CVE Feed

In the Linux kernel, the following vulnerability has been resolved: fprobe: Terminate the fgraph_data list when the reservation is not filled fprobe_fgraph_entry() reserves shadow stack space for every fprobe with an exit handler, but only fills it for those whose entry handler returns 0. fgraph_reserve_data() does not clear the area, so fprobe_return() parses the unused tail as headers left over from an earlier call, and an exit handler can run twice or despite its entry handler asking to skip it. Write a zero word after the last entry to terminate the walk. A zeroed slot does not decode to a NULL fprobe on the arches that encode the header into one unsigned long, since arch_decode_fprobe_header_fp() ORs in FPROBE_HEADER_MSB_PATTERN, so make read_fprobe_header() return NULL for a zeroed slot.

In the Linux kernel, the following vulnerability has been resolved: perf/core: Fill branch entries with a single assignment perf_clear_branch_entry_bitfields() clears the bitfields of struct perf_branch_entry one by one and leaves from/to alone, since callers overwrite those straight away. The list has to be kept in sync with the struct by hand and has already fallen behind: new_type and priv were added to perf_branch_entry and never added here. Only BRBE writes those two, and neither for every record. brbe_set_perf_entry_type() leaves new_type alone for a branch type it does not recognise, and priv is not set for source-only records. arm_pmuv3.c allocates the per-CPU branch stack with kmalloc(), so such a record reaches userspace with whatever the slot held: uninitialised kmalloc() data on the first pass over the buffer, the previous record's values after that. Nothing under arch/x86/events/ writes either field, so only arm64 is affected. Assign the whole entry at each site instead. Everything not named is then zero, and there is no list to keep in sync. The bitfields add up to exactly 64 bits, so the struct has no padding to leave undefined. perf_clear_branch_entry_bitfields() has no callers left, so remove it. perf_entry_from_brbe_regset() assigns an empty literal instead, since it fills from/to conditionally. PERF_BR_SPEC_NA is 0, so dropping the explicit spec assignment changes nothing.

In the Linux kernel, the following vulnerability has been resolved: vlan: ensure sufficient headroom in vlan_dev_hard_header() Callers that only reserve ETH_HLEN or less (such as llc_alloc_frame()), or skbs allocated before dynamic device/headroom changes (e.g. toggling VLAN_FLAG_REORDER_HDR or bonding/team switching slaves), can reach vlan_dev_hard_header() with insufficient headroom and trigger skb_under_panic(). Use skb_cow_head() in vlan_dev_hard_header() when VLAN_FLAG_REORDER_HDR is not set to ensure sufficient headroom for the VLAN header(s) and the underlying device hard header. Use READ_ONCE() to read dev->hard_header_len and dev->needed_headroom as they can be updated concurrently under RTNL (e.g. in vlan_transfer_features()) while vlan_dev_hard_header() runs locklessly on the transmit path. Also avoid LL_RESERVED_SPACE(dev) here so that the extra HH_DATA_MOD alignment padding does not trigger unnecessary pskb_expand_head() reallocations on inner stacked VLAN devices after the outer VLAN header has been pushed.

In the Linux kernel, the following vulnerability has been resolved: llc: fix skb UAF and leaks on llc_mac_hdr_init() failure In llc_conn_ac_resend_i_xxx_x_set_0_or_send_rr(), if llc_mac_hdr_init() fails, kfree_skb(skb) is called instead of kfree_skb(nskb). This leaks the newly allocated nskb, reads from the freed skb via LLC_I_GET_NR(pdu), and double-frees skb when llc_conn_state_process() drops its reference. In llc_sap_action_send_xid_r() and llc_sap_action_send_test_r(), nskb is leaked if llc_mac_hdr_init() returns an error. Free nskb in all three error paths.

In the Linux kernel, the following vulnerability has been resolved: gve: DQO: fix header length used by gve_can_send_tso() for UDP GSO gve_can_send_tso() computes how many buffers each segment of a GSO packet would span, and for this it needs the length of the headers that the device replicates in front of every segment. It unconditionally uses skb_tcp_all_headers(), which reads the doff field of the TCP header. SKB_GSO_UDP_L4 packets have no TCP header: tcp_hdrlen() then reads one byte of the UDP payload, and header_len can be anything in [0, 60] instead of the transport offset plus the eight bytes of the UDP header that gve_prep_tso() programs into the TSO context descriptor. A wrong header length shifts all the segment boundaries computed in the loop, so the number of buffers per segment can be over or under estimated. In the first case, GSO is needlessly disabled for this packet by gve_features_check_dqo() and the stack has to segment it. In the second case, the driver hands the device a packet whose segments span more than GVE_TX_MAX_DATA_DESCS buffers. Use the UDP header length for SKB_GSO_UDP_L4 packets, matching what gve_prep_tso() does.

In the Linux kernel, the following vulnerability has been resolved: net: flush skb_defer_nodes in dev_cpu_dead() When a CPU goes offline, dev_cpu_dead() drains its softnet queues (completion_queue, output_queue, poll_list, process_queue, and input_pkt_queue), but leaves net_hotdata.skb_defer_nodes untouched. If oldcpu goes offline while holding pending skbs in its skb_defer_nodes lists (e.g. below the sysctl_skb_defer_max >> 1 IPI threshold, or if the IPI races with CPU teardown), those skbs remain stranded until oldcpu is brought back online. If any of these skbs hold page_pool fragments, page_pool_destroy() will stall indefinitely waiting for inflight pages to be returned when a netdev or driver is torn down while oldcpu is offline. Additionally, if smp_call_function_single_async() fails in kick_defer_list_purge() because the target CPU went offline, reset defer_ipi_scheduled to 0 so future IPI kicks are not blocked when the CPU comes back online. Also, if oldcpu was the last online CPU on its NUMA node, drain that node's slot across all CPUs so no skbs deferred from that node remain stranded on idle remote CPUs (or if the node itself is subsequently offlined). Finally, in skb_attempt_defer_free(), re-check cpu_online(cpu) and whether the caller migrated CPUs after llist_add(), flushing the node list if so, to close the preemption TOCTOU race against CPU/node teardown.

In the Linux kernel, the following vulnerability has been resolved: net: ipv6: keep room for the mac header in dst_dev_overhead() The seg6, ioam6 and rpl lwtunnels size their skb_cow_head() request as the length they are about to push plus dst_dev_overhead(), then push the new headers and rebuild the mac header below them with skb_mac_header_rebuild(). That rebuild needs skb->mac_len of headroom, but dst_dev_overhead() leaves LL_RESERVED_SPACE() of the egress device, 16 bytes for plain Ethernet. Where the mac header is longer than that, as it is on ingress through a VLAN device with reorder_hdr off, the rebuild runs out of room: skb_set_mac_header(skb, -skb->mac_len) computes a negative offset, stores it unchecked in the u16 skb->mac_header, and the memmove that follows writes skb->mac_len bytes about 64 KB past skb->head. Forwarding plain ping6 traffic through such a device reproduces it on all five seg6 encapsulation modes and on the rpl and ioam6 inline paths; skb->mac_header comes back as 65534 on a 704-byte head. Return the larger of the two. The helper already returns skb->mac_len when it has no dst, so this only makes the other branch agree, and it covers every caller rather than each call site in turn.

In the Linux kernel, the following vulnerability has been resolved: net: stmmac: clear stale buf->page after recycling on skb build failure In stmmac_rx(), when napi_build_skb() fails the descriptor page is recycled back to the page pool with page_pool_recycle_direct(), but buf->page is left pointing at the recycled page, unlike every other consumption site in the function which clears the pointer after handing the page away. With the stale pointer stmmac_rx_refill() skips the replacement allocation and programs the already-recycled page back into the RX descriptor. Clear buf->page on the napi_build_skb() failure path to keep the buffer lifecycle consistent with the other consumption sites.

