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Linux

In the Linux kernel, the following vulnerability has been resolved: dm-integrity: fix NULL pointer dereference when the 'R' flag is used If the dm-integrity device has the SB_FLAG_DIRTY_BITMAP flag set and the user activates the device in the 'R' mode, a crash in dm_integrity_resume happens because the function attempts to read the journal containing the bitmap. This patch makes dm-integrity skip any writes to the device in dm_integrity_resume if the device is activated in the 'R' mode.

Linux

In the Linux kernel, the following vulnerability has been resolved: dm-integrity: validate the superblock on resume dm_integrity_resume() re-reads the superblock from the device so that it picks up the flags and the recalculate position. It performs no validation on the result, while the constructor validates the superblock it reads and sizes all the in-memory structures according to it. The user may modify the superblock on the underlying device while the dm-integrity device is suspended, so that the two no longer agree. In particular, access_journal_data() shifts the journal entry index by ic->sb->log2_sectors_per_block, while ic->journal_section_sectors and the journal page list were computed with the value that was present at constructor time. Increasing log2_sectors_per_block makes the index run past the end of the journal, and dm-integrity then writes 512 bytes through lowmem_page_address(NULL). Snapshot the validated superblock in the constructor and refuse to resume if any of the fields that describe the on-disk geometry changed.

Linux

In the Linux kernel, the following vulnerability has been resolved: sched_ext: Fix CPU hotplug hang when a dying CPU's tasks sit in the BPF scheduler A CPU going down has to empty its own rq. Its hotplug thread waits in sched_cpu_wait_empty() until nothing else is left, and the only thing that wakes it is balance_push(), which runs from __schedule() on the dying CPU and pushes off the migratable tasks the CPU picks. sched_ext breaks this. A task sitting on a user DSQ or held by the BPF scheduler is still counted on its rq, but an inactive CPU no longer calls ops.dispatch() and can't pull the task back. The dying CPU goes idle with the task still counted, and one of two things happens: - Another CPU consumes the task. The rq empties without a __schedule() on the dying CPU, so the hotplug thread is never woken and cpu_down() hangs holding cpu_hotplug_lock. - The task is affine only to the dying CPU. Nothing can move it, and the offline stalls until the watchdog ejects the BPF scheduler. When the rq goes offline, re-enqueue every task on it that isn't already on the local DSQ. An enqueue on an offline rq lands on the local DSQ, so the dying CPU itself picks the tasks and balance_push() pushes them off, the same as for the other sched classes. From then on no sched_ext path on another CPU can pull a task off the rq. ops.dispatch() currently stops as soon as the CPU goes inactive, and CPU hotplug then waits for an RCU grace period before taking the rq offline. A BPF-held task affine only to the dying CPU and preempted inside an RCU read-side critical section would block that grace period, and the rq would never go offline. Test SCX_RQ_ONLINE directly so that ops.dispatch() keeps running until the rq goes offline. Only the dying CPU's own hotplug thread clears the flag during teardown, and task_can_run_on_remote_rq() still keeps other rqs' tasks off the CPU.

Linux

In the Linux kernel, the following vulnerability has been resolved: readdir: take no inode lock on an immutable directory iterate_dir() takes the directory's i_rwsem shared and holds it across ->iterate_shared(). For a directory that never has an entry and is never removed the lock keeps nothing still, it only orders every reader and every writer of that inode behind each other. For the directory of a nullfs instance that matters. The instance of the initial mount namespace is the root of every empty mount namespace and the private instance is the root of every kernel thread, so one inode is shared across users who have nothing else in common. And a reader can hold the lock for as long as it likes: back the getdents() buffer with a mapping of a file on a FUSE mount of your own, let the copy of "." and ".." fault and let the server wait. Queue an exclusive taker behind it, a mkdir() in that directory goes through start_dirop() before the read-only mount is reported, and from then on every lookup that misses the dcache in that directory, every create and every mount on it waits until the server answers. One user of an empty mount namespace stalls all the others. Add FOP_IMMUTABLE for the file operations of a directory that never changes and is never removed and let iterate_dir() skip the lock for it. The flag never changes for a file, ->f_pos is protected by f_pos_lock since directories are FMODE_ATOMIC_POS, IS_DEADDIR can't be set on such a directory and neither touch_atime() nor fsnotify take i_rwsem. Set it on the nullfs directory. The placeholder directories of libfs never have an entry either but their owners remove them, so they keep the lock.

Linux

In the Linux kernel, the following vulnerability has been resolved: ksmbd: fix heap out-of-bounds write in krb5_authenticate() In krb5_authenticate(), out_len is calculated to determine available headroom in response_buf for the incoming Kerberos AP-REP token. However, when SMB2_SESSION_SETUP is processed as a non-first element of a compounded SMB2 request, the calculation omits work->next_smb2_rsp_hdr_off. This causes out_len to overstate remaining buffer headroom by the cumulative size of prior responses in the compound chain. Consequently, the length check in ksmbd_krb5_authenticate() (*out_len <= resp->spnego_blob_len) passes erroneously, allowing memcpy() to write the AP-REP blob past the end of response_buf into adjacent kernel heap memory. Fix this by subtracting work->next_smb2_rsp_hdr_off when computing out_len, ensuring it accurately reflects physical remaining buffer space.

Linux

In the Linux kernel, the following vulnerability has been resolved: KVM: x86: Use active memslots for the per-vCPU MMIO cache KVM tags the per-vCPU MMIO cache with a memslot generation so that a memslot update invalidates cached MMIO information. Both the fill and validation paths use kvm_memslots(), which unconditionally selects address space 0, even when the vCPU is running in SMM and using address space 1. Commit 56f17dd3fbc4 ("kvm: x86: fix stale mmio cache bug") added the memslot generation to the cache key so that a memslot update could not leave a stale entry valid. When commit 699023e23965 ("KVM: x86: add SMM to the MMU role, support SMRAM address space") added the SMM address space, these helpers were not converted to use the active memslots. Consequently, an update to the SMM memslots can leave an entry from the old SMM address space apparently valid after the vCPU returns to the normal address space. A guest can then cause an access to valid RAM at the same GFN to be returned to userspace as KVM_EXIT_MMIO. Completing that exit through the VMM's RAM address space writes the backing page without going through KVM's write-tracking path. If the page backs a nested EPT table, KVM can continue using shadow translations derived from the old contents. The WARN_ON_ONCE() in kvm_mmu_write_protect_fault() catches this invalid cache and memslot combination. Use the vCPU's active memslots when caching and validating the entry. Memslot generations are unique across address spaces, so the generation also distinguishes the normal and SMM views. This matches the MMIO SPTE cache, which already uses kvm_vcpu_memslots(). On an unpatched current-mainline kernel, a regression test that fills the cache in SMM and updates only the SMM memslots fails the intended RAM write and triggers the kvm_mmu_write_protect_fault() warning. With this change, the same test completes the RAM write without a warning. The x86/smm_test, set_memory_region_test, and memslot_modification_stress_test selftests also pass.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: fs: Free any excess xarray nodes in clear_inode() For many years we've had a hard to hit leak of xarray nodes. Hugh documented it well in commit 786b31121a2c. Recently people and syzbot have found ways to force it to happen with madvise. Rather than fix the leaks where they happen, just call xa_destroy() which has the side-effect of cycling the i_pages lock.

Linux

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.

Linux

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.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: ipv6: udp: Create exceptions before socket matching Currently, when ICMPv6 Packet Too Big and Redirect Message packets are locally delivered and quote a UDP packet, an exception is only created in the IPv6 exception cache 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 exception cache in order to discover the ephemeral port used by a connected UDP socket. Commit 4785305c05b2 ("ipv6: use siphash in rt6_exception_hash()") and commit a00df2caffed ("ipv6: make exception cache less predictible") tried to mitigate such attacks by making it harder for attackers to discover hash collisions in the exception 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 an exception, even before trying to find a matching socket. Do that by calling ip6_update_pmtu() and ip6_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 an exception being created or updated in the cache that the attacker can observe. Pass the ifindex of the ingress device and the default uid, in a similar fashion to icmpv6_err(). Unlike IPv4, an oif of 0 would not match any nexthop in ip6_redirect_nh_match() and no exception would be created in response to a Redirect Message. Note that this does not allow attackers to create exceptions that they could not create before, as both helpers can already be reached with little to no validation. For example, by sending an ICMPv6 error that quotes an ICMPv6 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 ip6_sk_update_pmtu() / ip6_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 exception cache of the nexthop resolved by the no socket helpers and another in the exception cache of the nexthop used by the socket.

Linux

In the Linux kernel, the following vulnerability has been resolved: rust_binder: reschedule node refcount update on thread exit When a thread exits via BINDER_THREAD_EXIT, its pending work items are cancelled. If a thread exits while holding a pending node refcount increment (e.g. pushed as deferred work to that thread), the refcount increment was previously dropped because Node::cancel() and NodeWrapper::cancel() were no-ops. Dropping the refcount update leaves the node's delivery state and count state desynchronized, and userspace will not receive the notification, which can cause the node to never be freed from the process's nodes tree when all external references are dropped. Fix this by implementing DeliverToRead::cancel() for Node and NodeWrapper to move the pending refcount update to the process's work queue on thread exit so another thread can deliver it to userspace.

Linux

In the Linux kernel, the following vulnerability has been resolved: ipv4: stop PMTU walk when nexthop group shrinks Commit 7d3f3b4367f3 ("net: ipv4: Cache pmtu for all packet paths if multipath enabled") made __ip_rt_update_pmtu() update every path. For nexthop objects, fib_info_num_path() and fib_info_nhc() independently load nh->nh_grp. RCU protects each group's lifetime but does not make the loads observe the same group. If replacement shrinks the group after the loop accepts an index, fib_info_nhc() returns NULL and update_or_create_fnhe() dereferences it. A deterministic probe only widened the existing window between the real operations. On v7.2, an RTM_NEWNEXTHOP replacement published a one-member group after the PMTU reader accepted index 1 from a two-member group, producing: KASAN: null-ptr-deref in range [0x0000000000000000-0x0000000000000007] RIP: update_or_create_fnhe+0x45/0x15b0 Stop when the indexed path is absent. Groups are dense, so no later index exists in that snapshot. The same real-writer window completed 114,423 PMTU iterations without an oops with this check. A reproducer is available on request. Replacement requires CAP_NET_ADMIN in the network namespace. Where unprivileged user namespaces are permitted, a local user can obtain it in a new user and network namespace.

Linux

In the Linux kernel, the following vulnerability has been resolved: ipv4: stop exception dump when nexthop group shrinks Commit 4ce5dc9316de ("inet: switch inet_dump_fib() to RCU protection") allowed IPv4 FIB dumps to run concurrently with nexthop replacement. fib_dump_info_fnhe() uses fib_info_num_path() as the loop bound, then fib_info_nhc() independently reloads nh->nh_grp. If RTM_NEWNEXTHOP replaces a group with fewer paths between those loads, fib_info_nhc() returns NULL and the dump dereferences it through nhc_flags. Stop the walk when the indexed path is absent. Nexthop groups are dense, so the current snapshot has no later path.

Linux

In the Linux kernel, the following vulnerability has been resolved: ipv4: stop route notification sizing when nexthop group shrinks Commit 680aea08e78c ("net: ipv4: Emit notification when fib hardware flags are changed") added an RCU-only fib_nlmsg_size() call for asynchronous hardware flag notifications. fib_nlmsg_size() checks fib_info_num_path() before each iteration, then fib_info_nhc() independently reloads nh->nh_grp. If RTM_NEWNEXTHOP replaces a group with fewer paths between those loads, fib_info_nhc() returns NULL and fib_nexthop_nlmsg_size() dereferences it. Stop sizing when the indexed path is absent. Nexthop groups are dense, so the current snapshot has no later path.

Linux

In the Linux kernel, the following vulnerability has been resolved: ipv4: do not warn on route notification size race Commit 680aea08e78c ("net: ipv4: Emit notification when fib hardware flags are changed") added asynchronous route notifications for hardware flag changes. fib_alias_hw_flags_set() sizes the skb with fib_nlmsg_size() and later fills it with fib_dump_info() while holding only RCU. With nexthop compatibility mode enabled, a concurrent replacement can grow the group between these independent snapshots. fib_dump_info() can then legitimately return -EMSGSIZE, so the warning does not prove a sizing bug. Remove the warning. The existing error path still frees the skb and reports the error to listeners.

Linux

In the Linux kernel, the following vulnerability has been resolved: ipv6: do not warn on route notification size race fib6_info_hw_flags_set() sizes the skb with rt6_nlmsg_size() and later fills it with rt6_fill_node(). With nexthop compatibility mode enabled, a concurrent replacement can grow the group between these independent snapshots. rt6_fill_node() can then legitimately return -EMSGSIZE, so the warning does not prove a sizing bug. Remove the warning. The existing error path still frees the skb and reports the error to listeners.

Linux

In the Linux kernel, the following vulnerability has been resolved: Input: ati_remote2 - fix type confusion in device attribute handlers Previously, the channel_mask and mode_mask attributes were registered on the parent usb_device via sysfs_create_group(). When commit b20d6bf8014b ("Input: ati-remote2 - use driver core to instantiate device attributes") converted the driver to let the driver core manage them via dev_groups in struct usb_driver, the attributes moved to the bound usb_interface devices, but the show and store callbacks were not updated accordingly. As a result, the callbacks still invoked to_usb_device() on a device pointer belonging to a struct usb_interface, producing an invalid usb_device pointer. Use to_usb_interface() directly in the attribute handlers to obtain the interface and its driver data.

Linux

In the Linux kernel, the following vulnerability has been resolved: ALSA: ctxfi: Fix memory leakage when clearing DAO input mappers Commit e576e7843c0d ("ALSA: ctxfi: Simplify dao_clear_{left,right}_input() functions") centralized the DAO input mapper clearing logic in dao_clear_input(), but omitted the kfree() calls. It removes the mapper entries from the list and sets their pointers to NULL without freeing the allocated array. This leaks memory each time dao_clear_input() is called. Store the base pointer of the allocation before clearing the mapper entries and free the array after removing entries from the list.

Linux

In the Linux kernel, the following vulnerability has been resolved: btrfs: scrub: fix local_root reference leak in scrub_print_warning_inode() When paths_from_inode() fails, scrub_print_warning_inode() jumps to err without dropping the reference taken by btrfs_get_fs_root(), leaking a reference to the root every time path resolution fails while printing scrub warnings. Every other error and success path of the function drops the reference. Drop the reference on the paths_from_inode() failure path too.

