196
tipc_disc_rcv net/tipc/discover.c:252
Freed by task 92:
kfree mm/slub.c:6801
rcu_core kernel/rcu/tree.c:2919
Last potentially related work creation:
__call_rcu_common.constprop.0 kernel/rcu/tree.c:3181
tipc_node_timeout net/tipc/node.c:814
Acquire a reference to the selected node with kref_get_unless_zero() before
leaving RCU, returning NULL if the node has already been released. Release
that reference on every caller exit after the last node access, including
transmission in LINK_SET. Keep the existing link lookup order and locking
so concurrent link removal still takes the existing error paths. |
| In the Linux kernel, the following vulnerability has been resolved:
fsverity: RCU-delay the freeing of struct fsverity_info
__fsverity_get_info() calls rhashtable_lookup_fast(), which just uses
rcu_read_lock() and doesn't directly synchronize with
fsverity_remove_info().
For the same inode this isn't a problem: its fsverity_info is removed
only at inode eviction time or upon failure to enable verity, when the
inode no longer needs its fsverity_info and it will no longer be
accessed via that inode.
However, this is broken and can cause a use-after-free for concurrent
__fsverity_get_info() for *different* inodes. Those rely on following
fsverity_info::rhash_head in the rhashtable under rcu_read_lock() only.
They also use fsverity_info::inode to do the key comparison.
Fix this by RCU-delaying the freeing of 'struct fsverity_info' after
it's been removed from the rhashtable. |
| In the Linux kernel, the following vulnerability has been resolved:
gpio: mpsse: fix race when arming the IRQ poll worker
gpio_mpsse_irq_enable() arms the poll worker before publishing it:
schedule_work() runs before the worker is added to priv->workers. If
gpio_mpsse_disconnect() walks the list in that window it misses the
worker, and once disconnect returns, the USB core frees mpsse_priv
while the orphaned gpio_mpsse_poll() work keeps accessing it, causing
a use-after-free.
Fix this by publishing and arming the worker in one irq_spin
critical section. Teardown walks the same list under irq_spin, so a
worker found on the list is guaranteed to be armed, and
cancel_work_sync() handles it whether it is queued or running.
A worker armed after the disconnect walk would still be missed, so
also set a new priv->dying flag under irq_spin before the teardown
walk, and check it in the same critical section, freeing it via
kfree_rcu() instead when the device is going away, since kfree()
would be called with the IRQ core's raw spinlock held and may sleep
on PREEMPT_RT. schedule_work() is safe to call with irq_spin held,
and the next probe gets a fresh mpsse_priv, so the flag never needs
to be cleared.
This issue was found by an in-house static analysis tool. |
| In the Linux kernel, the following vulnerability has been resolved:
gve: fix XSK buffer leak on error descriptor
When the error bit is set in the RX completion descriptor, the buf_state
and its attached buffer should be freed. In the case of AF_XDP ZC, the
XSK buffer was not freed, leading to a leak.
This issue was caught by LLM. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/net: populate CONNECT filter fields for 24-byte sockaddr_in6
io_connect_bpf_populate() fills connect.port and connect.v6_addr for
AF_INET6 only when addr_len >= sizeof(struct sockaddr_in6), i.e. 28
bytes.
The IPv6 connect paths accept SIN6_LEN_RFC2133 (24 bytes), which
already holds sin6_port and sin6_addr. io_connect() passes addr_len to
the socket layer unchanged.
At addr_len 24 the filter sees family AF_INET6 with port and address
zero, while the kernel connects to the real destination. A filter that
denies an IPv6 port or address is bypassed. Denying connect.port == P
rejects [::1]:P with -EACCES at addr_len 28 but connects at addr_len 24.
Use SIN6_LEN_RFC2133 as the threshold. sin6_scope_id stays zero when
absent.