In the Linux kernel, the following vulnerability has been resolved: ip_gre: Reject enabling collect metadata through changelink ipgre_netlink_parms() can enable collect_md on an existing GRE, GRETAP or ERSPAN device. Unlike newlink, changelink does not enforce metadata tunnel uniqueness. Converting a non-metadata device can therefore replace the metadata receive entry for another device of the same type in the same netns. Deleting either device then clears the shared entry, breaking metadata receive lookup for the surviving device. If parameter validation fails after collect_md is set, deleting the modified device can also clear an entry it never owned. Reject enabling metadata mode in both changelink callbacks before any encapsulation or tunnel parameters are modified. Allow requests that repeat the metadata attribute on an existing metadata device.

In the Linux kernel, the following vulnerability has been resolved: nfc: st21nfca: validate received frame size st21nfca_hci_i2c_repack() trims a received frame at its EOF marker before removing byte stuffing. It then assumes the truncated frame contains the LLC header and two CRC bytes, and it unconditionally reads the byte after an escape marker. A malformed frame can place EOF immediately after the start marker or can end its data portion with an escape marker. The former leaves too few bytes for check_crc(), while the latter makes the unstuffing loop read past the current skb length. Require the minimum framing bytes both before and after unstuffing. Use separate input and output cursors while removing byte stuffing, and reject an escape marker without its encoded byte. This keeps malformed frames within the received frame boundary before CRC processing.

In the Linux kernel, the following vulnerability has been resolved: nfc: nfcmrvl: validate helper command length before pull The firmware download receive path removes the NCI data header and reads the helper command before validating the remaining packet length. A short frame can therefore reach the data access before the malformed packet is rejected. Validate the complete helper command length before stripping the NCI data header.

In the Linux kernel, the following vulnerability has been resolved: bpf, sockmap: Reject max_entries > INT_MAX in sock_map_alloc sock_map_alloc() only rejects max_entries == 0 and otherwise allows any u32 value. sock_map_free() then walks the sks[] array with a signed int iterator: int i; for (i = 0; i < stab->map.max_entries; i++) struct sock **psk = &stab->sks[i]; When a SOCKMAP is created with max_entries = 0xffffffff (UINT_MAX), the allocation of 32 GiB can succeed on large-memory hosts. During free the counter reaches 0x80000000, wraps to INT_MIN, is sign-extended by movslq and turned into a ~16 GiB negative offset from stab->sks, pointing far below the allocation. The faulting access is an xchg() write in sock_map_free(). Without KASAN, the same out-of-bounds write can fault on an unmapped vmalloc page or corrupt an unrelated allocation if that vmalloc address is populated. On a KASAN kernel with CONFIG_KASAN_VMALLOC=y, the shadow check for that address hits an unmapped shadow page and oopses first: BUG: unable to handle page fault for address: fffff521b59c5a00 RIP: 0010:kasan_check_range+0x107/0x190 Call Trace: sock_map_free+0x93/0x190 map_create+0x68d/0xb30 __sys_bpf+0x21e/0x2e70 Vmcore confirmed stab->map.max_entries == 0xffffffff, stab->sks == 0xffffc911ace2d000, and the faulting address sks + (s64)INT_MIN * 8 exactly at 0xffffc90dace2d000. The same buggy path is reached on the normal close()/bpf_map_free_deferred() path whenever such a map is destroyed. sock_map_alloc() used to bound its allocation size through bpf_map_charge_init(), but the bound was dropped when rlimit-based memory accounting was removed. Reject max_entries > INT_MAX at creation time so the signed iterator in sock_map_free() never sees a value that would overflow. Triggered by syzkaller and reproduced on both a 6.6-based KASAN kernel and the upstream v7.3-rc2 kernel.

In the Linux kernel, the following vulnerability has been resolved: bpf: Reject pkt arguments in mutating subprogs The verifier tracks changes in how PTR_TO_PACKET registers' bounds are modified across subprog boundaries. PTR_TO_PACKET registers are actually passed as PTR_TO_MEM, which is assumed valid for the entire call. This is not the case with packet memory, where a pskb_* call may invalidate its memory region. Reject BPF code that passes PTR_TO_PACKET pointers to subprogs that may mutate a packet. We cannot pass the pointer as a true PTR_TO_PACKET because we would also need to somehow pass the PTR_TO_PACKET_META or PTR_TO_PACKET_END to the subprog. Since we cannot avoid representing the pointer in the subprog as PTR_TO_MEM, only permit it if the subprog is guaranteed not to mutate the packet.

In the Linux kernel, the following vulnerability has been resolved: s390/debug: Fix NULL pointer dereference in debug_info_copy() When debug_register_static() fails, it clears areas, active_pages and active_entries but leaves the area bounds unchanged. Copying such an area, either by opening its view file or via debug_dump(), makes debug_info_copy() dereference the NULL pointers. Skip the copy loop when the source has no areas.

In the Linux kernel, the following vulnerability has been resolved: net: usb: sr9700: include receive overhead in the length check The receive fixup subtracts the Ethernet CRC from the reported packet length, but compares that payload length against the whole remaining receive buffer. The following copy starts after the three-byte header, and the cursor advance consumes both that header and the four-byte CRC. Require the payload to fit after SR_RX_OVERHEAD before copying it or advancing to the next packet. The loop already ensures that the remaining buffer is larger than the overhead, so the subtraction is safe. The issue was found by our static-analysis tool.

In the Linux kernel, the following vulnerability has been resolved: tg3: clean up PHYLIB resources on probe failure tg3_get_invariants() can register an MDIO bus and connect a PHY for USE_PHYLIB devices. If tg3_init_one() later fails, its common error path releases the mappings and netdev without undoing those PHYLIB resources. Disconnect the PHY and unregister the MDIO bus before the remaining teardown. Guard PHY cleanup with USE_PHYLIB to match tg3_phy_init(), and call tg3_mdio_fini() unconditionally to match tg3_mdio_init(). The existing IS_CONNECTED and MDIOBUS_INITED flags make both helpers safe when initialization only completed partially. This issue was identified during our ongoing static-analysis research while reviewing kernel code.

In the Linux kernel, the following vulnerability has been resolved: udp: remove a disconnected socket from the 4-tuple hash table A UDP socket bound to a specific address and port keeps its entry in the 4-tuple hash table after it is disconnected: sk binds to 127.0.0.1:21001 sk connects to 127.0.0.2:20001 // filed in the 4-tuple table sk disconnects, connect(AF_UNSPEC) // still filed, peer now 0.0.0.0:0 __udp_disconnect() takes a socket out of that table only as a side effect of ->rehash() or ->unhash(), and it skips ->rehash() when SOCK_BINDADDR_LOCK is set and ->unhash() when SOCK_BINDPORT_LOCK is set. commit 6996a2d2d0a6 ("udp: Unhash auto-bound connected sk from 4-tuple hash table when disconnected.") fixed the same end state for a wildcard-bound socket, by a path this one does not take. The entry is counted whether or not anything hits it. hash4_cnt on the hash2 slot stays raised for as long as the socket lives, so udp_has_hash4() keeps sending every packet for that address and port through the 4-tuple lookup first. On IPv6 it can also be hit. __udp_disconnect() does not clear sk_v6_daddr, so udp_v6_rehash() files the entry under the peer the socket was connected to with a zero dport, and inet6_match() compares that same field: a datagram from the former peer with a zero source port matches, and source port zero is accepted on receive. On IPv4 the peer is cleared, so a match would need a zero source address as well, which the routing layer rejects as martian. The stale sk_v6_daddr is a separate defect, not addressed here; removing the entry closes this path either way. The entry can also be relocated. __udp_disconnect() clears sk_bound_dev_if, so a subsequent SO_BINDTODEVICE calls ->rehash(), and because the receive address is still specific udp_lib_rehash() moves the entry instead of removing it, into the bucket that (rcv_saddr, num, 0, 0) hashes to -- a pure function of the address and port, so every socket reaching this state on one address and port collects in one bucket. The bucket cannot be chosen from outside, as udp_ehashfn() is seeded with a per-boot secret. This last one became reachable only with commit 644f9108f3a5 ("udp: Make rehash4 independent in udp_lib_rehash()"), which moved the hash4 handling out of a branch a disconnected socket does not take; the stale entry itself dates from the commit in Fixes. Take the socket out of the table before __udp_disconnect() runs, while it still matches how it was filed. This also reaches the wildcard case ahead of udp_lib_rehash()'s udp_unhash4() branch, leaving that branch unreachable from udp_disconnect(); removing it belongs in net-next. udp_disconnect() and udp_abort() are the only UDP entries into __udp_disconnect(), which is shared with raw, ping and l2tp sockets that are not struct udp_sock: ping_prot.obj_size is sizeof(struct inet_sock), so udp_hashed4() on one would read past the allocation.