Linux

In the Linux kernel, the following vulnerability has been resolved: can: dev: init_can_skb(): restore skb header initialization Commit 9f10374bb024 ("can: remove private CAN skb headroom infrastructure") removed the skb_reset_mac_header()/skb_reset_network_header()/ skb_reset_transport_header() calls from init_can_skb(). As a result, RX skbs from alloc_can_skb() and friends again carry mac_header = 0xFFFF. When such an skb reaches packet_rcv_spkt() (SOCK_PACKET), the push length calculation overflows and triggers skb_under_panic -> kernel BUG -> full machine panic. The same issue was originally reported in 2014 on linux-can and fixed by commit 969439016d2c ("can: add missing initialisations in CAN related skbuffs"). packet_rcv_spkt() itself has never been hardened: only packet_rcv() and tpacket_rcv() gained dev_has_header() checks in commit d549699048b4 ("net/packet: fix packet receive on L3 devices without visible hard header").

Linux

In the Linux kernel, the following vulnerability has been resolved: cgroup/cpuset: Handle cpu hotplug race in guarantee_active_cpus() With commit 2125c0034c5d ("cgroup/cpuset: Make cpuset hotplug processing synchronous"), the cpuset hotplug operation becomes synchronous. That commit also removes the code that handles race between cpuset_hotplug_work and cpu hotplug notifier with the assumption that race is now gone. Later commit 7a0aabd9ce69 ("cgroup/cpuset: Always use cpu_active_mask") updates the cpuset code to always use cpu_active_mask instead of cpu_online_mask in some places including guarantee_online_cpus() which is renamed to guarantee_active_cpus() in that commit. In the case of CPU offline operation, cpu_active_mask is updated first in sched_cpu_deactivate() to remove the offline CPU before cpuset_handle_hotplug() is called to update the effective_cpus of the affected cpusets. The cpu_online_mask is updated after that near the end of the offline operation to remove the offline CPU. As a result, the race comes back and the top cpuset may not have any active CPU leading to NULL pointer dereference during the race window when guarantee_active_cpus() is called after cpu_active_mask is updated to remove the CPU to be torn down but before cpuset_handle_hotplug() is able to properly update the effective_cpus of the top cpuset. Fix this by adding back the NULL cs check to avoid this problem.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: dm ioctl: don't let data_start run past the buffer struct dm_ioctl has seven bytes of trailing padding, so the smallest buffer userspace is allowed to ask for, offsetof(struct dm_ioctl, data), is seven bytes shorter than sizeof(struct dm_ioctl). get_result_buffer() sets data_start to sizeof(struct dm_ioctl) without looking at how big the buffer actually is. Ask for the minimum and data_start ends up seven bytes past the end. Callers add it to data_size, and ctl_ioctl() copies that much back to userspace: BUG: KASAN: slab-out-of-bounds in _copy_to_user+0xad/0xd0 Read of size 312 at addr ffff8880169e4c00 by task repro-dm/140 ctl_ioctl+0x5e3/0xcb0 dm_ctl_ioctl+0x25/0x40 allocated 305-byte region [ffff8880169e4c00, ffff8880169e4d31) A DM_LIST_DEVICES with data_size set to 305 is enough to hit it, and the seven bytes that leak are whatever sat next to the allocation. Clamp data_start to the buffer size. Note that this is not a security bug because the out-of-bounds read can only be triggered by root.

Linux

In the Linux kernel, the following vulnerability has been resolved: vsock/virtio: account only unread bytes in read_skb() After a partial stream receive, rx_bytes tracks the unread suffix while buf_used and peer credit still cover the whole packet. virtio_transport_read_skb() dequeues that packet but passes pkt_len for both counters, underflowing rx_bytes. On a connected CID_LOCAL stream with an SK_SKB verdict, a 40-byte read from a 100-byte packet made SIOCINQ report -40, an empty recv return ELOOP, and poll report the empty socket readable. After verdict detach, receiving 40 bytes wrapped the counter to zero and hid those queued bytes until one more byte arrived. The fixed run kept the counter balanced and exposed all queued data immediately. This was found during an LLM-assisted manual source audit of VSOCK receive accounting while re-evaluating virtio_transport_read_skb() after CVE-2024-50169. Subtract only skb->len minus the VSOCK offset from rx_bytes. Retain the full packet length for buf_used and peer credit.

Linux

In the Linux kernel, the following vulnerability has been resolved: wifi: brcmfmac: fix P2P device removal race in brcmf_detach() When the driver is removed while the user space process that created the P2P device (e.g. wpa_supplicant) is still exiting, the P2P device interface is removed twice and the kernel crashes. The two paths are: - wpa_supplicant, on exit, sends NL80211_CMD_DEL_INTERFACE for the P2P device. nl80211_del_interface() holds RTNL and the wiphy mutex and calls brcmf_p2p_del_vif(), which disables discovery and waits up to 1.5 s for BRCMF_E_IF_DEL. Then it calls brcmf_remove_interface(), which unregisters the wdev and frees the vif and the ifp. - brcmf_detach() sets the bus down, so BRCMF_E_IF_DEL never arrives, and calls brcmf_remove_interface(ifp, false) for the same ifp. brcmf_p2p_ifp_removed() reads ifp->vif and then blocks on rtnl_lock(). When brcmf_p2p_del_vif() releases RTNL, it calls cfg80211_unregister_wdev() on the wdev that was already unregistered and freed. The second list_del_rcu() of wdev->list faults on LIST_POISON2: brcmfmac: brcmf_p2p_del_vif delete P2P vif brcmfmac: brcmf_p2p_deinit_discovery enter brcmfmac: brcmf_p2p_del_vif P2P: GO_NEG_PHASE status cleared brcmfmac: brcmf_detach Enter brcmfmac: brcmf_bus_change_state 1 -> 0 brcmfmac: brcmf_remove_interface Enter, bsscfgidx=1, ifidx=0 brcmfmac: brcmf_del_if Enter, bsscfgidx=1, ifidx=0 brcmfmac: brcmf_p2p_ifp_removed P2P: device interface removed [1.5 s later, brcmf_p2p_del_vif() timed out] brcmfmac: brcmf_remove_interface Enter, bsscfgidx=1, ifidx=0 brcmfmac: brcmf_del_if Enter, bsscfgidx=1, ifidx=0 brcmfmac: brcmf_p2p_ifp_removed P2P: device interface removed Unable to handle kernel paging request at virtual address dead000000000122 Internal error: Oops: 96000044 [#1] PREEMPT_RT SMP pc : _cfg80211_unregister_wdev+0x6c/0x264 [cfg80211] Call trace: _cfg80211_unregister_wdev+0x6c/0x264 [cfg80211] cfg80211_unregister_wdev+0x10/0x1c [cfg80211] brcmf_p2p_ifp_removed+0x84/0xb0 [brcmfmac] brcmf_remove_interface+0x1f4/0x254 [brcmfmac] brcmf_detach+0x88/0x180 [brcmfmac] brcmf_sdio_remove+0x90/0x7b0 [brcmfmac] brcmf_sdiod_remove+0x20/0xa0 [brcmfmac] brcmf_ops_sdio_remove+0xbc/0x12c [brcmfmac] sdio_bus_remove+0x38/0x144 device_release_driver_internal+0x1e8/0x2c0 device_release_driver+0x14/0x20 brcmf_sdio_bus_remove+0x2c/0x40 [brcmfmac] brcmf_fwvid_unregister_vendor+0xd8/0x170 [brcmfmac] __exit_compat+0x18/0x418 [brcmfmac_cyw] __arm64_sys_delete_module+0x1ac/0x244 The log was taken with debug=0x406 on an i.MX 8M Plus board with a CYW55513 (Sona IF513) on SDIO. The driver is the Ezurio backport of brcmfmac from v6.18.22 on a 5.4-rt kernel. The code involved is the same in mainline. To reproduce: 1. Start wpa_supplicant on wlan0 and let it associate. It creates the P2P device (iw dev shows "type P2P-device"). 2. Stop wpa_supplicant and remove the driver without waiting for wpa_supplicant to exit: kill $(pidof wpa_supplicant) rmmod brcmfmac_cyw brcmfmac The crash is reproducible on that board. For the interface without a netdev (the P2P device), take RTNL and the wiphy mutex in brcmf_detach() before reading iflist, and remove it with locked=true. If brcmf_p2p_del_vif() runs first, brcmf_detach() finds the slot empty. If brcmf_detach() runs first, nl80211 does not find the wdev any more. The interfaces with a netdev are removed as before, because brcmf_del_if() takes RTNL on its own for the primary interface.

Linux

In the Linux kernel, the following vulnerability has been resolved: wifi: mt76: mt7603: initialize global station WCID mt7603 uses a driver-private global station WCID for frames that have neither a station nor a vif, including frames injected through a monitor interface. Unlike the core global WCID initialized by mt76_register_device(), this WCID is published without ever being passed to mt76_wcid_init(). Since commit 0335c034e726 ("wifi: mt76: fix race condition related to checking tx queue fill status"), mt76_tx() queues every such frame on the WCID's tx_pending or tx_offchannel queue and links its tx_list. The first injected frame therefore reaches __skb_queue_tail() with an uninitialized sk_buff_head and dereferences a NULL prev pointer. Initialize the WCID before publishing it. The fix was tested on an MT7603 PCIe device. Three wildcard probe requests were injected through a monitor interface. The hardware queue head and tail advanced from 0 to 3 and a second radio captured the third request over the air. No oops or reset occurred. After restoring AP mode, a 150 MB transfer also completed without a reset.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: workqueue: Fix NULL current_pwq deref in chained work check current_wq_worker() only tells us that %current is a kworker. It does not guarantee that the worker is currently executing a work item, as worker->current_pwq is populated only while process_one_work() is running the work function. is_chained_work() can be reached while queueing on a draining or destroying workqueue. If a kworker outside work-item execution reaches that path, current_wq_worker() returns a worker but worker->current_pwq is NULL, and the chained-work test can fault before it emits the intended warning. Guard the chained-work test with worker->current_pwq. A kworker without current_pwq is not executing work on the target workqueue, so the helper should return false and let the draining/destroying warning fire.

Linux

In the Linux kernel, the following vulnerability has been resolved: xen/netfront: don't leak the skb when xennet_fill_frags() fails When a response chain has more slots than fit in the skb's frags, xennet_fill_frags() returns an error and xennet_poll() jumps to its error path. That path moves what's left on tmpq to errq to be freed, but the skb being filled was already dequeued from tmpq, so it's never freed. Each chain that overflows leaks the skb and the pages attached to it as frags, and the backend decides how many slots it sends. Put the skb back on tmpq before taking the error path, like the xennet_set_skb_gso() failure just above it does.

Linux

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].

Linux

In the Linux kernel, the following vulnerability has been resolved: tcp: reject net_iov in zerocopy receive mapping hints After copying a readable prefix, receive_fallback_to_copy() asks tcp_zerocopy_set_hint_for_skb() where page mapping can resume. If the next skb is unreadable, find_next_mappable_frag() passes its net_iov fragment to can_map_frag(). skb_frag_page() returns NULL for a net_iov, but can_map_frag() dereferences it in PageCompound(). A v7.2 KASAN run on a connected TCP socket with 64 readable bytes followed by a 4096-byte NET_IOV_DMABUF fragment reported: BUG: KASAN: null-ptr-deref in can_map_frag tcp_zerocopy_receive -> can_map_frag Kernel panic - not syncing: KASAN: panic_on_warn set The diagnostic inserted the net_iov directly because the test host has no devmem-capable NIC. Hardware end-to-end reachability remains untested and requires CONFIG_NET_DEVMEM plus a supported DMA-buf-bound RX queue. Reject all net_iov fragments before skb_frag_page(). This covers both DMABUF and IOURING net_iov types while leaving page-backed checks unchanged. With the guard, the same queue copied the readable prefix, returned a 4096-byte skip hint, and completed without a fault.

Linux

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.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: selinux: preserve NATIVE_LABELS on already-initialized sb in set_mnt_opts When an NFSv4 filesystem exported with security_label is mounted a second time, sget_fc() returns the existing superblock and nfs_get_root() calls security_sb_set_mnt_opts() again on it. Since the superblock is already SB_INITIALIZED, selinux_set_mnt_opts() takes an early exit without setting SECURITY_LSM_NATIVE_LABELS in *set_kern_flags. nfs_get_root() then sees kflags without SECURITY_LSM_NATIVE_LABELS set and clears NFS_CAP_SECURITY_LABEL on the shared nfs_server, silently disabling per-file labeling for both the new and the original mount. Commit 3815a245b501 ("security/selinux: fix SECURITY_LSM_NATIVE_LABELS on reused superblock") fixed the same problem in selinux_sb_clone_mnt_opts(). The reused-superblock path for a plain second mount now goes through sb_set_mnt_opts rather than sb_clone_mnt_opts, so it needs the same treatment: on both SE_SBINITIALIZED early exits, set SECURITY_LSM_NATIVE_LABELS in *set_kern_flags when the superblock was already using SECURITY_FS_USE_NATVE.

Linux

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.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: gpio: xilinx: fix runtime PM leak on request error path pm_runtime_get_sync() leaves the usage counter incremented even when it fails, and xgpio_request() returns the error without dropping it. gpiolib does not call ->free() when ->request() fails, so xgpio_free(), which holds the only matching pm_runtime_put(), never runs. The reference is leaked and the controller can no longer runtime-suspend, so its clock stays enabled. Switch to pm_runtime_resume_and_get(), which only increments the usage counter on success.

Linux

In the Linux kernel, the following vulnerability has been resolved: gve: fix napi_disable deadlock when attempting to disable XSK pools When disabling XSK pools, GVE calls the unlocked versions of napi_disable and napi_enable. However, the netdev lock has already been acquired before ndo_bpf is called because GVE supports queue management ops. Calling the unlocked versions of napi_disable/enable results in a deadlock when attempting to disable XSK pools, as the thread attempts to re-acquire a lock it already holds. Update the NAPI calls to use the locked versions. This issue was found in production.

Linux

In the Linux kernel, the following vulnerability has been resolved: gve: fix XSK buffer leak when rings are stopped GVE does not free XSK buffers when resetting ring state as a part of stopping queues. This causes all XSK buffers which are posted to the NIC to be leaked. Free XSK buffers attached to an allocated buf_state when stopping rings. This issue was hit in production.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: ibmveth: h_free logical LAN on open-fail after register ibmveth_open() registers the logical LAN, then allocates the RX buffer pools. A pool failure jumps to out_free_buffer_pools without h_free_logical_lan(), so PHYP still owns the buffer-list page when it is unmapped. request_irq() failure already issued the hypercall before taking the same label. Move that h_free to the shared post-register unwind so both paths deregister before the buffer list is unmapped.