The threshold mismatch was found by an AI coding assistant (Claude Code)
comparing this hook against the net/ipv6 length checks on my
instructions; it also drafted the reproducer and this patch. I verified
the bypass and the fix in a VM on 7.3-rc6 and that 7.2 has the same
code. |
| In the Linux kernel, the following vulnerability has been resolved:
net/smc: protect clcsock lifetime in smc_getname
smc_getname() dereferences smc->clcsock without holding
clcsock_release_lock. Link-group termination can release the CLC socket
through smc_close_active_abort() while the SMC socket is still open,
for example after shutdown(SHUT_WR).
A getsockname() caller can load smc->clcsock, then the termination worker
can clear the pointer and call sock_release() before the caller accesses
clcsock->ops or invokes getname(). This causes a use-after-free; if the
worker clears the pointer before the load, it causes a NULL dereference.
The syscall's file reference keeps the SMC socket alive but does not
prevent asynchronous release of its CLC socket.
KASAN reported the following with test-only timing instrumentation:
BUG: KASAN: slab-use-after-free in smc_getname+0x19e/0x1b0
Read of size 8 at addr ffff888109abb4e0 by task poc/103
Call Trace:
smc_getname+0x19e/0x1b0
do_getsockname+0xe5/0x170
__sys_getsockname+0x8c/0x100
Allocated by task 95:
sock_alloc_inode+0x1e/0x280
sock_alloc+0x3d/0x240
__sock_create+0x7e/0x430
smc_create+0x121/0x240
Freed by task 0:
kmem_cache_free+0xcc/0x340
rcu_core+0x50a/0x1850
Last potentially related work creation:
evict+0x446/0x6c0
smc_clcsock_release+0xa8/0xd0
smc_close_active_abort+0x26a/0x3a0
__smc_lgr_terminate.part.0+0x137/0x2e0
Hold clcsock_release_lock across the pointer check and the getname()
callback to serialize with smc_clcsock_release(). Return -EBADF if the
CLC socket has already been released, preserving the existing peer
state check and the callback's return value otherwise. |
| In the Linux kernel, the following vulnerability has been resolved:
net: skbuff: don't leave stale bytes in skb_copy_and_csum_bits()
When skb_copy_and_csum_bits() reaches unreadable frags it returns 0
after copying only the linear part, and the rest of the caller's buffer
is left as it was. The callers copy into a buffer that is about to go
out on the wire: an ICMP error quoting the offending packet, or a
driver's TX bounce buffer in skb_copy_and_csum_dev(). Neither buffer
is zeroed beforehand, so whatever was in memory there gets sent.
Zero the part of the buffer we didn't fill. The checksum usually
won't match the data any more, so the receiver will usually drop the
packet, but either way it no longer carries anything it shouldn't.
Only zero for a positive @len, a negative one from a broken caller must
not turn into a huge memset(). |
| In the Linux kernel, the following vulnerability has been resolved:
riscv: errata: select RISCV_ALTERNATIVE_EARLY for THEAD GHOSTWRITE
errata_probe_ghostwrite() only marks the CPU as vulnerable when it is
called at the RISCV_ALTERNATIVES_EARLY_BOOT stage, and the mitigation
in ghostwrite_enable_mitigation() only takes effect if the vulnerable
state was recorded before the boot alternatives are patched.
The early boot stage is only executed when CONFIG_RISCV_ALTERNATIVE_EARLY
is enabled, which currently is only selected by ERRATA_THEAD_MAE. In a
kernel built with ERRATA_THEAD_GHOSTWRITE=y but ERRATA_THEAD_MAE=n, the
early alternatives never run, so ghostwrite_state is never set to
VULNERABLE: ghostwrite_enable_mitigation() returns false, xtheadvector
remains enabled, and sysfs reports "Not affected" on a vulnerable
T-Head C9xx CPU.