In the Linux kernel, the following vulnerability has been resolved: eth: fbnic: Fix payload page pool error cleanup The payload page pool pointer contains an error pointer when its allocation fails. The cleanup path passes that error pointer to page_pool_destroy() instead of destroying the header page pool. This can dereference the error pointer and leave the header page pool allocated. Destroy the header page pool instead.

In the Linux kernel, the following vulnerability has been resolved: KVM: s390: Fix _gaccess_shadow_fault() In some circumstances, it is possible that the page of nested guest memory that is being shadowed is not present at all in the parent guest gmap. dat_entry_walk() will not find any leaf entry and return with -ENOENT, which will erroneously be propagated all the way to userspace. Fix by manually calling gmap_link() on the memory of the nested guest that is being shadowed if the mapping was not already present.

In the Linux kernel, the following vulnerability has been resolved: net/sched: cls_u32: fix manual hash table handle IDR aliasing A u32 hash table created with an explicit handle ('tc filter add ... handle 801: u32 divisor N') keys its IDR entry on the raw handle, while the destroy paths free it under handle2id(handle). The two key domains disagree for handles in the 0x800..0xFFF htid range: handle2id() folds them back into the auto-allocated id space (1..0x7FF). A manual table therefore leaves its raw-keyed IDR entry unreachable on delete (a permanent leak), and its delete can drop the idr entry of an unrelated live auto table. A later auto allocation can then hand out a handle that aliases the live manual table; u32_lookup_ht() first-match routes lookups and TCA_U32_LINK for that htid to the wrong table. Key the divisor-path alloc on handle2id(handle) so allocation and removal share one key domain. A manual handle that maps onto an id already in use is rejected with -ENOSPC, and auto allocation skips ids held by live manual tables. Conditions to recreate: ip link add test0 type dummy tc qdisc add dev test0 clsact tc filter add dev test0 ingress protocol ip pref 1 \ handle 801: u32 divisor 16 tc filter add dev test0 ingress protocol ip pref 2 u32 divisor 16 tc -d filter show dev test0 ingress | grep 'fh 801:' # unpatched: two live tables with handle 0x80100000 (the pref 2 root # hnode is auto-allocated id 1); patched: the auto hnode takes id 2. Also tested with a poc with a live u32 table on the block, add/delete a manual table 'handle 901: u32 divisor 1' twice; unpatched, the re-add fails with -ENOSPC because the raw key leaked on the first delete.

In the Linux kernel, the following vulnerability has been resolved: bpf: Avoid soft lockup in __htab_map_lookup_and_delete_batch() __htab_map_lookup_and_delete_batch() has no rescheduling point. The batch count bounds how many entries are copied out, not how many buckets are visited, so one BPF_MAP_LOOKUP_BATCH call can walk the map end to end. The empty-bucket fast path is worse: it stays inside a single rcu_read_lock() / bpf_disable_instrumentation() section for any run of consecutive empty buckets. That holds up on small maps, but it falls apart at scale. On a 144-CPU arm64 host running a CONFIG_PREEMPT_NONE kernel, periodic BPF_MAP_LOOKUP_BATCH calls against an LRU hash map with 16,777,216 buckets held a CPU inside the batch op for 77+ seconds and triggered the soft lockup watchdog. Commit 75134f16e7dd ("bpf: Add schedule points in batch ops") fixed this same problem in the generic batch ops, but not in this htab-native path, which every htab-based hash map variant uses for its lookup[_and_delete] batch ops. Complete that fix here. Leave the critical section after 64 consecutive empty buckets, call cond_resched_tasks_rcu_qs(), and resume at the saved bucket cursor. No locks are held at that point, and resuming from the cursor is already the function's behavior for non-empty buckets. Add the same call to the per-bucket loop after copy_to_user(), where every lock has been dropped. cond_resched_rcu() is not enough here: sleeping with bpf_prog_active elevated makes tracing programs on that CPU silently skip their invocations. Plain cond_resched() is not enough either. It is a no-op under PREEMPT and PREEMPT_LAZY, the only models arm64 and x86 have offered since commit 7dadeaa6e851 ("sched: Further restrict the preemption modes"). It is also never a Tasks RCU quiescent state, in any model: the reschedule counts as a preemption. The walking task stays a holdout and stalls every synchronize_rcu_tasks() caller, ftrace and BPF trampoline teardown included, until the syscall returns [1]. cond_resched_tasks_rcu_qs() is the usual tool for that [2]. It reports the quiescent state at each yield and still reschedules as cond_resched() does on PREEMPT_NONE and PREEMPT_VOLUNTARY kernels.

In the Linux kernel, the following vulnerability has been resolved: tcp: Skip cond_resched() in inet_csk_listen_stop() under BPF context bpf_sock_destroy() runs from the tcp iterator, under rcu_read_lock(). If the sock is a listener that still has children in its accept queue, tcp_abort() ends up in inet_csk_listen_stop() and the cond_resched() there trips the debug check: BUG: sleeping function called from invalid context at net/ipv4/inet_connection_sock.c:1523 in_atomic(): 0, irqs_disabled(): 0, non_block: 0, pid: 628, name: test_progs preempt_count: 0, expected: 0 RCU nest depth: 1, expected: 0 locks held by test_progs/628: 3, last CPU#3: #0: ffff8881158cee18 (&p->lock){+.+.}-{4:4}, at: bpf_seq_read+0x56/0x1210 #1: ffff8881106bb858 (sk_lock-AF_INET6){+.+.}-{0:0}, at: bpf_iter_tcp_seq_show+0x32b/0x4b0 #2: ffffffffb435af20 (rcu_read_lock){....}-{1:3}, at: bpf_iter_run_prog+0x46b/0xde0 CPU: 3 UID: 0 PID: 628 Comm: test_progs Tainted: G W 7.2.0+ #65 PREEMPT Tainted: [W]=WARN Call Trace: <TASK> dump_stack_lvl+0xc1/0xf0 dump_stack+0x10/0x20 __might_resched+0x3d2/0x610 inet_csk_listen_stop+0x7b/0xbf0 tcp_abort+0x23b/0x3b0 bpf_sock_destroy+0xfc/0x140 bpf_prog_448133d24601754f_iter_tcp6_server+0x81/0x8a bpf_iter_run_prog+0x538/0xde0 bpf_iter_tcp_seq_show+0x26b/0x4b0 bpf_seq_read+0x424/0x1210 vfs_read+0x197/0xe40 ksys_read+0x119/0x240 __x64_sys_read+0x72/0xc0 x64_sys_call+0x647/0x27e0 do_syscall_64+0xe5/0x610 entry_SYSCALL_64_after_hwframe+0x76/0x7e RIP: 0033:0x7fad39b28aca RSP: 002b:00007ffc381c61c0 EFLAGS: 00000246 ORIG_RAX: 0000000000000000 RAX: ffffffffffffffda RBX: 00007ffc381c6a88 RCX: 00007fad39b28aca RDX: 0000000000000032 RSI: 00007ffc381c6250 RDI: 0000000000000014 RBP: 00007ffc381c61e0 R08: 0000000000000000 R09: 0000000000000000 R10: 0000000000000000 R11: 0000000000000246 R12: 0000000000000003 R13: 0000000000000000 R14: 000055f077c1bbb0 R15: 00007fad3a0f3000 </TASK> The commit that added the kfunc already guards lock_sock() in tcp_abort() and udp_abort() with has_current_bpf_ctx(), but missed the listener path. Do the same for the cond_resched(). The loop runs inside the iterator's rcu_read_lock(), it must not reschedule or report a quiescent state there.