Linux

In the Linux kernel, the following vulnerability has been resolved: ice: fix use-after-free in dynamic port cleanup ice_dealloc_dynamic_port() uses dyn_port->vsi->idx to erase the dynamic port from pf->dyn_ports. However, it frees the VSI before reading the index for the erase, resulting in a use-after-free. Follow the reverse of the allocation order in ice_alloc_dynamic_port() by erasing the xarray entry before freeing the VSI.

Linux

In the Linux kernel, the following vulnerability has been resolved: ice: fix metadata_dst refcount handling on representor teardown ice_eswitch_release_repr() uses metadata_dst_free() to release the representor's metadata_dst. metadata_dst_free() directly frees the underlying memory without checking the dst_entry refcount. When ice_eswitch_port_start_xmit() processes a packet, it takes a reference via dst_hold() and attaches the metadata_dst to the skb. If the representor is torn down while packets are still queued on the lower device (e.g. in a qdisc), the metadata_dst is freed while references are still held. Use dst_release() instead, which correctly decrements the refcount and only frees the object when all references are dropped. The dst subsystem already handles metadata_dst cleanup in dst_destroy() when DST_METADATA is set. Other drivers sharing this pattern (nfp, airoha, bnxt) already use dst_release() for their metadata_dst lifecycle.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: ip6_gre: remove obsolete ERSPAN PMTU update A tc tunnel_key action can attach a METADATA_IP_TUNNEL dst to an skb, and mirred can redirect it to a native ip6erspan device. The legacy PMTU update in ip6erspan_tunnel_xmit() treats that metadata dst as a route. Since it has no output device, dst_dev(dst)->mtu dereferences NULL. On v7.2 with KASAN, an initial UID/GID 65534 process with CapEff 0 used self-created user and network namespaces to trigger: KASAN: null-ptr-deref RIP: ip6erspan_tunnel_xmit+0x10fc/0x2cc0 Kernel panic - not syncing: Fatal exception in interrupt The update is obsolete. Commit fe1a4ca0a2b7 ("ip6_gre: process toobig in a better way") changed IPv6 GRE Too Big handling to update the underlay route and removed the equivalent block from __gre6_xmit(). ip6_tnl_xmit() then propagates the underlay PMTU to the overlay route. Remove the stale ERSPAN copy and its now-unused dst variable. This eliminates the metadata-dst dereference and keeps ERSPAN aligned with IPv6 GRE PMTU handling.

Linux

In the Linux kernel, the following vulnerability has been resolved: KEYS: Fix add_key() race with keyring restriction __key_create_or_update() snapshots keyring->restrict_link before taking the destination keyring's semaphore. keyring_restrict() installs a restriction while holding that semaphore. This allows a writer to observe no restriction, wait for the keyring owner to install a reject-all restriction and return successfully, and then link a key using the stale NULL snapshot. The writer only needs write permission on the destination keyring. Move the restrict_link read after __key_link_lock() and __key_link_begin(). The read and the subsequent restriction check are then serialized with restriction installation by keyring->sem. The race was reproduced on v7.2.8 in 19 executions where restriction installation returned before the link completed. All 19 linked the key despite the reject-all restriction. With this change, 312 executions reached the same ordering and every add_key() call failed with -EPERM. The issue was found by manual concurrency analysis assisted by AI-based analysis and independently verified with a QEMU reproducer and kernel instrumentation.

Linux

In the Linux kernel, the following vulnerability has been resolved: memstick: core: wait for request completion before freeing card In memstick_alloc_card() and memstick_set_rw_addr(), wait_for_completion_timeout() is used to wait for device identification and address configuration requests to complete. The memstick subsystem does not provide an abort or cancellation callback in struct memstick_host for in-flight requests. If wait_for_completion_timeout() expires (for instance, when a USB transfer or command retry is delayed), memstick_alloc_card() returns an error and frees the memstick_dev structure via kfree(card). However, the host controller driver (such as rtsx_usb_ms) is still running the request asynchronously in worker context. When the transfer eventually finishes, the driver writes into host->req->data, which points to memory inside the already freed card structure. This leads to a KASAN slab-use-after-free write: BUG: KASAN: slab-use-after-free in ms_read_bytes+0x8fe/0x990 Write of size 1 at addr ffff88813c188079 by task kworker/1:5/5239 Workqueue: events rtsx_usb_ms_handle_req Call Trace: ms_read_bytes+0x8fe/0x990 drivers/memstick/host/rtsx_usb_ms.c:450 rtsx_usb_ms_issue_cmd drivers/memstick/host/rtsx_usb_ms.c:477 [inline] rtsx_usb_ms_handle_req+0x557/0x1a00 drivers/memstick/host/rtsx_usb_ms.c:531 process_one_work+0xac7/0x1b10 kernel/workqueue.c:3396 worker_thread+0x5ef/0xe50 kernel/workqueue.c:3560 kthread+0x373/0x450 kernel/kthread.c:436 Allocated by task 1662: memstick_alloc_card drivers/memstick/core/memstick.c:383 [inline] memstick_check+0x2c1/0x10d0 drivers/memstick/core/memstick.c:452 process_one_work+0xac7/0x1b10 kernel/workqueue.c:3396 Freed by task 1662: kfree+0x20c/0x650 mm/slub.c:6792 memstick_alloc_card drivers/memstick/core/memstick.c:420 [inline] memstick_check+0x72b/0x10d0 drivers/memstick/core/memstick.c:452 process_one_work+0xac7/0x1b10 kernel/workqueue.c:3396 Host controller drivers (rtsx_usb_ms, jmb38x_ms, tifm_ms, r592) already implement their own hardware and bus timeouts, and callers in mspro_block and ms_block already rely on wait_for_completion() unconditionally. Replace wait_for_completion_timeout() with wait_for_completion() in memstick_set_rw_addr() and memstick_alloc_card() so that the card structure is never freed while an asynchronous request is in flight.

Linux

In the Linux kernel, the following vulnerability has been resolved: memstick: rtsx_usb_ms: complete requests after eject instead of dropping them memstick_check() can pass its host->removing check just before rtsx_usb_ms_drv_remove() sets eject/removing. Its next request is then silently dropped: rtsx_usb_ms_request() skips schedule_work() once eject is set, and drv_remove's cancel_work_sync() can also cancel a queued handle_req before it picks the request up. Nobody completes card->mrq_complete. Once memstick core waits for requests without a timeout ("memstick: core: wait for request completion before freeing card"), this hangs removal: memstick_check() never returns and memstick_remove_host() blocks in flush_workqueue(): INFO: task kworker/u10:3:65 blocked for more than 20 seconds. Workqueue: kmemstick memstick_check __wait_for_common memstick_check INFO: task kworker/1:1:33 blocked for more than 20 seconds. Workqueue: usb_hub_wq hub_event __flush_workqueue memstick_remove_host rtsx_usb_ms_drv_remove Never drop a request. rtsx_usb_ms_request() always schedules handle_req, and handle_req fails requests with -ENOMEDIUM once eject is set, without touching the device. drv_remove flushes handle_req instead of cancelling it, and cancels it only after memstick_remove_host(), when no new request can arrive, so it cannot run on a freed host. The host_mutex drain in drv_remove is removed: handle_req always leaves host->req NULL when it finishes, so the drain never had anything to do, and it would now race with handle_req.

Linux

In the Linux kernel, the following vulnerability has been resolved: dm-crypt: reject the lmk IV mode with AEAD ciphers A "capi:" table line that pairs an AEAD cipher with the lmk IV mode, for example "capi:gcm(aes)-lmk" with the integrity:28:aead option, causes an oops when it is loaded. lmk sets tfms_count to 64, but crypt_alloc_tfms_aead only allocates one transform, so crypt_setkey reads past the end of the tfms_aead array and calls crypto_aead_setkey on the values it finds there. lmk exists for compatibility with Loop-AES, which uses CBC, so refuse it with an AEAD cipher rather than allocate 64 AEAD transforms. The check cannot go in crypt_iv_lmk_ctr with the other lmk checks, because the IV constructor runs after crypt_set_key. Note that this is not a security bug because the memory corruption can only be triggered by root.

Linux

In the Linux kernel, the following vulnerability has been resolved: dm: fix reading free memory in do_resume When do_resume calls dm_table_get_mode(new_map), the call is done without holding any locks (it only holds a reference to the md). It may be possible that another concurrent ioctl on the same device will swap table after dm_swap_table and before dm_table_get_mode. In this case, dm_table_get_mode(new_map) reads freed memory. This race condition was triggered by syzbot. This commit fixes the race by reading mode of the new table before dm_swap_table. Note that new_map is also passed to dm_ima_need_measure, but this function doesn't attempt to dereference it, so this call should be safe. Note that this is not a security bug because only root can trigger it and the commonly used tools such as lvm or cryptsetup do not call the ioctls concurrently, so that they can't trigger it acceidentally.

Linux

In the Linux kernel, the following vulnerability has been resolved: mmc: mtk-sd: Cancel request timeout work on remove The driver queues req_timeout while a request is active. msdc_request_done() uses cancel_delayed_work(), which does not wait for a timeout callback that has already started. The timeout callback calls mmc_request_done(), which wakes the request waiter, and then continues to use host->dev_comp and check the SDIO IRQ. During unbind, msdc_drv_remove() can return and the managed mmc_host can be freed before the callback finishes. KASAN reported use-after-free accesses in msdc_request_done() and msdc_recheck_sdio_irq() in each of three unbind tests. The same tests completed without a kernel diagnostic after this change. Call cancel_delayed_work_sync() after mmc_remove_host() has stopped new requests and before the driver releases the host resources.

Linux

In the Linux kernel, the following vulnerability has been resolved: mmc: sdhci-sprd: disable runtime PM on remove sdhci_sprd_probe() enables runtime PM, while sdhci_sprd_remove() does not perform the corresponding runtime PM cleanup. The probe failure path disables runtime PM before disabling the clocks, but the normal remove path directly disables clocks that are also managed by the runtime PM callbacks. If the device is runtime suspended, those clocks may already be disabled. Resume the device before removal, disable runtime PM and drop the temporary runtime PM reference before disabling the clocks. This issue was found by manual code inspection.

Linux

In the Linux kernel, the following vulnerability has been resolved: net/packet: guard the ll header push in packet_rcv_spkt() packet_rcv_spkt() restores the link layer header with skb_push(skb, skb->data - skb_mac_header(skb)); That subtraction is only meaningful when the device actually has a link layer header. packet_rcv() and tpacket_rcv() both wrap it in dev_has_header(), which is also the predicate the block comment at the top of the file states the restore in terms of; packet_rcv_spkt() does not. commit d549699048b4 ("net/packet: fix packet receive on L3 devices without visible hard header") introduced the helper and changed the two call sites, and this one stayed behind. A device without a visible ll header can leave skb->mac_header at the 0xFFFF sentinel that __alloc_skb() initialises it to. A CAN skb does: init_can_skb() sets pkt_type and ip_summed but does not reset the headers, and commit 9f10374bb024 ("can: remove private CAN skb headroom infrastructure") dropped the skb_reset_*_header() calls that used to be there. skb_mac_header() is then 0xFFFF, the length becomes a large negative number and skb_push() reports it through skb_under_panic() -- from softirq context, so it is a full system panic even with panic_on_oops=0: skbuff: skb_under_panic: text:ffffffff8bd21bc1 len:-65455 put:-65471 head:... data:... tail:0x50 end:0x180 dev:can0 kernel BUG at net/core/skbuff.c:214! RIP: 0010:skb_panic+0x50/0x60 Call Trace: <IRQ> skb_push+0x38/0x40 packet_rcv_spkt+0xe1/0x170 __netif_receive_skb_core.constprop.0+0x7e8/0xd30 ... Kernel panic - not syncing: Fatal exception in interrupt The socket type is reachable: packet_create() accepts SOCK_PACKET alongside SOCK_RAW and SOCK_DGRAM behind the same CAP_NET_RAW check, and neither the socket length nor a capability check keeps it away from a CAN interface. The missing skb_reset_*_header() calls in init_can_skb() are a regression in their own right and are being fixed separately, but a packet socket should not turn a link layer that did not initialise its mac header into a kernel panic. Guard the push the way the other two receive paths do. Tested on v7.3-rc5 under QEMU with a slcan device on a pty, which is the driver RX path: vcan does not reproduce it, because can_send() resets the headers on the way out. One SOCK_PACKET socket bound to can0 and one frame written into the line discipline panics an unpatched kernel with the trace above; the same image with this patch prints no panic and powers off normally. Both kernels are this tree, defconfig plus CONFIG_CAN_SLCAN=y, differing only in this hunk.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: net: bridge: avoid recursive multicast port cleanup br_multicast_sg_del_exclude_ports() removes automatically added STAR_EXCL port groups by calling br_multicast_del_pg(). For S,G entries, that deletion calls back into br_multicast_sg_del_exclude_ports(), so a large number of STAR_EXCL ports consumes one kernel stack frame per port and can hit the stack guard page. Keep the complete port-group deletion path for cleanup, but suppress the S,G exclude-port cleanup while that helper is already walking the same entry. Normal deletion callers retain the existing behavior.

Linux

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().

Linux

In the Linux kernel, the following vulnerability has been resolved: net: sparx5: free the matchall entry on destroy sparx5_tc_matchall_replace() allocates a struct sparx5_mall_entry for every offloaded matchall filter and adds it to sparx5->mall_entries. sparx5_tc_matchall_destroy() removes the entry from the list, but never frees it, so the entry of every deleted mirror and goto matchall filter is leaked. Free the entry after unlinking it. The leak was discovered by an AI code review agent, and reproduced with kmemleak on a lan969x EV board (EV23X71A) by repeatedly adding and deleting matchall mirror and goto filters. With the fix, kmemleak no longer reports the leak.