Select RISCV_ALTERNATIVE_EARLY from ERRATA_THEAD_GHOSTWRITE so that the
vulnerability is detected in the early boot stage and the default
mitigation is applied. |
| In the Linux kernel, the following vulnerability has been resolved:
riscv: futex: untag the user pointer before the atomic access
access_ok() checks untagged_addr(uaddr), but arch_futex_atomic_op_inuser()
and futex_atomic_cmpxchg_inatomic() hand the raw uaddr to the inline asm
amoswap/amoadd/.../lr.w+sc.w through the "+m" (*uaddr) operand. When the
tagged address ABI is enabled (CONFIG_RISCV_ISA_SUPM, prctl
PR_SET_TAGGED_ADDR_CTRL with PMLEN != 0), those two addresses differ: a
user pointer whose top PMLEN bits hold a tag passes access_ok() after
untagging, but the atomic is then performed on the still-tagged raw
address.
Supervisor-mode data accesses are not subject to the U-mode pointer
masking (that is governed by menvcfg.PMM, which the kernel does not set),
so hardware does not strip the tag for the kernel's own access. With
Sv57 and PMLEN=16 the tag bits overlap the canonical-address bits, so a
tagged pointer can name a canonical kernel virtual address (e.g. in the
linear map) whose untagged form is a valid user address. An unprivileged
process can thus make FUTEX_WAKE_OP perform an atomic read-modify-write on
an arbitrary kernel address, with the matching FUTEX_OP_CMP_* result
serving as a read oracle.
get_user()/put_user()/raw_copy_{to,from}_user() already untag the pointer
after the access_ok() check; do the same in the futex helpers so the
atomic operates on the address that was actually validated. |
| In the Linux kernel, the following vulnerability has been resolved:
Bluetooth: hci_sync: Fix command skb lifetime during scan setup
During PA sync scan setup, hci_le_set_ext_scan_param_sync() gets the
address used for the PA_LINK lookup from hci_sent_cmd_data(). The
returned pointer borrows storage from sent_cmd or req_skb, without
holding a reference to either skb.
The scan worker runs on req_workqueue while hci_cmd_work() runs on the
separate device workqueue. After the scan worker loads sent_cmd, the
command worker can free and replace it before hci_sent_cmd_data()
dereferences the skb. The command payload can also be freed between
returning from the helper and comparing the address. The connection
lookup's RCU critical section does not protect the command skb.
KASAN reported:
BUG: KASAN: slab-use-after-free in hci_sent_cmd_data+0x27c/0x2f0
Workqueue: hci0 hci_cmd_sync_work
Call Trace:
hci_sent_cmd_data+0x27c/0x2f0
hci_le_set_scan_param_sync+0x43a/0x850
hci_passive_scan_sync+0xb09/0x1460
hci_update_passive_scan_sync+0x47c/0x6e0
hci_le_pa_create_sync+0x200/0xa70
hci_cmd_sync_work+0x13c/0x290
Allocated by task 95:
skb_clone+0x13f/0x340
hci_cmd_work+0x2a9/0x7e0
Freed by task 95:
kmem_cache_free+0xba/0x3a0
hci_cmd_work+0x29c/0x7e0
Use hdev->lock to serialize the address snapshot with sent_cmd
replacement. Protect req_skb replacement and completion cleanup with
the same lock because the helper falls back to that skb. Copy the
address before releasing the lock and use the copy for the existing
RCU-protected connection lookup. Release the mutex before sending or
waiting for commands, preserving the lookup and completion ordering. |
| In the Linux kernel, the following vulnerability has been resolved:
Bluetooth: ISO: Reject concurrent BIS listener setup
iso_sock_listen() drops the socket lock before calling iso_listen_bis()
to preserve the hdev->lock ordering. Two concurrent listen() calls on
the same broadcast socket can therefore both observe BT_BOUND. Caller
A can attach connA and release the locks, then caller B can acquire
them and attach connB before A publishes BT_LISTEN.