In the Linux kernel, the following vulnerability has been resolved: bpf: Fix divide-by-zero in btf_struct_walk() When an access goes past the struct and the last member is a flexible array, btf_struct_walk() folds the offset back into a single element with (off - moff) % t->size, but never checks that the element type has a size. BTF takes an empty struct, so this in program BTF /* event could be empty */ struct event { #ifdef HAVE_TIMESTAMP __u64 ts; #endif }; struct batch { int nr; struct event events[]; }; divides by zero at prog load time. Getting there needs a PTR_TO_BTF_ID that is not MEM_ALLOC, e.g. a plain read of a local kptr stashed in a map from a sleepable program. Oops: divide error: 0000 [#1] SMP KASAN PTI RIP: 0010:btf_struct_walk+0x53f/0x1570 Call Trace: <TASK> btf_struct_access+0x42a/0xcd0 check_ptr_to_btf_access+0x4dc/0x1160 check_mem_access+0x3a45/0x8740 check_load_mem+0x36a/0xd10 do_check_common+0x3ef0/0xb210 bpf_check+0x6d3b/0x8580 bpf_prog_load+0xf7c/0x2720 __sys_bpf+0xa83/0x3690 __x64_sys_bpf+0xc7/0x150 x64_sys_call+0x1f3f/0x27e0 do_syscall_64+0xe5/0x610 entry_SYSCALL_64_after_hwframe+0x76/0x7e </TASK> Reject a zero-sized element type. The fixed array path in the same function already bails out on the same thing: btf_struct_walk() ... /* skip empty array */ if (moff == mtrue_end) continue; msize /= total_nelems;

In the Linux kernel, the following vulnerability has been resolved: KVM: s390: Add missing srcu in kvm_s390_set_irq_state() Like kvm_s390_inject_vcpu(), kvm_s390_set_irq_state() also needs the kvm->srcu or the slots lock when performing the Store status operation. Fix by taking kvm->srcu in kvm_s390_set_irq_state().

In the Linux kernel, the following vulnerability has been resolved: KVM: x86/mmu: Bail from shadow walks if the root is invalid or a dummy When walking shadow page tables, immediately terminate the walk if the root is a "dummy" root, i.e. a root whose top-level page table is backed by the zero page, but otherwise doesn't exist. If memslot creation races with a stage-2 page fault (EPT violation or #NPF) from L2, then if the stars align, KVM will attempt to walk shadow page tables using the zero page and hit a NULL pointer deref. BUG: kernel NULL pointer dereference, address: 0000000000000021 #PF: supervisor read access in kernel mode #PF: error_code(0x0000) - not-present page PGD 0 P4D 0 Oops: Oops: 0000 [#1] SMP CPU: 30 UID: 1000 PID: 941 Comm: qemu Not tainted 7.2.0-rc2-1b731e5ded48-next-vm #1741 PREEMPTLAZY Hardware name: QEMU Standard PC (Q35 + ICH9, 2009), BIOS 0.0.0 02/06/2015 RIP: 0010:__kvm_mmu_invalidate_addr+0xea/0x210 [kvm] Call Trace: <TASK> kvm_mmu_invalidate_addr+0x92/0xe0 [kvm] __kvm_inject_emulated_page_fault+0x67/0x80 [kvm] ept_page_fault+0x160/0x850 [kvm] kvm_mmu_do_page_fault+0x102/0x1f0 [kvm] kvm_mmu_page_fault+0x8e/0x6b0 [kvm] vmx_handle_exit+0x163/0x640 [kvm_intel] kvm_arch_vcpu_ioctl_run+0x960/0x2120 [kvm] kvm_vcpu_ioctl+0x2c7/0x970 [kvm] __x64_sys_ioctl+0x90/0xd0 do_syscall_64+0x67/0x5f0 entry_SYSCALL_64_after_hwframe+0x4b/0x53 RIP: 0033:0x7f606d0b53bb Opportunistically harden the shadow walks against fully invalid roots, but WARN, as all callers are expected/required to pre-check for a valid root. Don't WARN in the dummy root case as the whole point of using a dummy root is to provide a root that's valid enough to enter the guest, i.e. it should Just Work for all flows except those that *need* to know about dummy roots. Alternatively, KVM could allocate a dedicated page and associated shadow page structure for the dummy root, which is very tempting as it such an approach should be more resilient against unexpected behavior. But that would be a much larger and thus riskier change than simply terminating walks of dummy roots.

In the Linux kernel, the following vulnerability has been resolved: iio: proximity: aw96103: validate firmware data length The firmware parser reads a length from the binary header and then uses it to walk six-byte register/value entries. A truncated firmware image can make the parser read beyond the copied firmware buffer. A value smaller than the four-byte count field also underflows, producing a very large length. Read the little-endian field with get_unaligned_le32(), as previously proposed by David Lechner. Validate the header size, reject an underflowed length, ensure the resulting payload is contained in the firmware image, and require it to contain whole register/value entries.

In the Linux kernel, the following vulnerability has been resolved: perf: Require kernel access for text poke events Perf events with exclude_kernel=1 can be opened without kernel perf access. However, exclude_kernel does not suppress text-poke sideband records. Every PERF_RECORD_TEXT_POKE is marked PERF_RECORD_MISC_KERNEL and contains a raw kernel instruction address. An unprivileged task can therefore open and mmap a task-local software event with text_poke=1. Both opening a count-only tracepoint event and configuring UDP GRO for ESP-in-UDP cause updates to inline static calls; the observer receives the relocated addresses of the modified instructions. For a known kernel image, any such address reveals the runtime kernel text base despite KASLR. Call perf_allow_kernel() whenever attr.text_poke is set, regardless of exclude_kernel. Events that neither monitor kernel execution nor request text-poke records retain their existing permissions.

In the Linux kernel, the following vulnerability has been resolved: iio: buffer: Ensure bounce buffer used for unaligned case is zeroed. iio_push_to_buffers_with_ts_unaligned() leaks uninitialized heap memory to userspace if the data passed in is not a multiple of 8 bytes and the timestamp is enabled. Specify __GFP_ZERO for the devm_krealloc() to ensure any extra space is cleared.

In the Linux kernel, the following vulnerability has been resolved: iio: adc: stm32-adc: fix possible division by zero in processed channel In case the conversion has failed or returned zero, processing *val can lead to a division by zero. Need to check for errors, or converted value is zero, before processing the data. In case the converted value is zero, e.g. the Vrefint channel, this should be considered as invalid in all cases.

In the Linux kernel, the following vulnerability has been resolved: iio: adc: ade9000: handle Phase C dip events in IRQ1 handler ADE9000_ST1_CROSSING_DEPTH is the exclusive upper bound for scanning the STATUS1 event bits, so a value of 25 left out the highest event bit, ADE9000_ST1_DIPC_BIT (bit 25). Phase C dip events were never reported. Bump the value to 26 so bit 25 is scanned.

In the Linux kernel, the following vulnerability has been resolved: iio: adc: pac1934: check ACPI label duplication A label duplication failure currently reaches the terminating-byte write. Release the ACPI result and abort before accessing the missing label; earlier labels are managed by devres. Detected by static analysis and reviewed with AI-assisted source auditing.