Linux

In the Linux kernel, the following vulnerability has been resolved: net: xilinx: axienet: Free outstanding DMA buffers on dmaengine stop In the dmaengine path the driver pre-submits RX buffers and holds in-flight TX buffers whose SKBs are DMA-mapped by the driver and freed only in the completion callbacks. On ndo_stop() dmaengine_terminate_sync() aborts these descriptors without running their callbacks, and the driver then frees only the ring shells, leaking every SKB still owned by the engine and its DMA mapping on each ifdown. With 128 RX buffers pre-posted per channel, the mapping leak can eventually exhaust a limited IOMMU aperture. Clear the slot's skb in the TX and RX callbacks so a non-NULL skb marks a slot that still owns a live, DMA-mapped buffer, and on stop unmap and free every such buffer. axienet_dma_rx_cb() runs from the DMA tasklet and re-arms the RX ring on each completion, so it can race axienet_stop(): a completion may submit a fresh buffer after dmaengine_terminate_sync() has returned, leaving the channel armed with a buffer the teardown then frees while the engine may still write into it (dma_release_channel() does not stop it either). Add a lock that axienet_dma_rx_cb() holds across the @stopping check and the resubmit, and axienet_stop() holds to set @stopping before terminating. Once @stopping is set no callback can arm a new buffer, and any armed just before is aborted by the terminate, so teardown only frees buffers the engine no longer owns.

Linux

In the Linux kernel, the following vulnerability has been resolved: net: team: stop reusing skb after queue override A nonzero dev_queue_xmit() result does not return ownership of the skb. When an override port's qdisc drops it, the queue override helper used to try another port or the mode transmit op with the freed skb. Return the handoff decision separately from transmit success. Select only the first override port under RCU and stop after handing it the skb, regardless of the lower transmit result. Keep the existing success and drop accounting. BUG: KASAN: slab-use-after-free in sk_skb_reason_drop Write of size 4 by task poc/5040 sk_skb_reason_drop (net/core/skbuff.c:1240) dev_kfree_skb_any_reason (net/core/dev.c:3470) ab_transmit (drivers/net/team/team_mode_activebackup.c:49) team_xmit (drivers/net/team/team_core.c:1849) __dev_direct_xmit (net/core/dev.c:4934) Freed by task 5040: __tcf_kfree_skb_list (net/sched/sch_generic.c:43) __dev_queue_xmit (net/core/dev.c:4831) team_xmit (drivers/net/team/team_core.c:1846) Kernel panic - not syncing: KASAN: panic_on_warn set ...

Linux

In the Linux kernel, the following vulnerability has been resolved: net: team: free skb when broadcast has no txable port bc_transmit() sends the original skb through the last txable port. When no port is txable, it returns false without consuming the skb, but team_xmit() still returns NETDEV_TX_OK. AF_PACKET sends then retain the skb and its socket write-memory charge. A TX-enabled, link-down port keeps the broadcast transmit op installed. Carrier can stay up through a second link-up, TX-disabled port or be forced on by userspace. Free the original skb on the no-port path while leaving the existing drop accounting intact. Discovered by Sashiko during review of previous commit.

Linux

In the Linux kernel, the following vulnerability has been resolved: openvswitch: fix soft lockup in the netlink flow dump A production compute node carrying a few thousand datapath flows hit a soft lockup inside a single netlink flow dump and panicked. ovs_flow_cmd_dump() calls ovs_flow_stats_get() for every flow it emits, and that releases stats->lock with spin_unlock_bh() once per CPU that has touched the flow. Every release is a local_bh_enable(), and each one runs the pending softirq backlog in the dumping thread's own context. The skb bounds the flows one callback emits, but not the softirq work it absorbs. On a CPU that carries the box's packet load the backlog refills as fast as it drains, so the dumping thread becomes that CPU's softirq engine. It never sleeps and it has no reschedule point, so under voluntary preemption nothing can take the CPU away from it: neither the ksoftirqd the kernel woke to take the work over, nor the stopper thread the softlockup detector dispatches to refresh its timestamp. Hold BH off across the whole callback instead, the way ctnetlink_dump_table() does, so the nested spin_unlock_bh() stop draining softirqs. The loop already runs under rcu_read_lock() and cannot sleep. What it gives up is preemption under CONFIG_PREEMPT, since a BH-off region is not preemptible outside PREEMPT_RT. The region stays short: the skb caps the flows one callback emits, and empty buckets cost no skb space but are each visited once per dump, as the cursor only moves forward. The table grows on insert and shrinks only on flush, so the walk is bounded by the largest flow count the datapath has held. The softirq backlog the callback used to absorb has no bound at all.

Linux

In the Linux kernel, the following vulnerability has been resolved: optee: register TEE devices only once fully initialized optee_probe() called tee_device_register() on both the client and the supplicant device before the rest of struct optee had been initialized. tee_device_register() calls cdev_device_add(), which does two things at once: it creates /dev/tee0 and /dev/teepriv0, and it links the device into the tee class so that class_find_device() can find it. From that moment on the device is reachable both from user space via tee_open() and from kernel space via tee_client_open_context(). The only gate in teedev_open() is tee_device_get(), which merely checks that teedev->desc is non-NULL, which was already set by tee_device_alloc(). Therefore, there is effectively no gate at all. A context opened in that window can run against a struct optee where - optee->call_queue.mutex is not initialized by optee_cq_init() - optee->supp mutex and completions is not initialized by optee_supp_init() - optee->rpmb_dev_mutex is not initialized yet, - the message argument cache is not initialized by optee_shm_arg_cache_init() - optee->ctx is still NULL Any open session or invoke during this window takes uninitialized mutexes and dereferences a NULL pointer. Fix it by moving both tee_device_register() calls down to the point where all of struct optee is set up. [JW: adding fixes and cc]

Linux

In the Linux kernel, the following vulnerability has been resolved: power: sequencing: pcie-m2: Fix leaking array from of_regulator_bulk_get_all() of_regulator_bulk_get_all() allocates the consumers array and hands it to the caller, who is expected to free it at an appropriate time. pwrseq-pcie-m2.c never did, leading to a memory leak. Add an explicit kfree() for the consumers array.

Linux

In the Linux kernel, the following vulnerability has been resolved: regulator: ltc3676: don't read the write-only HRST and CLIRQ registers HRST (0x1e) and CLIRQ (0x1f) are write-only command strobes. Writing HRST triggers a hard reset of the chip and writing CLIRQ clears the interrupt; reading them returns nothing useful, and the read itself fires the command. The result is that a plain register dump resets the board: cat /sys/kernel/debug/regmap/2-003c/registers walks every readable register, hits HRST, and the board reboots. Split the callback back in two so HRST and CLIRQ stay writable but are no longer readable. The readable range is left unchanged (LTC3676_BUCK1 ... LTC3676_IRQSTAT).

Linux

In the Linux kernel, the following vulnerability has been resolved: regulator: of: fill in supply names in of_regulator_bulk_get_all() of_regulator_bulk_get_all() returns an array it allocated itself, and fills in only the consumer of each entry. Every other way of getting a bulk array has the supply name set, because the caller provides it, and the core expects it to be there: regulator_bulk_enable() prints it when a supply fails to enable, so a caller that hands such an array to it dereferences uninitialised memory on that path. Copy each name into the array's own allocation, right behind the entries, so that it shares the array's lifetime and callers still have nothing extra to free. That also retires the fixed 64 byte stack buffer the names were assembled in, which is_supply_name() never bounded the copy against.

Linux

In the Linux kernel, the following vulnerability has been resolved: regulator: tps6594-regulator: Fix n_linear_ranges for TPS65224 BUCK2-4 The TPS65224 BUCK2, BUCK3 and BUCK4 descriptors pass 4 as n_linear_ranges, but tps65224_bucks_2_3_4_ranges[] only has two entries, covering selectors 0x00-0x45. The linear range helpers loop over n_linear_ranges entries, so any lookup of a selector above 0x45 reads past the end of the array and treats whatever follows it as two more ranges. This happens on the normal registration path when the core checks the voltage constraints against every selector. TPS652G1 uses the same descriptors and is affected too. Use ARRAY_SIZE() so the count always matches the table.

Linux

In the Linux kernel, the following vulnerability has been resolved: scsi: ch: Do not keep references to data transfer element devices ch_readconfig() looks up the scsi_device of every data transfer element whose SCSI id the changer reports in READ ELEMENT STATUS, and stores it in ch->dt[]. scsi_device_lookup() takes a reference, and nothing ever drops it: ch_destroy() frees the array with kfree(). ch->dt[] is read nowhere else - it only supplies the vendor, model and revision printed in the same loop. Once such a drive is removed, its scsi_device can never be released. It stays on the host's device list at its address, so a new device there is refused by anything that walks the list - target_core_pscsi reports "scsi_device_get() failed for H:C:T:L" - and the low-level driver's module can no longer be unloaded. Only a reboot recovers. It shows with any changer that reports its drives' ids; with the mhvtl virtual library (IBM 3573-TL personality), each create and remove of a library with four drives leaves four references behind, counted by the module's use count in lsmod. With ch not bound the count is unchanged, and with this patch applied it is unchanged too. Drop the reference as soon as the name has been printed, and remove the now unused dt[] array. That also removes the array leaked when ch_probe() fails after ch_readconfig(). The same leak was reported with an RFC patch in 2022, which was not merged.

Linux

In the Linux kernel, the following vulnerability has been resolved: riscv: entry: gate shadow call stack reload on sstatus.SPP, not tp Under CONFIG_SHADOW_CALL_STACK, handle_exception() reloads the kernel shadow call stack pointer (gp) only when the trap came from userspace. It made that decision with scs_load_current_if_task_changed, which compares the trap-time tp (read back from sscratch into s5) against the current tp: beq \prev, tp, _skip_scs scs_load_current _skip_scs: tp is a user-writable register. A user task that sets tp to the linear-map address of its own task_struct (&current) before trapping makes the comparison succeed (can achieve by brute force with KASLR enable under our test), so the kernel skips scs_load_current and keeps running with the gp that userspace left in place. Under CONFIG_SHADOW_CALL_STACK gp is the shadow call stack pointer and load_global_pointer is a no-op, so the compiler-emitted "sd ra, -8(gp)" / "ld ra, -8(gp)" in every non-leaf function then write to, and return through, an attacker-controlled address -- an unprivileged kernel arbitrary-write and control-flow-hijack primitive. The poisoned gp is saved to task_struct.scs_sp on return, so the reload is skipped on every subsequent entry as well. Decide "came from userspace" from the hardware sstatus.SPP bit (already in s1 at this point) instead of from tp, mirroring the check that ret_from_exception already uses a few instructions later. SPP is written by hardware on trap entry and cannot be forged from U-mode. Legitimate behaviour is unchanged: user entries (SPP=0) reload the task shadow call stack, nested kernel entries (SPP=1) skip it.

Linux

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.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: Bluetooth: hci_sock: Serialize dead-device detachment Rebinding an HCI socket after its controller is unregistered can race monitor control replay and crash the kernel. hci_sock_bind() clears hci_pi(sk)->hdev and drops the device reference under only the socket lock, whereas send_monitor_control_replay() holds hci_sk_list.lock for reading. Replay can observe a non-NULL hdev, then bind can clear it and drop the last reference before create_monitor_ctrl_open() reads hdev->id. This can cause a NULL pointer dereference if the pointer is reloaded, or a use-after-free if the old pointer is retained. The kernel reported: KASAN: null-ptr-deref in range [0x0000000000000068-0x000000000000006f] RIP: 0010:create_monitor_ctrl_open+0x579/0x800 Call Trace: hci_sock_bind+0xd96/0x1190 __sys_bind+0x166/0x200 __x64_sys_bind+0x6d/0xb0 do_syscall_64+0xdd/0x4a0 entry_SYSCALL_64_after_hwframe+0x77/0x7f The same missing serialization lets hci_send_to_sock() select a socket for the old device and queue its frame after bind has detached the socket or installed a new binding. Take hci_sk_list.lock for writing when clearing the device pointer and resetting the socket state. This waits for existing replay and delivery readers to finish before detachment and prevents later readers from using the old binding. Keep hci_dev_put() outside the critical section because the final device release can sleep.

Linux

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.

Linux

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

Linux

In the Linux kernel, the following vulnerability has been resolved: Bluetooth: ISO: Serialize concurrent connect calls iso_sock_connect() checks the socket state before taking the socket lock and drops the lock again before setting up a connection. Two callers can both pass the admission check while the socket is open or bound. Caller A can copy its destination for route selection, then caller B can replace the socket destination before A binds or connects the CIS. A then uses B's destination with the route selected for its own request. B can also proceed after A attaches a connection and sets BT_CONNECT. For deferred BIS setup, B can bind a second BIS and overwrite iso_pi(sk)->conn, leaving the first connection's reference and conn->sk back-pointer stranded. Concurrent CIS connects can likewise replace the socket's connection. Later teardown only detaches the current connection, so a callback on the old connection can race with socket release and access a freed socket. KASAN reported: BUG: KASAN: slab-use-after-free in iso_sock_hold+0xf7/0x1b0 Call Trace: iso_sock_hold+0xf7/0x1b0 iso_conn_del+0x7b/0x1d0 iso_connect_cfm+0x186/0x16a0 hci_conn_failed+0x154/0x280 hci_abort_conn_sync+0x3dc/0x7d0 hci_cmd_sync_work+0x173/0x300 Allocated by task 121: sk_alloc+0x2b/0x6d0 bt_sock_alloc+0x29/0x370 iso_sock_alloc.constprop.0+0x19/0x300 iso_sock_create+0x94/0x100 Freed by task 121: kfree+0x121/0x3c0 __sk_destruct+0x42b/0x540 iso_sock_release+0x29d/0x340 __sock_release+0xa1/0x260 Check admission under the socket lock and mark the connect operation in progress before publishing its destination. Keep that flag set across route lookup and connection setup, rejecting another connect with -EBADFD before it can change the destination or attach a connection. Clear the flag after either helper returns, including on failure, so a failed setup can be retried. A separate flag is needed because the socket lock must be released before acquiring the HCI device lock, while BT_CONNECT requires an attached connection. Keep the existing state transitions and the deferred CIS completion through iso_sock_recvmsg(). Reject an existing socket attachment in __iso_chan_add() before taking a new reference or publishing either pointer, so every caller preserves the connection association. Keep the same-socket, same-connection success case for deferred CIS setup. Reject an attached socket before BIS setup so a reusable BIS is not claimed before the attachment is rejected. In CIS setup, reject a different HCI connection before iso_conn_add() and drop the per-attempt HCI hold directly. An unowned iso_conn may still carry another reference after detachment, so its destruction cannot be relied on to release that hold. iso_listen_bis() creates a fresh PA HCI/ISO pair under the device lock, so the temporary iso_conn_put() frees the ISO candidate and drops its HCI hold on rejection. iso_conn_ready() uses a freshly allocated, unpublished child with no connection, making the socket-side rejection unreachable.

Linux

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.