__iso_chan_add() checks whether connB already has a socket, but not
whether the socket already has a connection. The second attachment
overwrites iso_pi(sk)->conn, leaking the connA reference and leaving
connA->sk intact. Socket teardown follows connB, so later events on
connA can cause a stale socket access.
Recheck BT_BOUND and require iso_pi(sk)->conn to be NULL after taking
both the HCI device and socket locks in iso_listen_bis(). The connection
check also covers the interval between attachment and publication of
BT_LISTEN. Reject a competing setup with -EBADFD through the existing
unlock path before allocating another connection.
A recorded KASAN report from an instrumented kernel shows:
BUG: KASAN: slab-use-after-free in iso_sock_hold+0x119/0x1f0
Call Trace:
iso_sock_hold+0x119/0x1f0
iso_conn_del+0xe4/0x270
hci_disconn_complete_evt+0x351/0x8c0
hci_event_packet+0x71b/0xb20
hci_rx_work+0x293/0x730
Allocated by task 93:
sk_prot_alloc+0x113/0x220
sk_alloc+0x2b/0x6d0
bt_sock_alloc+0x29/0x370
iso_sock_alloc.constprop.0+0x19/0x300
iso_sock_create+0x94/0x100
Freed by task 93:
kfree+0x121/0x3c0
__sk_destruct+0x42b/0x540
iso_sock_release+0x29d/0x340
__sock_release+0xa1/0x260
sock_close+0x10/0x20 |
| In the Linux kernel, the following vulnerability has been resolved:
Bluetooth: SCO: serialise sco_conn lifetime against sco_recv_scodata()
sco_recv_scodata() upgrades the weak hcon->sco_data back-pointer to a
strong reference under hdev->lock:
hci_dev_lock(hdev);
hcon = hci_conn_hash_lookup_handle(hdev, handle);
...
conn = sco_conn_hold_unless_zero(hcon->sco_data);
hci_dev_unlock(hdev);
but the pointer is cleared from the other side without that lock. When
the last sco_conn reference is dropped, sco_conn_free() ran
conn->hcon->sco_data = NULL;
with no hdev->lock held, so the RX path could read hcon->sco_data and
call kref_get_unless_zero() on an sco_conn that was concurrently freed:
BUG: KASAN: slab-use-after-free in sco_conn_hold_unless_zero+0xbe/0x160
Write of size 4 by task kworker/u17:0
Workqueue: hci0 hci_rx_work
Call Trace:
sco_conn_hold_unless_zero+0xbe/0x160
sco_recv_scodata+0x13f/0x490
hci_rx_work+0x3af/0x730
kref_get_unless_zero() only guards against a zero refcount, not against
the backing memory already being freed.
Give hcon->sco_data an actual reference on the sco_conn it points to, so
the object cannot be freed while the pointer is still installed, and only
clear and drop it from sco_conn_del(), which runs under hdev->lock (its
callers, sco_connect_cfm() and sco_disconn_cfm(), hold it). The reader in
sco_recv_scodata() also takes hdev->lock, so the store and the read are
now serialised: the RX path either observes NULL or a reference that is
guaranteed to stay valid until it drops its own.
sco_conn_add() no longer hands its kref_init() reference to the caller as
the connection's only reference; that initial reference is the one owned
by hcon->sco_data, and callers get their own via sco_conn_hold().
Because the sco_conn can now outlive the socket (the association keeps it
alive until the link goes down), the hci_conn can no longer be dropped
from sco_conn_free() without regressing socket close: closing a connected
SCO socket must still tear the link down. Move hci_conn ownership to the
socket instead: __sco_chan_add() takes an hci_conn reference and it is
released from sco_chan_del() and sco_sock_destruct(), exactly once. The
sco_conn no longer owns an hci_conn reference, so sco_conn_add() stops
consuming one and sco_connect() drops the hci_connect_sco() reference on
every path; sco_connect_cfm() no longer needs its hci_conn_hold() dance. |