In the Linux kernel, the following vulnerability has been resolved: iio: accel: kxcjk-1013: reject duplicate event disable The IIO core does not filter duplicate writes to the event enable attribute. kxcjk1013_write_event_config() already ignores repeated enable requests, but a repeated disable request still calls kxcjk1013_set_power_state(data, false), dropping a runtime PM reference that was not acquired for this request. This can underflow the runtime PM usage count and trigger a "Runtime PM usage count underflow" warning. Return early when the requested state already matches ev_enable_state.

In the Linux kernel, the following vulnerability has been resolved: smb: client: split cifsFileInfo bitfields to avoid shared-byte RMW races The invalidHandle, swapfile, oplock_break_cancelled, offload, and status_file_deleted fields are stored in the same bitfield byte in struct cifsFileInfo, but are updated in different code paths that may run simultaneously, and are protected by different locks. Since bitfield assignments generate byte-level read-modify-write operations, a modification to one flag can overwrite a concurrent modification to another flag. To avoid these races, convert these flags from a bitfield to separate bool fields.

In the Linux kernel, the following vulnerability has been resolved: KVM: SVM: Update control fields on #VMEXIT if and only if VMRUN succeeded Leave control.erap_ctl and control.clean as-is in the VMCS if VMRUN fails, because as per AMD: there's no explicit architectural guarantee about the behavior in the presence of VMRUN failures. So the best thing to do would be to assume that if VMRUN fails, the actions requested in the control fields may not have been performed.

In the Linux kernel, the following vulnerability has been resolved: KVM: SVM: Preserve TLB control (i.e. pending TLB flush) on failed VMRUN Don't reset the VMCB's TLB control back to "do nothing" on a failed VMRUN, as empirical testing shows that the CPU performs the requested TLB flush if and only if VMRUN is successful, i.e. clearing TLB control on a failed VMRUN effectively drops a TLB flush. Explicitly track the need to flush all ASIDs on a per-CPU basis, as the ASID reuse condition is tied to the pCPU, not to the vCPU. As a bonus, this also obviates the need to avoid clobbering FLUSH_ALL_ASID with TLB_CONTROL_FLUSH_ASID, e.g. in svm_flush_tlb_asid(). Deliberately don't bother saving/restoring the "old" tlb_ctl on failure, in quotes because it's not exactly the old tlb_ctl, it's the tlb_ctl from after pre_svm_run(), but before updating tlb_ctl for flush_all_asids. If VMRUN fails and TLB_CONTROL_FLUSH_ALL_ASID is forced, then the next successful run of the VMCB *may* unnecessarily flush all ASIDs, which strictly speaking could result in noisy neighbor issues. However, the fact that new_asid() is already guest-triggerable, because of KVM's flawed behavior of clearing the ASID on emulated INIT, means that a guest can already trigger a flush of all ASIDs at roughly the same rate. And once KVM stops clobbering the ASID on emulated INIT, *or* assigns a static ASID to each vCPU, this flaw goes away.

In the Linux kernel, the following vulnerability has been resolved: mtd: block2mtd: Fix divide error when erase_size is zero The erase size is parsed from the "block2mtd" module parameter and can be set to zero. add_device() then evaluates if (size % erase_size) with a zero divisor, which triggers a divide error: divide error: 0000 [#1] PREEMPT SMP PTI RIP: 0010:add_device drivers/mtd/devices/block2mtd.c:296 [inline] RIP: 0010:block2mtd_setup2+0x592/0xda0 drivers/mtd/devices/block2mtd.c:459 Call Trace: block2mtd_setup+0x27/0xe0 drivers/mtd/devices/block2mtd.c:476 param_attr_store+0x214/0x310 kernel/params.c:589 module_attr_store+0x65/0x90 kernel/params.c:904 kernfs_fop_write_iter+0x3a4/0x540 fs/kernfs/file.c:345 ... Reject a zero erase size before performing the modulo operation so the existing "erasesize must be a divisor of device size" error path reports the invalid argument and frees the device.

In the Linux kernel, the following vulnerability has been resolved: Bluetooth: hci_conn: Lock parent access during enhanced SCO setup Bluetooth: hci_conn: Lock parent access during enhanced SCO setup hci_enhanced_setup_sync() runs on the request workqueue without the hci_dev_lock held by its caller when setup was queued. Its CVSD capability check and find_next_esco_param() dereference conn->parent while the receive workqueue can unlink and release that parent. The following interleaving can cause a use-after-free: hci_enhanced_setup_sync() hci_disconn_complete_evt() load conn->parent hci_dev_lock() hci_conn_del(ACL parent) unlink SCO child drop link's parent reference clear child->parent release ACL parent bt_link_release() kfree(parent) hci_dev_unlock() read parent->features[0][3] The reference held for the queued SCO child does not keep its ACL parent alive after unlinking. Commit 42de40abe25d ("Bluetooth: hci_conn: fix the SCO setup context lifetime") protects the child stored in the queued context, but leaves these parent accesses unprotected. With a 40 ms diagnostic delay after loading conn->parent, an instrumented kernel based on fd179f8a05be, which already contains 42de40abe25d, BUG: KASAN: slab-use-after-free in hci_enhanced_setup_sync+0xda5/0xdf0 Read of size 1 at addr ffff888102204047 by task kworker/u17:0/93 Call Trace: hci_enhanced_setup_sync+0xda5/0xdf0 hci_cmd_sync_work+0x13c/0x290 process_one_work+0x6b4/0x10e0 Allocated by task 92: __hci_conn_add+0x304/0x1df0 hci_connect_acl+0x349/0x3e0 hci_connect_sco+0x3b/0x9a0 sco_sock_connect+0x475/0xca0 Freed by task 94: kfree+0x121/0x3c0 bt_link_release+0x79/0xa0 device_release+0xc8/0x240 kobject_put+0x14d/0x280 hci_conn_del+0x524/0xe30 hci_disconn_complete_evt+0x403/0x8c0 hci_event_packet+0x71b/0xb20 hci_rx_work+0x293/0x730 The accessed address is 71 bytes into the freed ACL parent, at its features[0][3] byte; the queued SCO child is a different object. The diagnostic preserves the loaded parent across the delay, matching the unmodified compiled capability check, and does not change the parent references or teardown path. Hold hci_dev_lock() across the codec switch, including every call to find_next_esco_param(), and release it on all selection errors. Keep configure_datapath_sync() outside the critical section because it waits for HCI events. Preserve parameter selection and existing return values.

In the Linux kernel, the following vulnerability has been resolved: EDAC/altera: Drop __init from ECC setup paths for re-probe safety Sashiko reports: "Does suppressing sysfs unbinding fully prevent the execution of freed __init memory? If altr_sysmgr_regmap_lookup_by_phandle() returns -EPROBE_DEFER, the probe is deferred until after __init memory is freed." The a10 EDAC .setup callbacks (sdmmc, ethernet, nand, dma, usb, qspi) and their helpers (altr_init_a10_ecc_device_type, altr_init_a10_ecc_block) were marked __init. These run from the probe path, which may execute after init memory is freed -- e.g. a probe deferred via -EPROBE_DEFER that only succeeds once a late/module dependency appears, or a manual unbind/rebind. Calling __init code then dereferences freed memory. Remove __init so these functions remain valid at runtime.