Linux

In the Linux kernel, the following vulnerability has been resolved: Bluetooth: RFCOMM: free the skb when the DLC has no owner rfcomm_recv_data() hands the skb to d->data_ready() and returns 0 without freeing it, so the callback owns the skb from that point on. The tty implementation frees it when the DLC has been detached from its device: static void rfcomm_dev_data_ready(struct rfcomm_dlc *dlc, struct sk_buff *skb) { struct rfcomm_dev *dev = dlc->owner; if (!dev) { kfree_skb(skb); return; } The socket implementation returns without freeing it, and the skb is leaked. A connected DLC can have no owner: rfcomm_sock_destruct() clears d->owner, while the DLC itself stays on the session until it is closed. Data frames that arrive in that window still find d->state == BT_CONNECTED in rfcomm_recv_data() and are handed to the callback, one leaked skb each. The leak does not stop on its own either. The remote runs out of credits only once RFCOMM_RX_THROTTLED is set, and for a socket the only place that sets it is rfcomm_sk_data_ready() itself, below the return. rfcomm_process_dlcs() keeps granting credits, so the remote keeps sending and the leak keeps growing. Free the skb, as the tty side does.

Linux

In the Linux kernel, the following vulnerability has been resolved: Bluetooth: RFCOMM: connect the session socket without rfcomm_mutex An RFCOMM connect() issued while a BR/EDR link is being authenticated makes lockdep report a circular dependency, and the reported cycle is a real AB/BA between rfcomm_mutex and hdev->lock. rfcomm_security_cfm() is called from the HCI event path, which already holds hdev->lock: hci_rx_work() hci_event_packet() hci_cc_read_enc_key_size() [hdev->lock] hci_encrypt_cfm() [hci_cb_list_lock] rfcomm_security_cfm() [rfcomm_mutex] while an RFCOMM connect() from userspace takes the same two locks the other way round: rfcomm_sock_connect() rfcomm_dlc_open() [rfcomm_mutex] __rfcomm_dlc_open() rfcomm_session_create() kernel_connect() l2cap_sock_connect() l2cap_chan_connect() [hdev->lock] WARNING: possible circular locking dependency detected kworker/u131:1/1128 is trying to acquire lock: rfcomm_mutex, at: rfcomm_security_cfm+0x31/0x3e0 [rfcomm] but task is already holding lock: hci_cb_list_lock, at: hci_cc_read_enc_key_size+0x1d2/0xcc0 Chain exists of: rfcomm_mutex --> &hdev->lock --> hci_cb_list_lock hci_auth_complete_evt() and hci_encrypt_change_evt() reach the callback the same way. Both orders have to be seen in the same boot, which is why a BR/EDR connection alone is not enough to show it: a session set up by the remote side is created by rfcomm_accept_connection() in krfcommd, which calls kernel_accept() and never takes hdev->lock under rfcomm_mutex. Connecting a device that authenticates and encrypts the link and then calling connect() on an RFCOMM socket towards any address - the connect does not have to succeed, the order is recorded before the page timeout - reports it every time. Only the session socket has to be connected with the lock held, and it does not: nothing else can see the socket before it is put on the session list. So connect it first and take rfcomm_mutex afterwards, which removes the rfcomm_mutex -> hdev->lock order for good, rather than keeping the HCI event path out of rfcomm_mutex. rfcomm_session_create() becomes rfcomm_session_connect(), which returns the connected socket without touching the session list, and rfcomm_dlc_open() adds the session once it holds the lock again. If another opener added a session for the same pair while this socket was connecting, that session is used and this socket is dropped. __rfcomm_dlc_open() now takes the session it should use, and its state check runs after the lock is re-acquired, so a DLC that was opened or closed in the meantime is still handled. Over an existing ACL link the connection can complete before the session reaches the list, and the wakeup from the socket callback is then lost, so krfcommd is woken once the session is visible. Connecting without the lock opens a window in which the socket can be closed. The DLC is not on a session yet, so rfcomm_dlc_close() finds nothing to do and returns without touching it, and rfcomm_dlc_open() would then attach a DLC whose owner is gone. rfcomm_dlc_close() now marks such a DLC with RFCOMM_CLOSED, and rfcomm_dlc_open() checks the mark once it holds rfcomm_mutex again and drops the socket instead of attaching the DLC. The mark is cleared when an open starts, so a DLC that is reused is not affected. This covers both callers of rfcomm_dlc_open(), the socket and the tty layer.

Linux

In the Linux kernel, the following vulnerability has been resolved: Bluetooth: RFCOMM: Fix NULL tty_dev dereference in rfcomm_dev_shutdown syzbot hit a NULL pointer dereference in rfcomm_dev_shutdown() when an RFCOMM tty is closed. rfcomm_dev_add() puts the new device on rfcomm_dev_list first and only then registers the tty device and saves it in dev->tty_dev. If that registration fails, dev->tty_dev is never set at all. The device can be looked up in the meantime, so an open() of /dev/rfcommN gets through rfcomm_tty_install() and attaches to it. When that tty is closed again, rfcomm_dev_shutdown() uses dev->tty_dev without checking it: Oops: general protection fault, probably for non-canonical address KASAN: null-ptr-deref in range [0x0000000000000040-0x0000000000000047] RIP: 0010:rfcomm_dev_shutdown+0x48/0xb0 net/bluetooth/rfcomm/tty.c:131 Call Trace: tty_port_shutdown+0x1bf/0x220 drivers/tty/tty_port.c:372 tty_port_close+0x4e/0x150 drivers/tty/tty_port.c:705 tty_release+0x359/0x1670 drivers/tty/tty_io.c:1718 __fput+0x418/0xa50 fs/file_table.c:512 rfcomm_dev_destruct() already copes with a NULL dev->tty_dev. Add the same check to rfcomm_dev_shutdown(), so the tty device is only moved when it was actually registered.

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/fb-helper: defer setup when fbdev_probe() fails, not just on -EAGAIN __drm_fb_helper_initial_config_and_unlock() calls drm_client_modeset_probe() first, which already fills in each mode_set->mode/num_connectors for every connected output, and only afterwards calls drm_fb_helper_single_fb_probe() to create fb_helper->fb and wire it into those mode_sets via drm_setup_crtcs_fb(). If drm_fb_helper_single_fb_probe() fails with anything other than -EAGAIN (e.g. dev->driver->fbdev_probe() itself fails), the function bails out before drm_setup_crtcs_fb() runs and before setting fb_helper->deferred_setup. The client is left with mode_sets that have a mode and connectors but mode_set->fb == NULL, and nothing marks this fb_helper as needing a retry. __drm_fb_helper_restore_fbdev_mode_unlocked() only skips committing when fb_helper->deferred_setup is set, so a later .restore() call (drm_lastclose() on process exit, in this report) commits the stale mode_set through drm_client_modeset_commit() -> drm_client_modeset_commit_atomic() -> __drm_atomic_helper_set_config(), which hits WARN_ON(!set->fb). Mark fb_helper->deferred_setup on any drm_fb_helper_single_fb_probe() failure, not just -EAGAIN, so restore keeps skipping the commit until a later hotplug event successfully re-probes and re-populates mode_set->fb. The real error code is still returned to the caller unchanged.

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/i915/gem: Prevent overstepping exec array boundary It is possible for eb_relocate_parse_slow() to call kvfree() on entries outside the original exec array during following scenario inside eb_relocate_parse_slow(): 1) eb_copy_relocations() allocates as many as eb->buffer_count copies for exec entries; 2) eb_parse() appends up to two VMAs, right after incrementing eb->buffer_count; 3) cleanup under "out" label uses this incremented buffer_count to release relocation pointers from exec array via kvfree(), not from VMA array. To amend this problem, ensure that relocation cleanup uses the original exec object count, excluding the VMAs appended by the command parser. (cherry picked from commit 2a0158925fb943c7ba8f3216e2734952128defc9)

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/i915/gt: Unmask interrupts only when ACTIVE is true There is a small race window between intel_rps_boost() reading a stale ACTIVE state and queuing rps->work, where ACTIVE may be changed via asynchronous calls triggered by wakeref drop. This means that rps_work() could be triggered after the GT has been parked, so unmasking RPS PM interrupts would happen without holding PM wakelock. This sequence of events may be observed during live_rc6_ctx_wa() selftest: 1) live_rc6_ctx_wa() -> intel_engine_pm_put() -> wakeref_put() ASYNC1) wakeref_put() callback -> __engine_park() -> intel_gt_pm_put_async() -> intel_gt_pm_put_async_untracked() -> intel_wakeref_put_async() ASYNC2) wakeref_put_async() callback -> __gt_park() -> intel_rps_park() -> intel_rps_clear_active() [clear ACTIVE] WITH INTERRUPTS: -> rps_disable_interrupts() -> cancel_work_sync(rps->work) -> rps_reset_interrupts() [rps->pm_iir = 0] WITHOUT INTERRUPTS (TIMER mode): -> rps_stop_timer() -> cancel_work_sync(rps->work) 2) live_rc6_ctx_wa() -> intel_gt_wait_for_idle() -> intel_gt_retire_requests_timeout() -> dma_fence_wait_timeout() -> ops->wait() [i915_fence_wait()] -> i915_request_wait_timeout() -> intel_rps_boost() -> intel_rps_is_active() [check ACTIVE] [OPEN RACE WINDOW] -> queue_work(rps->work) -> retire_requests() -> i915_request_retire() -> intel_rps_dec_waiters() [from 1 to 0] If ASYNC1 and ASYNC2 would execute during [OPEN RACE WINDOW], then: 3) rps_work() executes from the workqueue and may satisfy "(!pm_iir && !client_boost)" test (in TIMER mode rps->pm_iir is zeroed by an earlier rps_work() call or never gets a new value after previous intel_rps_unpark() event), so it jumps to "out" label, bypassing intel_rps_is_active() guard and executes gen6_gt_pm_unmask_irq(); 4) GT is parked and runtime PM wakelock is dropped, so only a raw reference in __gt_park() -> intel_display_power_put_async() remains for 100ms, then the MMIO write in gen6_gt_pm_unmask_irq() detects missing wakelock via assert_rpm_wakelock_held() and emits "RPM wakelock ref not held during HW access" tainting kernel. Commit 3e7abf814193 ("drm/i915: Extract GT render power state management") dropped the "if (rps->interrupts_enabled)" guard that gen6_pm_rps_work() had around this unmask in i915_irq.c. INTEL_RPS_ACTIVE brackets the same window, since it is set in intel_rps_unpark() before rps_enable_interrupts() and rps_start_timer() and cleared in intel_rps_park() before rps_stop_timer() and rps_disable_interrupts(). Skip unmasking when RPS is no longer active (reaching unmask operation is independent of INTERRUPT mode). This does not prevent RPS operation from resuming on the next unpark. In interrupt mode, intel_rps_unpark() calls rps_enable_interrupts(), which enables and unmasks the PM interrupts. In timer mode, it calls rps_start_timer(), restarting periodic engine-busyness evaluation. The timer generates software RPS events and queues rps_work() directly, independently of PM interrupt delivery. Neither mode requires the late worker to unmask interrupts while the GT is parked (cherry picked from commit ca4a6deadad15b91ee083afe445fa6bdbfb68206)

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/imagination: Size page table preallocation by device address Commit 0a8224058a58 ("drm/imagination: Fix page count for page table for map() interface") passed the device address to pvr_mmu_op_context_create(), but the preallocation is still sized from device_addr + sgt_offset with an exclusive end. The offset into the object's pages has no place in a device-virtual range, and the exclusive end allocates one table too many when the range ends on a table boundary. Count the tables from device_addr and the inclusive end of the range. The new calculation does not take page table indices into account. This fixes another issue with the previous implementation, where the difference between start and end indices could underflow when a mapping crossed a page table boundary. This would lead to the driver preallocating a huge number of MMU pages, exhausting system memory and triggering an out of memory panic. [Alessio: add paragraph to point out OOM bug being fixed, cc stable]

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/panthor: validate userspace queue count against initialized firmware slot count In panthor_drv.c:panthor_ioctl_group_create(), the queue count is validated against the hardcoded maximum: if (!args->queues.count || args->queues.count > MAX_CS_PER_CSG) return -EINVAL; However, if the firmware reports fewer slots than MAX_CS_PER_CSG, panthor_fw_init_ifaces() will only initialize that smaller number of stream interfaces. If unprivileged userspace requests more queues than the firmware reported (but no more than MAX_CS_PER_CSG), cs_slot_prog_locked() fetches an uninitialized stream interface and may crash when dereferencing cs_iface->input->ringbuf_base.

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/vmwgfx: validate pitch coming from userspace Validate the pitch, alongside the box dimensions before trying to copy data from the underlying surface. Fixes possible out of bounds reads with cursor snooping.