In the Linux kernel, the following vulnerability has been resolved: ALSA: ua101: reject mismatched capture/playback packet sizes detect_usb_format() cross-checks bSubframeSize, bBitResolution and tSamFreq between the capture and playback interfaces, but never relates the two endpoints' wMaxPacketSize and bNrChannels. Each playback URB gets a buffer of ua->playback.max_packet_bytes, while the number of bytes written into it is derived from the capture stream: capture_urb_complete() computes frames from the received capture packet and capture.frame_bytes, and start_usb_playback() and playback_work() multiply that by playback.frame_bytes. A device declaring a large capture wMaxPacketSize with few capture channels and a small playback wMaxPacketSize with many playback channels therefore memset()s and memcpy()s past the end of the playback buffer, in open() of the PCM node the driver registers during probe. usb_submit_urb() rejects the over-long iso_frame_desc[0].length with -EMSGSIZE, but only after the write. Reject such descriptors at probe time. Genuine UA-101/UA-1000 hardware declares proportional packet sizes and is unaffected. BUG: KASAN: slab-out-of-bounds in start_usb_playback (sound/usb/misc/ua101.c:586) Write of size 2048 at addr ffff8881098f3c00 by task exploit/5021 Call Trace: __asan_memset (mm/kasan/shadow.c:84) start_usb_playback (sound/usb/misc/ua101.c:586) playback_pcm_open (sound/usb/misc/ua101.c:679) snd_pcm_open_substream (sound/core/pcm_native.c:2829) snd_pcm_open (sound/core/pcm_native.c:2865 sound/core/pcm_native.c:2932) snd_pcm_playback_open (sound/core/pcm_native.c:2891) snd_open (sound/core/sound.c:166) chrdev_open (fs/char_dev.c:411) do_dentry_open (fs/open.c:996) vfs_open (fs/open.c:1101) path_openat (fs/namei.c:4837 fs/namei.c:5000) do_file_open (fs/namei.c:5029) do_sys_openat2 (fs/open.c:1417) __x64_sys_openat (fs/open.c:1423 fs/open.c:1439 fs/open.c:1434) do_syscall_64 (arch/x86/entry/syscall_64.c:61 arch/x86/entry/syscall_64.c:84) entry_SYSCALL_64_after_hwframe (arch/x86/entry/entry_64.S:121) The buggy address belongs to the object at ffff8881098f3c00 which belongs to the cache kmalloc-192 of size 192 The buggy address is located 0 bytes inside of allocated 168-byte region [ffff8881098f3c00, ffff8881098f3ca8)

In the Linux kernel, the following vulnerability has been resolved: arm64/fpsimd: signal: Forbid non-streaming SVE payload on SME-only systems When system_supports_sme() is true but system_supports_sve() is false, restoring a specifically crafted SVE signal context can result in the task erroneously having non-streaming SVE state. Subsequent attempts to manipulate the task's FPSIMD/SVE/SME state can result in a variety of problems, including fatal EL1 UNDEFs. In such configurations, the kernel always creates an SVE signal context when delivering a signal, and this can only be in one of two states: (1) SVE_SIG_FLAG_SM is set, and an SVE payload is present containing streaming mode SVE state. The recorded VL is the task's live streaming VL. (2) SVE_SIG_FLAG_SM is clear, and an SVE payload is not present. The FPSIMD context contains the non-streaming mode FPSIMD state. The recorded VL is 0. Currently restore_sve_fpsimd_context() correctly rejects cases where SVE_SIG_FLAG_SM is set and an SVE payload is not present, but fails to reject cases where SVE_SIG_FLAG_SM is clear and an SVE payload is present. Consequently, restore_sve_fpsimd_context() can place the task in a state where it has non-streaming SVE state even when this is not supported by HW. For example, this can cause a later EL1 UNDEF when the kernel attempts to restore the task's ZCR_EL1 value: | # ./sme-sigcontext-to-sve | Internal error: Oops - Undefined instruction: 0000000002000000 [#1] SMP | Modules linked in: | CPU: 0 UID: 0 PID: 131 Comm: sme-sigcontext- Not tainted 7.3.0-rc1 #1 PREEMPT | Hardware name: linux,dummy-virt (DT) | pstate: 61402009 (nZCv daif +PAN -UAO -TCO +DIT -SSBS BTYPE=--) | pc : fpsimd_restore_current_state+0x258/0x458 | lr : exit_to_user_mode_loop+0xb8/0x188 | sp : ffff80008056be40 | x29: ffff80008056be40 x28: fff00000c1670000 x27: 0000000000000000 | x26: 0000000000000000 x25: 0000000000000000 x24: 0000000000000000 | x23: ffff80008056bec0 x22: 0000000000000008 x21: 0000000000000040 | x20: 0000000000000081 x19: 0000000008800010 x18: 0000000000000000 | x17: 0000fffffd412570 x16: 0000000000001000 x15: 0000fffffd4123b0 | x14: 0000fffffd412780 x13: 0000fffffd412be8 x12: 0000000047435300 | x11: 0000fffffd412570 x10: 0000000000000000 x9 : 0000000045585401 | x8 : fff00000c18a6c44 x7 : 0000000000000000 x6 : 0000000000000002 | x5 : 0000000000000002 x4 : ffff800080568000 x3 : 0000000000000001 | x2 : 0000000008800010 x1 : fff00000c1670000 x0 : 0000000008800000 | Call trace: | fpsimd_restore_current_state+0x258/0x458 (P) | exit_to_user_mode_loop+0xb8/0x188 | el0_svc+0x1cc/0x1d0 | el0t_64_sync_handler+0xa0/0xe4 | el0t_64_sync+0x198/0x19c | Code: d5384101 f9400020 53175c03 36b80de0 (d5381202) | ---[ end trace 0000000000000000 ]--- | Kernel panic - not syncing: Oops - Undefined instruction: Fatal exception in interrupt | Kernel Offset: 0x291fb1000000 from 0xffff800080000000 | PHYS_OFFSET: 0x40000000 | CPU features: 0x0,00000000,0052802f,ffb88f43,3afcf73f | Memory Limit: none Rework restore_sve_fpsimd_context() to reject cases where SVE_SIG_FLAG_SM is clear and an SVE payload is not present. As parse_user_sigframe() rejects SVE signal frames when neither SVE nor SME are supported, it isn't necessary for restore_sve_fpsimd_context() to handle the case where neither are supported.

In the Linux kernel, the following vulnerability has been resolved: tty: vcc: zero-initialize control packet in vcc_send_ctl() The stack-allocated struct vio_vcc pkt was partially initialized, leaving the tag.stype_env field uninitialized before being sent via ldc_write(), potentially leaking kernel stack data to the LDC peer.

In the Linux kernel, the following vulnerability has been resolved: usb: gadget: aspeed-vhub: cancel wake work on device removal wake_work is armed from the gadget .wakeup callback to resume suspended downstream ports, and it is never cancelled in ast_vhub_remove(), so a work item queued before or during removal can run after devm has freed vhub and its ports, a use-after-free. Cancelling the work alone is not sufficient: usb_gadget_wakeup() takes no lock and the unbind path never clears wakeup_en, so a remote-wakeup request in flight on another CPU can re-arm the work after the cancel and before vhub is freed. ast_vhub_del_dev() already clears d->registered under vhub->lock before unregistering the gadget. Refuse the wakeup in ast_vhub_udc_wakeup() once d->registered is clear: the check and the schedule_work() are then atomic against the unbind, so a wakeup that passed before the flag was cleared is drained by the cancel_work_sync() after the del_dev loop, and one that arrives later returns without arming. Also move INIT_WORK() to the top of ast_vhub_probe(), since the probe error path reuses ast_vhub_remove() and would otherwise cancel a never-initialized work item. This issue was found by an in-house static analysis tool.

In the Linux kernel, the following vulnerability has been resolved: nitro_enclaves: fix use-after-free on SLOT_ALLOC failure When ne_create_vm_ioctl() fails the SLOT_ALLOC request after anon_inode_getfile() has succeeded, the error path calls fput(enclave_file) and then frees ne_enclave. In normal userspace context, fput() defers the final __fput() via task_work. ne_enclave_release() therefore runs after ne_enclave has already been freed and dereferences ne_enclave->slot_uid, causing a use-after-free: KASAN: slab-use-after-free in ne_enclave_release. The enclave has no slot allocated and is not yet linked into the enclaves list on this error path, so ne_enclave_release() is expected to return early when slot_uid is zero. However, reading slot_uid already accesses the freed object. Clear enclave_file->private_data before fput() on the error path. ne_enclave_release() then returns immediately when private_data is NULL, leaving the ioctl error path as the sole owner of ne_enclave. This is safe because the file has not been fd_install()'d yet. Tested on an AWS EC2 m5.2xlarge with CONFIG_KASAN=y. Without the patch, the reproducer triggers a KASAN slab-use-after-free on every SLOT_ALLOC failure. With the patch, no KASAN report is produced and the SLOT_ALLOC error is still returned. Normal enclave creation and teardown are unaffected.

In the Linux kernel, the following vulnerability has been resolved: drm/mediatek: Fix pdev reference leak in mtk_drm_bind() When the device is not the mmsys master, mtk_drm_bind() returns early after taking a reference on the disp-mutex device via of_find_device_by_node(), without ever dropping it: mtk_drm_unbind() only puts mutex_dev for the master. Drop the reference before returning from the non-master path.

In the Linux kernel, the following vulnerability has been resolved: drm/mediatek: Fix ovl adaptor platform device leak mtk_drm_probe() creates an OVL adaptor platform device with platform_device_register_data() when the display pipeline requires the OVL adaptor. If a later initialization step fails, the probe error path releases the DRM resources without unregistering the already registered OVL adaptor device. The normal remove path likewise leaves the device registered after the DRM driver is unbound. Keep track of whether the OVL adaptor was successfully registered and unregister it on probe failure. Also recover the platform device from the stored DDP component device and unregister it during normal removal. The issue was identified by a static analysis tool I developed and confirmed by manual review.

In the Linux kernel, the following vulnerability has been resolved: net: cap skb->queue_mapping when the tx queue is picked skbedit can set skb->queue_mapping and raise the per-CPU skip_txqueue flag so __dev_queue_xmit() honours the mapping. __dev_queue_xmit() cleared the flag before sch_handle_egress() and only read it afterwards, so the flag was not confined to the xmit that set it: a nested xmit (mirred redirect or mirror, or a drop after skbedit) could set the flag and the outer xmit would consume it for an skb that never went through skbedit. A forwarded packet still carries the ingress NIC's rx_queue + 1 in skb->queue_mapping, so the outer device then indexes its tx queue state with that stale value. Taprio's child array q->qdiscs[] is sized to the device's queue count, so taprio_enqueue() indexes past its allocation and dereferences the result as a struct Qdisc *. We (ab)use the skb->nf_skip_egress which means "skip netfilter egress for this packet" to tag to "am I in tc egress?". Despite the overload I dont see it as a conflict since the marker is set only around the single sch_handle_egress() call and ingress path is guarded by tc_at_ingress. I will send a followup(net-next) patch once this hits net-next to rename the skb->nf_skip_egress bit/flag to skb->skip_egress Arm the flag only from the egress classifier that can use it: raise skip_txqueue from tcf_skbedit_act() only when it runs inside sch_handle_egress(), thanks to skb->nf_skip_egress. An egress qdisc classifier runs in q->enqueue(), after the tx queue has been picked, so a mapping it sets cannot affect the current packet; arming the flag there only pollutes it for a later xmit. Then own the flag for the xmit frame the egress hook runs in: save the incoming value and clear it just before sch_handle_egress(), and restore it after the hook - on the consumed (drop) path, or, in the same call that reads it, on the surviving path. The save and the restores stay inside the egress_needed_key static branch, so a packet pays for them only when egress hooks are active (2f1e85b1aee4). Store the value netdev_cap_txqueue() selected back into skb->queue_mapping in netdev_tx_queue_mapping(), as netdev_core_pick_tx() already does, so the skip_txqueue path never hands a later reader on the xmit path a mapping the device cannot serve. A store made still later in the same frame, by a tc BPF program attached to a transmit qdisc, is outside this path and is not re-capped; a separate followup will resolve that path. netdev_xmit_skip_txqueue() returns the previous flag value so the save-and-clear is one call, and a no-op stub is provided when CONFIG_NET_EGRESS is disabled. skb->nf_skip_egress is compiled under CONFIG_NET_EGRESS rather than CONFIG_NETFILTER_SKIP_EGRESS, so skb_at_tc_egress() is valid whenever the egress path is built. A local user in a network namespace can redirect a packet from a device with more TX queues to one with fewer after setting a mapping valid only on the larger device. That reaches these reads and, under KASAN, faults with "slab-out-of-bounds in taprio_enqueue". Conditions to recreate the bug: the report's own trigger is a local user with CAP_NET_ADMIN in a network namespace, so no eBPF program is needed. With CONFIG_NET_SCH_TAPRIO=y, CONFIG_NET_ACT_SKBEDIT=y, CONFIG_NET_ACT_MIRRED=y, CONFIG_NET_CLS_MATCHALL=y, CONFIG_NET_SCH_PRIO=y and KASAN enabled, create qa (3 queues), qb (2 queues) and qc (1 queue) as dummy devices; put a taprio root on qb (num_tc 1, queues 2@0) and clsact on all three; then add an egress matchall filter on every device. On qa: "action skbedit queue_mapping 2 pipe action mirred egress redirect dev qb". On qb: "action mirred egress mirror dev qc". On qc: "action skbedit queue_mapping 0 pipe". Send one packet out qa. qc's skbedit sets the flag while qb's outer xmit is in flight; without the fix qb consumes it and reads its two-entry taprio child array with the forwarded packet's stale mapping. A qc whose skb ---truncated---

In the Linux kernel, the following vulnerability has been resolved: wifi: mac80211: fix slab-out-of-bounds read in ieee80211_monitor_select_queue() ieee80211_monitor_select_queue() validates the skb length using skb->len before dereferencing pointers into skb->data to access the 802.11 header. However, skb->len includes data in both the linear head buffer and non-linear fragments/pages. When AF_PACKET sends a packet large enough to become non-linear, the 802.11 header at skb->data + len_rthdr may extend past the linear head buffer even though skb->len appears sufficient. This leads to a slab-out-of-bounds read when accessing hdr->frame_control, as the kernel reads beyond the allocated skb head buffer: BUG: KASAN: slab-out-of-bounds in ieee80211_monitor_select_queue+0x1ed/0x220 syzbot has hit the same issues. Fix this by checking skb_headlen(skb) (which gives the length of the linear data region) instead of skb->len before dereferencing skb->data pointers. This ensures the required bytes are actually present in the linear portion of the skb. Apply the same fix to ieee80211_validate_radiotap_len() which has the same class of bug: it uses skb->len to validate accesses to skb->data, and to the corresponding checks in ieee80211_monitor_start_xmit(): for drivers advertising NETIF_F_SG the skb is not linearized before ndo_start_xmit, so the 802.11 header can be read past the linear buffer there as well.