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/xe/i2c: cancel the client work on remove xe_i2c_notifier() stores the DesignWare adapter in i2c->adapter and schedules i2c->work when the adapter is registered under the xe I2C platform device, and xe_i2c_client_work() then instantiates the AMC client device on that adapter. xe_i2c_remove() tears down the AMC, unregisters the client devices, the bus notifier and the adapter platform device, but it never drains i2c->work. A work item that is still queued or running when the adapter is unregistered dereferences i2c->adapter in i2c_new_client_device() after platform_device_unregister() has released the adapter. The work item is also embedded in the devm-allocated struct xe_i2c, so a work item still queued after the drm device devm unwind frees that allocation runs its callback on freed memory. The bus notifier is the only thing that schedules this work, and it is unregistered after the client devices. An instance that is still queued when the teardown runs can therefore write i2c->client[XE_I2C_CLIENT_AMC] while the loop is unregistering and clearing the same array, and an AMC client it instantiates late is only cleaned up by the adapter's own child sweep in i2c_del_adapter(). Move bus_unregister_notifier() in front of the client teardown loop and cancel the work right after it, so no new instance can be scheduled and a queued instance is drained before the client array is touched. A running instance still finds a live adapter, since the adapter is unregistered later. (cherry picked from commit a23f27d5ec11c3f418c5e5531eea34ccda2b5b20)

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/xe: Fix shrinker accounting double-subtraction on nested external pins xe_bo_pin_external() and xe_bo_unpin_external() support nested pins, since the underlying ttm_bo_pin()/ttm_bo_unpin() maintain a pin_count refcount rather than a boolean. However, the shrinker accounting calls xe_ttm_tt_account_subtract() and xe_ttm_tt_account_add() are invoked unconditionally on every pin_external()/unpin_external() call, instead of only on the transition into and out of the pinned state. An external BO (dma-buf) can be pinned more than once while already pinned, for example when it has multiple attachments/importers, each independently calling xe_bo_pin_external(). Each such nested pin subtracts the BO's pages from the shrinker's accounting again, even though the BO's pages were already removed from consideration by the first pin. Symmetrically, each nested unpin adds them back again before the BO is actually unpinned. This desynchronizes the shrinker's page and object counts from reality, and since they are declared as signed long but interpreted as unsigned long in xe_shrinker_count(), the underflow can turn into an enormous shrinkable/purgeable page count, causing the shrinker to be invoked excessively under memory pressure. Guard the accounting calls with the same pin-count transition checks already used to guard the pinned_link list maintenance, so accounting is only updated on the outermost pin and the final unpin. While centralizing this logic into xe_bo_account_pin()/ xe_bo_account_unpin(), also close a second, related accounting gap affecting imported dma-bufs. For an imported bo's ttm_tt, xe's ttm_tt_populate hook (xe_ttm_tt_populate()) early-returns without calling xe_ttm_tt_account_add(), since TTM_TT_FLAG_EXTERNAL is set without TTM_TT_FLAG_EXTERNAL_MAPPABLE. However, TTM core's ttm_tt_populate() wrapper unconditionally marks the tt as populated on driver-hook success, so ttm_tt_is_populated() still reports true for it. xe_bo_account_pin()/xe_bo_account_unpin() trusted ttm_tt_is_populated() as a proxy for "these pages are accounted for by the shrinker", and would subtract/add pages for such a tt that were never added in the first place, again underflowing the shrinker's counts. This is reachable today via xe_bo_pin_external(), and would have become newly reachable via the fb-pin path once this patch routes __xe_pin_fb_vma() through xe_bo_pin_account() below, so fix it here rather than in a later patch, to avoid a commit in this series that introduces the regression before fixing it. Skip the accounting for imported bos, using the same EXTERNAL && !EXTERNAL_MAPPABLE test xe_ttm_tt_populate()/xe_ttm_tt_unpopulate() use to decide whether to act, via a new xe_ttm_bo_is_imported() helper (hoisted up from its only prior use site). v2: - Fix the critical issue caught by Sashiko AI review: the pin-count transition check introduced here was still bypassed by the fb-pin path (__xe_pin_fb_vma()/__xe_unpin_fb_vma()), which could leave a stale pinned_link entry and a permanent shrinker accounting leak when the fb-unpin performed the true last unpin. Addressed in a separate follow-up patch in this series rather than folded in here, since it is a distinct bug with its own Fixes: tag. - Drop the unneeded local "last_unpin" variable in xe_bo_unpin_external(), restoring the original inline bo->ttm.pin_count == 1 check; it was a no-op rename not used by any of the new accounting helpers, and was going to be deleted again by the follow-up patch anyway. v3: - Exclude imported dma-bufs from the pin/unpin shrinker accounting, fixing another underflow caught by Sashiko AI review. - Add a short comment on the backup-object pin site noting it's already populated/accounted at pin time, addressing (as a false-positive clarification, no code fix needed) a separate Medium-severity accounting-leak concern also caught by Sashiko AI ---truncated---

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/xe: Fix stale pinned_link entry when fb-pin performs the final unpin xe_bo_pin_external() and xe_bo_unpin_external() maintain the bo's pinned_link list membership in xe->pinned.late.external, based on whether the current call is the outermost pin or the final unpin. However, the same external bo's pin_count can also be raised and lowered directly by __xe_pin_fb_vma()/__xe_unpin_fb_vma(), which pin the bo as a display scanout buffer without going through xe_bo_pin_external()/xe_bo_unpin_external() at all, and have no notion of, or ownership over, pinned_link. If a bo is pinned both externally (e.g. dma-buf export) and as an fb, and the external unpin happens first, xe_bo_unpin_external() correctly observes pin_count > 1 and leaves the bo on pinned_link. When the fb unpin later performs the true last unpin (pin_count 1 -> 0), it never touches pinned_link, leaving the bo linked on xe->pinned.late.external indefinitely. Once the bo is subsequently freed, this stale list entry points into freed memory, corrupting the list and risking a use-after-free the next time the list is walked or spliced. The backup object pin/unpin sites in xe_bo_notifier_prepare_pinned()/ xe_bo_notifier_unprepare_pinned() have the same bypass characteristic, though they never add their bo to a pinned list, so are not affected by this particular list-corruption issue. Move the pinned_link removal into xe_bo_account_unpin(), which already runs on every unpin path (kernel, external, framebuffer, backup object) right before the true 1 -> 0 pin_count transition. Since list_del_init() only operates on the node itself, this removal is list-agnostic and safe to perform regardless of which list (external or kernel_bo_present) the bo happens to be linked on, or which code path is performing the final unpin. Drop the now-redundant explicit list_del_init() calls in xe_bo_unpin_external() and xe_bo_unpin(). v2: - New patch v3: - Rebased on changes to previous patch. (cherry picked from commit ca58d36a1f93af586ee93b340c9f9fd81312cc01)

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/bridge: th1520-dw-hdmi: Fix remove() callback This driver stores struct th1520_hdmi * in platform drvdata, but th1520_dw_hdmi_remove() was reading it back as struct dw_hdmi * and passing it to dw_hdmi_remove(), so teardown runs on the wrong pointer. Retrieve struct th1520_hdmi * from platform drvdata and pass hdmi->dw_hdmi to dw_hdmi_remove().

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/amdgpu/userq: only accept doorbell BOs as queue doorbell amdgpu_userq_get_doorbell_index() pins any BO passed as doorbell_handle into the doorbell domain. For a regular BO, TTM moves it with ttm_bo_move_null() and its contents are lost. Since an importer on the same device gets the exporter's GEM object, a client can do this to a buffer another client shared with it. Reject BOs that were not created in the doorbell domain. v2: drop the error message, use a local abo variable (Christian) (cherry picked from commit f7497bbc19db4821c62b9d97b31c9086feb2d917)

Linux

In the Linux kernel, the following vulnerability has been resolved: fs: don't silently unmount busy mounts The propagate_mount_busy() helper exists to decide whether a synchronous can suceed. It does that by looking at the copies of the victims in the mounts its parent propagates to. One of those functions from propagation hell. It skips a copy that has child mounts except for the case of a single child mount covering the copy's root. That algorithm used to be exact. It isn't anymore. Originally only childless copies were unmounted during propagation. And 1064f874abc0 ("mnt: Tuck mounts under others instead of creating shadow/side mounts.") ended up adding covered mounts to both sides in one go. Along came 99b19d16471e ("mnt: In propgate_umount handle visiting mounts in any order") and that started to unmount more than before: [1]: A copy of a mount is unmounted when each of its children is either its overmount or a copy of the victim itself. In the former case the overmount gets reparented. In the latter case the copy of the victim will end up being unmounted. So propagate_umount() ended up being rewritten without adjusting propagate_mount_busy(). And currently propagate_mount_busy() doesn't look at copies of the victim that have two or more children nor a single child that is not the overmount. So any such copy in [1] is not checked for references. The result is that a synchronous unmount succeeds, pulling out a mount that is still in use. A container can run into this with its own mounts. The host shares a tree with the container and the host has a mount in that tree: mount -t tmpfs none /mnt mount --make-shared /mnt mount -t tmpfs none /mnt/a The container's /mnt is a slave copy of the host mount. The container now takes a detached copy of the tree and mounts that copy beneath its propagated copy of /mnt/a and keeps the file descriptor to the tree open: unshare -m --propagation unchanged mount --make-rslave / fd = open_tree(AT_FDCWD, "/mnt", OPEN_TREE_CLONE | AT_RECURSIVE) move_mount(fd, "", AT_FDCWD, "/mnt/a", MOVE_MOUNT_BENEATH) grep /mnt /proc/self/mountinfo 775 355 ... /mnt ... master:646 777 775 ... /mnt/a ... master:646 776 777 ... /mnt/a ... master:647 778 777 ... /mnt/a/a ... master:647 The now attached tree (777) is now mounted beneath the container's copy of /mnt/a. That's also where the host mount sits (as seen from the host mount namespace). So the propagated copy (776) of the host's mount is on top of its root and the copy of that (778) is inside it. Now the host runs: umount /mnt/a and it succeeds. The moved tree tree and the copy inside of it are gone from the container. Now only the propagated copy is left and it got reparented to where it was before: grep /mnt /proc/self/mountinfo 775 355 ... /mnt ... master:646 776 775 ... /mnt/a ... But the file descriptor still refers to the moved tree. So a synchronous umount detached a mount that is still in use. Teach propagate_mount_busy() the same checks that propagate_umount() does. It now checks for a single victim mount without children for which it can assume that MNT_LOCKED is cleared on every copy like propagate_umount() does. The copies form chains. IOW, when a copy itself receives propagation from the victim's parent then the mount at the same mountpoint inside that copy is a copy as well. Whether a mount receives propagation from the victim's can simply be checked by walking its masters. Walk each chain once from its top so we are sure that nested copies stay linear. Check every reference once of every copy that gets unmounted.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: reset the old parent's ->overmount in mnt_change_mountpoint() __umount_mnt() resets ->overmount of the parent when it unmounts the mount that overmounts it. mnt_change_mountpoint() doesn't. It moves the mount to its new parent and leaves the old one pointing at it. For a mount that is moved beneath in attach_recursive_mnt() that doesn't matter. The new mount has already reset ->overmount of the old parent by the time the old top mount is moved onto it. It matters for reparent() though. propagate_umount() moves a surviving overmount onto the first ancestor that isn't going to be unmounted. The mount it was taken from is unmounted immediately afterwards. If a file descriptor keeps that unmounted mount alive ->overmount points at the moved mount. That's a dangling pointer and UAF as the moved mount is unmounted and freed. do_lock_mount() follows ->overmount from the target path under mount_lock. With a file descriptor on such a mount as the target of a mount beneath topmost_overmount() walks into the freed mount and do_lock_mount() takes references on whatever it finds there. So reset ->overmount of the old parent in mnt_change_mountpoint() the same way __umount_mnt() does. A mount's ->overmount is then always a mount that is attached to it.

Linux

In the Linux kernel, the following vulnerability has been resolved: statmount: read the parent of an unmounted mount under mount_lock statmount_mnt_basic() reads the ids of the parent through m->mnt_parent with namespace_sem held for reading. That's fine for a mount that is in a namespace. Everything above it is pinned by that namespace. STATMOUNT_BY_FD also accepts a file descriptor on a mount that has been unmounted on purpose and that only the file descriptor keeps alive. If that mount stayed attached to its parent when it was unmounted the parent is unmounted as well and nothing pins it for the child. The final mntput() unmounts the children under mount_lock alone and frees the parent afterwards. Concurrently, statmount_mnt_basic() reads m->mnt_parent twice while that can happen and can easily dereference a mount that's already gone. Read the parent of an unmounted mount under mount_lock. The parent can't be put while it is held and both ids come from the same mount. A mounted mount's parent is pinned by the namespace and a mount can't leave its namespace while namespace_sem is held, so the common case takes no global lock.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: don't inherit MNT_UMOUNT in clone_mnt() MNT_UMOUNT gets raised on a mount that umount_tree() or umount_one() has taken out of its namespace. It's never cleared. __detach_mounts() unhooks a mount that carries it instead of unmounting it. disconnect_mount() lets an unmounted child stay attached only if its parent carries it. propagate_umount() takes it as the sign that a mount is going to be unmounted when it decides what else can be unmounted, where a chain of locked mounts ends and where a for reparenting an overmount. clone_mnt() copies the flags of the source and only masks MNT_INTERNAL_FLAGS which doesn't include MNT_UMOUNT. A mount that was lazily unmounted but is kept alive by a file descriptor can still be copied when its root is an nsfs or pidfs file because may_copy_tree() accepts those before checking whether the source is mounted at all: mount --bind /proc/self/ns/net /tmp/x exec 3< /tmp/x umount -l /tmp/x mount --bind /proc/self/fd/3 /tmp/y The mount on /tmp/y is attached to the caller's namespace with MNT_UMOUNT already set and every copy of it inherits the flag again. That causes a number of bugs: - Unlink its mountpoint from a mount namespace where it isn't one. __detach_mounts() takes it for a mount that was unmounted earlier, unhooks it from its parent and drops the reference the namespace holds on it. It stays in the namespace with one reference too few and as its own parent, so it can't be unmounted through the file descriptor anymore and the namespace never finds it again. When the namespace goes away it's leaked and points to a freed struct mnt_namespace that may_change_propagation() and anon_ns_root() read on the next propagation change or mount_setattr() through the file descriptor. - Make it shared and unmount a peer. trace_transfers() follows the flag from the peer into it and strips it from its peer group. - Let a propagated umount reach it. gather_candidates() marks it with T_UMOUNT_CANDIDATE but never puts it on the candidate list, so the mark is never cleared and every later propagated umount skips it. If it's the parent of a locked mount handle_locked() takes it as the cutoff and commits the locked child that should have stayed and disconnect_mount() leaves that child attached to it, unmounted under a mounted parent, which is what disconnect_mount() exists to prevent. When the namespace goes away umount_tree() visits the child a second time. And once __detach_mounts() has made it its own parent reparent() never finds the end of the chain of unmounted parents. A clone is a new mount that has never been unmounted. Add MNT_UMOUNT to MNT_INTERNAL_FLAGS so clone_mnt() never copies it. And warn when umount_tree() reaches a mount that already has it set.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: don't copy an unmounted mount tree may_copy_tree() accepts a mount whose root is an nsfs or pidfs file before it checks whether that mount is mounted. This exception exists for nsfs and pidfs mounts. They are part of no mount namespace so neither check_mnt() nor is_mounted() would let them through and bind-mounting a nsfsfd or a pidfd wouldn't work. Problem is that exception also admits a bind mount of such a file when it was lazily unmounted and is only kept alive by a file descriptor: mount --bind /proc/self/ns/net /tmp/x exec 3< /tmp/x umount -l /tmp/x mount --rbind /proc/self/fd/3 /tmp/y With MS_REC or AT_RECURSIVE copy_tree() doesn't stop at the unmounted mount. It copies every mount that stayed attached to it as well. So if you have a bunch of overmounts over the nsfsfd or the pidfd the whole unmounted stack is copied. That's a bug. So refuse to walk an unmounted tree. Copying its root alone stays possible. So a bind mount of a namespace that is gone can still be recreated from a file descriptor on the old one. The file descriptor pins that mount and its root and __has_locked_children() only reads the direct children which don't leave the parent's list without namespace_sem. The internal mounts carry MNT_NS_INTERNAL instead of NULL and are still walked. Copying an unmounted tree recursively fails with EINVAL.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: queue a mount only once for mount notifications mnt_notify_add() puts a mount on notify_list. It doesn't check whether the mount is on the list already. That's fine as long as every mount is queued at most once per namespace_sem hold. It isn't. propagate_umount() moves a surviving overmount off a stack of mounts that are going away and queues it as moved. shrink_submounts() and mark_mounts_for_expiry() call umount_tree() for several mounts under a single namespace_sem hold. So the second umount_tree() can take down exactly the mount the first one reparented, reparent it once more and so end up queueing it a second time. The second list_add_tail() cuts the mounts queued in between out of the list while the head still points at the last of them. notify_mnt_list() then only visits that one mount and the others are freed after the grace period with notify_list still pointing at them. Every later mount operation on the host walks freed memory: BUG: KASAN: slab-use-after-free in __list_add_valid_or_report Read of size 8 at addr ffff8881003d57e0 by task notify_dq/162 __list_add_valid_or_report+0x15c/0x1a0 mnt_add_to_ns+0x322/0x890 attach_recursive_mnt.isra.0+0xf2a/0x1a00 path_mount+0x139c/0x1d40 This needs a mount with MNT_SHRINKABLE to bind from. That's what finish_automount() creates (submounts of NFS, AFS and CIFS and the tracefs mount below debugfs). Bind mounts inherit the flag. A mount namespace with a FAN_MARK_MNTNS mark does the rest and root in a user namespace can have both. Initialize to_notify in alloc_vfsmnt() and leave a mount alone that is queued already. mnt_notify() looks at the state the mount has when the list is drained so a single entry per mount is enough. The move event for a mount that is unmounted under the same hold is lost. That's fine. The detach is what matters.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: check a submount for references right before unmounting it shrink_submounts() and mark_mounts_for_expiry() first collect all the mounts they are allowed to unmount and then unmount them. Whether a mount is busy is decided by propagate_mount_busy() on the tree. But unmounting the first mount changes the tree that the second one propagates into. propagate_umount() moves a surviving overmount off a copy it unmounts and mounts it where the copy was mounted. If that's where the propagated copy of the second victim is looked up the overmount becomes a candidate of the second umount_tree(). It's childless and so trim_one() commits it without looking at its reference count. So a synchronous umount pulls out a busy mount that is in use somewhere else and marks it MNT_SYNC_UMOUNT while it has users: umount2(/tmp/plshrink/p1, 0) = 0 errno 0 cwd is (unreachable) The same two umounts requested one after the other fail with EBUSY. This needs a mount with MNT_SHRINKABLE, i.e., automounted submounts of NFS, AFS and CIFS or the tracefs mount below debugfs. Bind mounts inherit the flag. Check each mount right before it is unmounted under the same hold of namespace_sem and mount_lock as the umount itself. Make sure that the algorithm stays linear.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: check a recursive bind mount for mount namespace loops do_loopback() refuses to bind mount the file of a mount namespace that is as old as the caller's or older because a mount namespace that holds a mount of its own file, or of an ancestor's, would cause a cycle. But it only checks the dentry the bind mount starts from. With MS_REC everything below it is copied and do_loopback() passes CL_COPY_MNT_NS_FILE so mount namespace files below the source are copied. That's fine for a source in the caller's own mount namespace. Every mount namespace file in there passed the same check when it was bind-mounted. It isn't fine for a source in another mount namespace. may_copy_tree() accepts a bind mount of any nsfs or pidfs file no matter what mount namespace it lives in so that /proc/<pid>/ns/<ns> and pidfds can be bind mounted. If we create a bind-mount stack of mount namespaces file descriptors the kernel will copy them irrespective of their ancestoral relationship to the mount namespace in question: 151 146 0:7 net:[4026531833] /tmp/nrc/y rw - nsfs nsfs rw 152 151 0:7 mnt:[4026532293] /tmp/nrc/y rw - nsfs nsfs rw Mount 152 is a mount of the mount namespace file of mount namespace 4026532293 inside mount namespace 4026532293. That namespace and every mount in it are leaked... All it takes is a process in an older mount namespace that stacks the file and repeating it leaks without limit: after control: Shmem: 380 kB child: mount(/proc/self/fd/6, MS_BIND|MS_REC): ok after cycle: Shmem: 65916 kB do_move_mount() runs check_for_nsfs_mounts() over a detached tree for exactly that reason. Do the same for the copy before it is grafted.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: keep covered mounts covered in OPEN_TREE_NAMESPACE open_tree(OPEN_TREE_NAMESPACE) creates a new mount namespace from a copy of the tree at the given path. The caller needs to be privileged over its current user namespace because the new mount namespace will be owned by it. It doesn't need to be privileged over the mount namespace the tree is copied from. An unprivileged user just needs to create a user namespace first. The copy follows the rules of a detached bind mount. Without AT_RECURSIVE only the mount itself is cloned and its children are left out. With AT_RECURSIVE unbindable mounts are skipped. Both is fine when the caller has privileges over the source mount namespace because it could unmount those mounts anyway. Not so for an unprivileged user in that mount namespace. For them a mount namespace copy is what unshare() does. copy_mnt_ns() copies everything including unbindable mounts and lock_mnt_tree() makes sure nothing can be unmounted in the copy. Nothing that was covered gets revealed. create_new_namespace() only does the locking and so the covered content is right there in the new mount namespace: # mount -t tmpfs none /tmp/otn; mkdir /tmp/otn/covered # echo hidden > /tmp/otn/covered/under.txt # mount -t tmpfs none /tmp/otn/covered uid 1000, after unshare(CLONE_NEWUSER): fd = open_tree(AT_FDCWD, "/tmp/otn", OPEN_TREE_NAMESPACE); setns(fd, CLONE_NEWNS); open("/covered/under.txt", O_RDONLY) -> "hidden" Covering paths with mounts is how container runtimes mask parts of /proc and /sys and how admins hide things. When the caller's user namespace doesn't own the source mount namespace copy the way unshare() does and refuse a non-recursive copy of a mount that has anything mounted below the requested directory and copy unbindable mounts in a recursive copy. A caller that is privileged over the source mount namespace sees no change.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: look at the topmost mount for a mount namespace file attach_recursive_mnt() preallocates a mountpoint for the root of the topmost mount of the source so that a mount that already sits at the destination can be put on top of the source or on top of one of its propagated copies. The copies are made without CL_COPY_MNT_NS_FILE so their chain of overmounts is shorter than the source's. It ends below the first bind mount of a mount namespace file. So while the loop walks up the chain of the source it remembers the mountpoint of that mount in "shorter" and the existing mount is put there for the copies. But the loop ends at the topmost mount without ever looking at it. If the topmost mount is the only mount namespace file in the chain "shorter" stays NULL and mnt_change_mountpoint() attaches the existing mount of the copy below the root of the mount namespace file. That's a dentry of nsfs. No path walk in the copy's mount namespace ever gets there. The mount is gone until its parent goes and the copy that took its place can't be unmounted synchronously because it has a child: S bind mount of a file N bind mount of a mount namespace file on top of S A shared mount, B a slave of A, Q mounted on B/file move_mount(S, A/file) cat B/file -> the content of S instead of Q umount B/file -> EBUSY Look at every mount of the chain including the topmost one.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: remove the fsnotify marks of a mount namespace in process context mnt_ns_release() drops the last passive reference of a mount namespace and removes its fanotify marks via fsnotify_mntns_delete(). That takes the mutex of every group with a mark on the namespace and the spinlock of the connector. Fine from process context. But mnt_ns_tree_remove() hands the reference the namespace was allocated with to call_rcu() and so the marks are removed from the RCU softirq: BUG: sleeping function called from invalid context at kernel/locking/mutex.c:623 in_atomic(): 1, irqs_disabled(): 0, non_block: 0, pid: 0, name: swapper/3 __mutex_lock+0x113/0x24b0 fsnotify_destroy_marks+0x11b/0x3d0 mnt_ns_release_rcu+0x57/0xa0 rcu_core+0x6b6/0x1e40 and lockdep complains about the connector lock being taken from softirq context. A fanotify group with a FAN_MARK_MNTNS mark on the mount namespace of another task is all that's needed. Root in a user namespace can do that for a mount namespace it owns. The task exits and the namespace is freed with the mark still on it. Remove the marks in free_mnt_ns() before the namespace is handed to RCU. That runs in process context once the last active reference is gone. A mark is added through a file descriptor to the namespace which holds an active reference so no mark can show up after that.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: check the mounts before reading their parents in pivot_root() path_pivot_root() reads the parents of new_root and of the caller's root and checks whether they are shared before it checks that either mount is in the caller's mount namespace. Only namespace_sem is held. That's fine for a mount that is in the caller's mount namespace. But new_root can be a file descriptor to a mount that has been unmounted and that only the file descriptor keeps alive. If that mount stayed attached to its parent when it was unmounted nothing pins the parent for it. The final mntput() of the parent unhooks the children under mount_lock alone and frees the parent afterwards: pivot_root() close(fd), last ref on the parent ------------ --------------------------------- ex_parent = new_mnt->mnt_parent mntput_no_expire_slowpath() __umount_mnt(new_mnt) cleanup_mnt() call_rcu() IS_MNT_SHARED(ex_parent) BUG: KASAN: slab-use-after-free in path_pivot_root+0xf1a/0x1840 Read of size 4 at addr ffff8881047382f8 by task pivot_widen/157 path_pivot_root+0xf1a/0x1840 __x64_sys_pivot_root+0x165/0x190 The same goes for the caller's root via chroot(). Both outcomes of the check end in EINVAL so nothing but the read itself goes wrong. It's the same thing commit bb4609405752 ("statmount: read the parent of an unmounted mount under mount_lock") fixed for statmount(). Check that both mounts are in the caller's namespace before their parents are read. Every path returns EINVAL either way.

Linux

In the Linux kernel, the following vulnerability has been resolved: dcache: don't put a mountpoint on a dentry that's being removed rmdir(), unlink() and rename() call dont_mount() on the victim and then detach_mounts() with the victim's inode locked. do_lock_mount() takes the inode lock of the mountpoint and checks cant_mount() so a mount can't show up after detach_mounts(). But attach_recursive_mnt() makes a second mountpoint for the root of the source mount so that the mounts already located at the destination can be put on top of it. No inode is locked for that one and d_set_mounted() only refuses a dentry that is unlinked. Between detach_mounts() and d_delete() the victim is still hashed: rmrace: b passed dont_mount() and detach_mounts(), sleeping T2: move_mount(S, /tmp/plcant/x, BENEATH) = 0 errno 0 () T1: rmdir(/tmp/plcant/d/b) = 0 errno 0 109 107 0:61 /d/b//deleted /tmp/plcant/x rw,relatime - tmpfs tmpfs rw 108 109 0:63 / /tmp/plcant/x rw,relatime - tmpfs T rw A bind mount of the directory that's being removed is moved beneath an existing mount while the rmdir() is located between the two calls. The mount ends up on the removed directory and nothing will ever detach it. Check cant_mount() in d_set_mounted() as well. dont_mount() raises the flag under d_lock before detach_mounts() runs so either the mountpoint is set first and detach_mounts() finds the mount or the flag is seen and the mount is refused.

Linux

In the Linux kernel, the following vulnerability has been resolved: unshare: don't drop active namespace references that were never taken Active references on the namespaces of an nsproxy are taken when the nsproxy is installed into a task in switch_task_namespaces() and copy_namespaces() and dropped again by put_nsproxy() through deactivate_nsproxy(). But ksys_unshare() calls put_nsproxy() on an nsproxy that was never installed when set_cred_ucounts() fails. The new namespaces of that nsproxy go from zero to minus one and the namespaces shared with the caller lose a reference that belongs to the nsproxy the caller keeps using: WARNING: kernel/nscommon.c:171 at __ns_ref_active_put+0x1cd/0x230 nsproxy_ns_active_put deactivate_nsproxy ksys_unshare Afterwards the namespaces the caller lives in aren't listed by listns() anymore. set_cred_ucounts() only fails when alloc_ucounts() can't allocate and it only allocates when the real uid of the caller differs from its effective uid. Commit cefd55bd2159 ("nsproxy: fix free_nsproxy() and simplify create_new_namespaces()") separated the two cases on purpose. nsproxy_free() frees an nsproxy that was prepared but never installed and that's what a failed setns() uses in put_nsset(). Export it and use it for a failed unshare() as well.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: don't let a pseudo dentry become the root of a mount d_alloc_pseudo() hands out dentries for pipes, sockets and other files that are never anyone's child or parent. They carry DCACHE_NORCU and dentry_free() frees them right away because no lockless path walk can ever reach them. That holds as long as such a dentry isn't the root of a mount. But bind mounting /proc/self/fd/<fd> of such a file does exactly that. Most callers of alloc_file_pseudo() put their files on kernel-internal mounts and may_copy_tree() refuses those. bpf_token_create() doesn't. It places the token file on the bpffs mount the caller handed it and that mount is in the caller's mount namespace so the clone goes through: mount --bind /proc/self/fd/<token> <file> open_tree(tokfd, "", AT_EMPTY_PATH | OPEN_TREE_CLONE) + move_mount() __follow_mount_rcu() then loads ->mnt_root of that mount, reads d_seq and d_flags of the dentry and only then checks mount_lock. The dentry can be gone by then. The final mntput() of an unmounted parent unhooks a child that stayed attached to it under mount_lock alone. When the child's holder does the final put right after that the root is dput() and freed immediately while a walker that found the mount hashed still looks at it. Refuse to clone a mount with a DCACHE_NORCU dentry as its root. Nothing sensible can be done with a bind mount of a bpf token anyway.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: unhash a dentry before detaching the mounts on it unlink(), rmdir() and rename() remove the entry from the filesystem, call detach_mounts() on the entry with the inode locked and then call d_delete() once the inode is unlocked. The thing is that between detach_mounts() and d_delete() the dentry is still hashed and positive. But after detach_mounts() nothing covers the dentry anymore. A lookup that finds this dentry in the dcache can uncover what the mounts hid. That's a problem when unlinking files or directories that are mountpoints in other mount namespaces. Everybody who had the underlying entry covered can race the detach_mount() call until d_delete() has run. The race window isn't all that small. It encompasses namespace_unlock() with a full synchronize_rcu_expedited() grace period and the inode_unlock() of the entry. In my experiments three walkers that kept trying read an overmounted file in 11938 of the 16140 unlinks that removed its mountpoint from a bind mount of the filesystem within a minute. Fun fact, d_invalidate() has the same ordering problem but gets it right. It unhashes the dentry first and detaches the mounts afterwards. Let's do the same in __detach_mounts(): - A lookup that hasn't found the dentry yet misses it in the dcache and waits for the directory lock that the caller holds until the name is gone for good. - A lockless lookup that found it already fails the mount_lock check while the dentry still counts as a mountpoint and the d_seq check once it doesn't, and retries. This fixes the lockless path. We still need to fix the reference count lookup in a follow-up patch. So d_drop() the dentry. The dentry stays positive and held but can't be found anymore. Then proceed with the detach and unlink. Reproducer: The reads were counted with the program below, three walkers for 60 s in a VM with 4 CPUs. The race window is stretched with the debug patch pasted here. fs.detach_race_walk_us sleeps in lookup_fast() once it found a mountpoint dentry and before its mounts are crossed. fs.detach_race_unlink_us sleeps in vfs_unlink() between detach_mounts() and d_delete(). // SPDX-License-Identifier: GPL-2.0 /* * unlink_covered: unlink a file that is a mountpoint in a detached copy of * its mount, against lookups that open it through that copy. * * A is a tmpfs with the file f1 (content SECRET_F) and the plain file * MARK_A. A' is an open_tree(OPEN_TREE_CLONE) copy of A with MARK_A bound * on A'/f1, held through an O_PATH fd on its root once the tree fd is * closed. Walkers open f1 through that fd while the driver unlinks A/f1, * where nothing is mounted on it. A walker may read MARK_A or get ENOENT. * A read of SECRET_F is a hit: the name was found after its mount was * gone. The lockless walks use openat2(RESOLVE_CACHED). * * usage: unlink_covered [-t seconds] [-w walkers] */ #ifndef _GNU_SOURCE #define _GNU_SOURCE #endif #include <errno.h> #include <fcntl.h> #include <pthread.h> #include <stdatomic.h> #include <stdio.h> #include <stdlib.h> #include <string.h> #include <unistd.h> #include <sys/mount.h> #include <sys/stat.h> #include <sys/syscall.h> #include <linux/openat2.h> #ifndef __NR_open_tree #define __NR_open_tree 428 #endif #ifndef __NR_move_mount #define __NR_move_mount 429 #endif #ifndef __NR_openat2 #define __NR_openat2 437 #endif #ifndef OPEN_TREE_CLONE #define OPEN_TREE_CLONE 1 #endif #ifndef OPEN_TREE_CLOEXEC #define OPEN_TREE_CLOEXEC O_CLOEXEC #endif #ifndef MOVE_MOUNT_F_EMPTY_PATH #define MOVE_MOUNT_F_EMPTY_PATH 0x00000004 #endif #define WORK "/tmp/uc" #define ADIR WORK "/A" static int duration = 60, nwalkers = 3; static atomic_int stop, writer_waiting; static pthread_rwlock_t cur_lock = PTHREAD_RWLOCK_INITIALIZER; static int cur_fd = -1; /* the root of A', -1 wh ---truncated---