In the Linux kernel, the following vulnerability has been resolved: bpf: Hold map BTF for the memory allocator destructor record bpf_ma_set_dtor() duplicates map->record so that the bpf_mem_alloc destructor can release the special fields of hash and rhash map elements once the allocator frees them for good. btf_record_dup() only acquires references on kernel and module BTF. Fields whose types live in the map BTF keep pointing into it: a kptr to a local type refers to map->btf, and a list_head or rb_root field carries a value_rec owned by the struct meta table of map->btf. That borrowed state can outlive the map. When RCU callbacks are still in flight, bpf_mem_alloc_destroy() copies the allocator and defers the final drain, together with the destructor context, to a workqueue. bpf_map_free() then frees map->record and drops the map's BTF reference as soon as map_free() returns. Once the program that loaded the BTF is gone as well, the BTF is freed while the worker still uses the duplicate. The worker dereferences it through btf_is_kernel() and btf_find_struct_meta() in bpf_obj_free_fields() when destroying elements that were freed under RCU, and through btf_is_kernel() in btf_record_free() when releasing the context itself: BUG: KASAN: slab-use-after-free in btf_is_kernel+0x19/0x30 Read of size 1 at addr ffff88811505dce8 by task kworker/u16:2/50 Workqueue: events_unbound free_mem_alloc_deferred Call Trace: btf_is_kernel+0x19/0x30 btf_record_free+0xc7/0xf0 htab_dtor_ctx_free+0x1a/0x30 free_mem_alloc_no_barrier+0x39/0x230 free_mem_alloc_deferred+0x2a/0x40 process_scheduled_works+0x5ce/0x9d0 worker_thread+0x42a/0x5e0 kthread+0x1f5/0x230 ret_from_fork+0x1dd/0x390 ret_from_fork_asm+0x1a/0x30 Allocated by task 358: __kmalloc_cache_noprof+0x287/0x510 btf_new_fd+0xf7/0x3d0 __sys_bpf+0x487/0x840 Freed by task 0: kfree+0x186/0x560 rcu_core+0x6f6/0xd40 Freeing a map that still holds more elements than the allocator's high watermark is enough to get there, because the bulk free during teardown queues the RCU callback that makes bpf_mem_alloc_destroy() defer. Non-preallocated hash maps and resizable hash maps both register the destructor and are affected alike. Keep one reference on the map BTF in the destructor context. Every non-kernel pointer in the duplicated record is either map->btf or memory owned by it, so a single reference covers kptrs, list heads, and rb roots, and mirrors what bpf_map_meta_alloc() does for the duplicated record of an inner map. Release it only after the duplicated record has been freed: the worker is preemptible and not in an RCU read-side critical section, so dropping the last reference first would let the BTF be freed while btf_record_free() still reads its fields.

In the Linux kernel, the following vulnerability has been resolved: net/sched: cls_api: reclaim an empty proto on the error path Two racing tc filter add requests on the same chain/prio of an unlocked classifier both run change() on the shared proto and both can fail: the winner's tcf_chain_tp_delete_empty() attempt gives up because the loser's handle is still in the idr, and the loser's error path drops only its own reference without a second reclamation attempt. The empty proto stays linked in the chain, holding the chain reference, a block reference and the classifier module reference until the chain or block is torn down. Reclaim the proto on the error path of any failed request that holds a proto reference. The reclamation is emptiness-gated: tcf_chain_tp_delete_empty() unlinks the proto only when delete_empty() admits it is empty, so a live shared proto is never unlinked. A proto the request created is reclaimed unconditionally - it is the only owner, so marking it for deletion is safe even without a delete_empty callback. Classifiers without one (the check marks the proto unconditionally) are rtnl-serialized, so the raced window this guard closes cannot arise for them. This is a follow-up to commit d4e359b3608a ("net/sched: cls_api: fix teardown of an adopted proto on insert-race loss"), which stopped the loser of the insert race from unlinking the winner's live proto but left the empty-proto residual in place. Conditions to recreate: - CONFIG_NET_CLS_FLOWER=y; veth pair - tc qdisc add dev veth0 ingress - two concurrent `tc filter add dev veth0 ingress protocol ip pref 1 flower skip_sw ... action drop` (both fail in fl_hw_replace_filter after publishing their handle in the idr); repeat in a loop - an empty flower tp stays linked after both requests fail; visible as a bare `filter protocol ip pref 1 flower chain 0` header in `tc filter show` with no filter entries - CAP_NET_ADMIN (namespace-local via unshare -Urn suffices)

In the Linux kernel, the following vulnerability has been resolved: mm/slab: do not wake up kswapd in __kfree_rcu_sheaf() Since kfree_rcu() can be called under pi_lock (a raw spinlock in the scheduler), kfree_rcu() itself should never allocate memory with __GFP_KSWAPD_RECLAIM as waking up kswapd ends up acquiring pi_lock, which leads to a deadlock. Reproducing the issue even intentionally was not straightforward. The set_cpus_allowed_force() path that is called under pi_lock is exercised very rarely, and competing tasks that allocate memory will most likely wake kswapd up. Therefore the existence of the deadlock was verified with a modified kernel that has a lockdep map for waking up kswapd: ====================================================== WARNING: possible circular locking dependency detected 7.2.0-rc1-slab-for-next+ #10 Not tainted ------------------------------------------------------ git/6660 is trying to acquire lock: ffff8e524533c550 (&p->pi_lock){-.-.}-{2:2}, at: _raw_spin_lock_irqsave+0x12/0x20 but task is already holding lock: ffff8e57bffff1a0 (&pgdat->kswapd_wait){....}-{3:3}, at: _raw_spin_lock_irqsave+0x12/0x20 which lock already depends on the new lock. the existing dependency chain (in reverse order) is: -> #2 (&pgdat->kswapd_wait){....}-{3:3}: __lock_acquire+0x5a4/0xc50 lock_acquire.part.0+0xb7/0x240 lock_acquire+0x70/0x170 __raw_spin_lock_irqsave+0x44/0x80 _raw_spin_lock_irqsave+0x12/0x20 __wake_up_common_lock+0x31/0xa0 __wake_up+0x20/0x40 kswapd_wakeup_wake+0x7e/0x110 wakeup_kswapd+0x295/0x320 wake_all_kswapds+0xac/0x1a0 __alloc_pages_slowpath.constprop.0+0x282/0xf80 __alloc_frozen_pages_noprof+0x32a/0x360 alloc_slab_page+0x2e/0x160 allocate_slab+0x82/0x420 new_slab+0x52/0xb0 refill_objects+0x13d/0x190 refill_sheaf+0x5d/0xd0 __pcs_replace_empty_main+0x230/0xb10 [...] -> #1 (kswapd_wakeup){-.-.}-{0:0}: __lock_acquire+0x5a4/0xc50 lock_sync.part.0+0x75/0x100 lock_sync+0x36/0x70 might_wakeup_kswapd+0x61/0xa0 __kfree_rcu_sheaf+0x33/0xd90 kvfree_call_rcu+0x1d4/0x3b0 set_cpus_allowed_force+0x163/0x220 cpuset_cpus_allowed_fallback+0x18b/0x240 select_fallback_rq+0x1e6/0x250 [...] -> #0 (&p->pi_lock){-.-.}-{2:2}: check_prev_add+0xe6/0xe00 validate_chain+0x51e/0x6e0 __lock_acquire+0x5a4/0xc50 lock_acquire.part.0+0xb7/0x240 lock_acquire+0x70/0x170 __raw_spin_lock_irqsave+0x44/0x80 _raw_spin_lock_irqsave+0x12/0x20 try_to_wake_up+0x77/0xa90 default_wake_function+0x27/0x60 autoremove_wake_function+0x23/0xb0 __wake_up_common+0xb8/0x170 __wake_up_common_lock+0x51/0xa0 __wake_up+0x20/0x40 kswapd_wakeup_wake+0x7e/0x110 wakeup_kswapd+0x295/0x320 wake_all_kswapds+0xac/0x1a0 __alloc_pages_slowpath.constprop.0+0x282/0xf80 __alloc_frozen_pages_noprof+0x32a/0x360 alloc_slab_page+0x2e/0x160 allocate_slab+0x82/0x420 new_slab+0x52/0xb0 refill_objects+0x13d/0x190 refill_sheaf+0x5d/0xd0 __pcs_replace_empty_main+0x230/0xb10 [...] other info that might help us debug this: Chain exists of: &p->pi_lock --> kswapd_wakeup --> &pgdat->kswapd_wait Possible unsafe locking scenario: CPU0 CPU1 ---- ---- lock(&pgdat->kswapd_wait); lock(kswapd_wakeup); lock(&pgdat->kswapd_wait); lock(&p->pi_lock); *** DEADLOCK *** 3 locks held by git/6660: #0: ffff8e519f0e9e20 (&type->i_mutex_dir_key#6){++++}-{4:4}, at: lookup_slow+0x2d/0x60 #1: ffffffff8a31c460 (kswapd_wakeup){-.-.}- ---truncated---