Linux

In the Linux kernel, the following vulnerability has been resolved: namei: don't reveal overmounted entries in refwalk In rcuwalk the dentry is validated before it is accepted. step_into() step_into() rechecks d_seq and __follow_mount_rcu() rechecks mount_lock. An entry that gets unlinked in between causes the lookup to retry and miss. A refwalk doesn't do this. lookup_fast() takes a reference on the hashed dentry and simply accepts it. So an unlink that happens after the reference was taken isn't seen by refwalk. That's fine for a simple file. We just happened to open it before it was unlinked, no problem. For a mountpoint and specifically a locked mountpoint it very much isn't. The unlink detaches all mounts and then removes the name. Any refwalk that hasn't traversed the mounts yet simply reveals the underlying entry. It's a very narrow window but it can be hit: reads of a covered file in 60 s, 3 walkers, ~15000 unlinks no widening 27 that step + 200 us 3741 See the appended patch for a more reliable reproducer. So check the entry after step_into(). unlink(), rmdir() and rename() mark the dentry with dont_mount() before they detach the mounts and remove the dentry. So a dentry that is marked with DCACHE_CANT_MOUNT and is unhashed by the time its mounts were looked at is a name that was unlinked under the refwalk. DCACHE_CANT_MOUNT is read after the mounts were looked up and it is set before they are detached by detach_mounts(). A refwalk that missed the mounts will see DCACHE_CANT_MOUNT. Plain d_unlinked() is fine. A rename takes the dentry off its hash chain but ___d_drop() leaves d_hash.pprev set. So only __d_drop() and a rename over the dentry unhash it. The only move that flips IS_ROOT splices in a disconnected alias. That doesn't have DCACHE_CANT_MOUNT set. Add the to step_into_slowpath(). ".." and LOOKUP_DOWN may legitimately land on an unhashed directory. A refwalk that crossed onto a mount has path.mnt different from nd->path.mnt. A dentry that a filesystem dropped on its own (d_invalidate(), d_drop()) doesn't have the flag set and is treated as before. // SPDX-License-Identifier: GPL-2.0 /* * unlink_covered: unlink a file that is a mountpoint in a detached copy of * its mount, against walkers that open it through that copy. * * A is a tmpfs with the file f1 (content SECRET_F) and the plain file * MARK_A. A' is an open_tree(OPEN_TREE_CLONE) copy of A with MARK_A bound * on A'/f1, held through an O_PATH fd on its root once the tree fd is * closed. Walkers open f1 through that fd while the driver unlinks A/f1, * where nothing is mounted on it. A walker may read MARK_A or get ENOENT. * A read of SECRET_F is a hit: the name was found after its mount was * gone. The lockless walks use openat2(RESOLVE_CACHED). * * usage: unlink_covered [-t seconds] [-w walkers] */ #ifndef _GNU_SOURCE #define _GNU_SOURCE #endif #include <errno.h> #include <fcntl.h> #include <pthread.h> #include <stdatomic.h> #include <stdio.h> #include <stdlib.h> #include <string.h> #include <unistd.h> #include <sys/mount.h> #include <sys/stat.h> #include <sys/syscall.h> #include <linux/openat2.h> #ifndef __NR_open_tree #define __NR_open_tree 428 #endif #ifndef __NR_move_mount #define __NR_move_mount 429 #endif #ifndef __NR_openat2 #define __NR_openat2 437 #endif #ifndef OPEN_TREE_CLONE #define OPEN_TREE_CLONE 1 #endif #ifndef OPEN_TREE_CLOEXEC #define OPEN_TREE_CLOEXEC O_CLOEXEC #endif #ifndef MOVE_MOUNT_F_EMPTY_PATH #define MOVE_MOUNT_F_EMPTY_PATH 0x00000004 #endif #define WORK "/tmp/uc" #define ADIR WORK "/A" static int duration = 60, nwalkers = 3; static atomic_int stop, writer_waiting; static pthread_rwlock_t cur_lock = PTHREAD_RWLOCK_INITIALIZER; static int cur_fd = -1; /* the root of A', -1 while there is none */ static atomic_long n_unlink, n_walk, n_mark, n_enoent, n_othe ---truncated---

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: refuse an automount below a mount that is in no namespace It's possible to add autmounts even when the parent mount isn't in the mount namespace of the caller. The only requirement we have is that the parent mount namespace must not be NULL, i.e., unmounted. Problem is that clone_private_mount() has MNT_NS_INTERNAL which makes that trivially true. So that passes the test and attach_recursive_mnt() accepts that as a mount point and funny enough, count_mounts() dereferences MNT_NS_INTERNAL. The problem is it is an error pointer... So we can reach this in userspace via fanotify. A filesystem mark on the lower filesystem of an overlay reports paths on the layer clone and reading the event hands out a descriptor on it. For example with debugfs as the lower layer it goes kaboom: openat(evfd, "tracing", O_DIRECTORY) Oops: general protection fault KASAN: null-ptr-deref in range [0x1d0-0x1d7] RIP: 0010:count_mounts+0x35/0x200 attach_recursive_mnt finish_automount And since that sleeping beauty happens under namespace_sem held for writing every mount operation on the system blocks from then on. Congrats. Use is_mounted() instead which rejects unmounted and internal mounts alike. The open fails with EINVAL just as it did before clone_private_mount() used MNT_NS_INTERNAL.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: never expire a locked mount Locked mounts are special. They protect the underlying files and directories from being revealed. do_umount() refuses to unmount locked mounts but shrink_submounts() doesn't. A shrinkable mount can become locked once the owner of a user namespace puts it beneath a locked mount with MOVE_MOUNT_BENEATH. The locked property now moves to the mount at the bottom making it possible to unmount the top mount. So umount() of an unlocked ancestor now expires the bottom mount first and the covered directory is revealed: move_mount(c -> x/hidden, MOVE_MOUNT_BENEATH) = 0 the cover after the lock moved: umount2(x/hidden) = 0 the holder of the lock: umount2(x/hidden) = EINVAL the root of the copy, busy: umount2(x) = EBUSY reads x/hidden/secret: "covered-by-root" So leave a locked mount alone as it dies together with its parent. A plain umount() of an unlocked mount with a locked one below it is EBUSY from now on. It's the same for any other locked child. A lazy umount still takes the whole tree. mark_mounts_for_expiry() never sees a locked mount. lock_mnt_tree() leaves a mount on an expiry list alone and nothing else puts a locked one on such a list. Add an assert for this.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: handle mount locking for automounts correctly If mounts are propagated across user namespaces, attach_recursive_mnt() locks every copy of the source mount to protect overmounts from vanishing and revealing the underlying files or directories. The user namespace is taken from the caller's mount namespaces since this is where the mounts end up. Except, that's not always true. Automounts may legitimately get popped in by tasks located in a different mount and user namespace during path lookup. Then check doesn't make sense at that point. The copy in the namespace of the parent - which may be the host's - is now locked and the host cannot change the flags of its own mounts anymore. Here's the reproducer: - host has debugfs mounted nosuid,nodev,noexec and shared - tracefs automount below it is not yet active - hand a child process a directory descriptor on that mount - child process enters auser and mount namespace - child process' namespace now holds a locked copy of the host's debugfs mount that receives propagation from it - child process stats "tracing/." through the descriptor - lookup runs on the host's mount so the automount lands below the host's mount - propagation puts a copy below the child's copy - both try to clear the flags on the mount they got, with a bind remount: host, on its own automount: MS_REMOUNT|MS_BIND = EPERM child, on the copy in its namespace: MS_REMOUNT|MS_BIND = 0 So the lock landed on the host's mount instead of the child's copy. Congrats. So we need to compare with the owner of the namespace the mount actually gets mounted on. For all regular cases that is the caller's mount namespace and so nothing changes. Detached trees in anonymous mount namespaces by be handed over via SCM_RIGHTS or inherited in other ways on purpose so the attaching task's mount namespace is authoritative, not the creator of the detached tree.

Linux

In the Linux kernel, the following vulnerability has been resolved: namespace: keep the lock on a mount that a propagated copy is moved beneath attach_recursive_mnt() transfers MNT_LOCKED from the top mount to the mount that is moved beneath it with MOVE_MOUNT_BENEATH. This allows the owner of a user namespace to replace its locked /proc or its root. The mount beneath takes on the job of covering the underlying mountpoint allowing the top mount to be unmounted. Consider two mount namespaces: (H) The host H has a shared mount P with a secret in P/d (Z) Z is a user namespace made from M. Its copy of P receives propagation from the host's P and its copy of M's cover on P/d is locked Z's owner must not get to see P/d. Now (H) mounts X on P/d. This propagates into (Z). The copy of X lands beneath (Z)'s locked cover. The locked property is transfered from (Z)'s cover to the copy of X propagated beneath it. The cover is now unlocked. Now (H) unmounts X again. The copy of X in (Z) gets unmounted and the covering mount is left unlocked on top of P/d. (Z) can now unmount it: Z: umount2(P/d) = EINVAL /* the cover is locked */ H: mount X on P/d, umount X /* both propagate into Z /* Z: umount2(P/d) = 0 /* the cover is now unlocked */ Z: read P/d/secret = "covered-by-root" /* secret revealed */ So only transfer the locked property to the mount beneath for mounts the caller has placed. A propagated copy that lands beneath a locked mount is locked as well so that the mount at the bottom of the stack carries a lock the way every check expects. The mount on top of it remains locked to ensure that it keeps covering even if the propagated mount is unmounted again.

Linux

In the Linux kernel, the following vulnerability has been resolved: xfs: stop exchanging reflink flags when exchanging mappings When an exchange covers both files from offset 0 to i_disk_size, xmi_can_exchange_reflink_flags() assumes every mapping moved and hands the reflink flag over to the other file. That assumption is wrong. A file can still own shared blocks past i_disk_size, e.g. after EXCHANGE_RANGE_TO_EOF sets a smaller size without unmapping anything. The flag is cleared anyway, and the next write to such a block happens in place instead of CoW, corrupting the other file that shares it. Commit a23eca88448e ("xfs: fix exchange-range reflink flag clearing issue with INO1_WRITTEN") fixed one instance of this. Rather than add another special case, stop moving the flags. xfs_exchmaps_ensure_reflink() already sets the flag on both inodes before the exchange, so both just keep it. Nothing sets CLEAR_INO{1,2}_REFLINK anymore. Intents logged by older kernels may still carry them; recovery already masks them off, so they stay in the log format mask and are ignored. XFS_IOC_SWAPEXT has the same problem: it moves the reflink flag from one inode to the other along with the extents. Just removing that is not enough here: the exchange path flags both inodes up front in xfs_exchmaps_ensure_reflink(), but swapext has no such step, so the inode receiving the shared extents would be left without the flag. Set it on both inodes instead, and make sure both have a CoW fork, since an inode with the reflink flag is expected to have one. The flag now stays until FALLOC_FL_UNSHARE_RANGE, xfs_scrub, or a truncate to zero removes it.

Linux

In the Linux kernel, the following vulnerability has been resolved: xfs: fix exchange-range-to-eof file size exchange Norbert Szetei posted a patch containing an incomplete description of a bug in XFS_IOC_EXCHANGE_RANGE's TO_EOF flag. The bug tricks exchange-range into modifying a file so that it has shared blocks starting beyond EOF, which is never allowed because files never have written data beyond EOF, and you can only share written data blocks. This is key to all the incorrect behavior that follows. Eventually I got from Norbert a description of what's going wrong: > Fair. Here it is with physical block numbers, 4k blocks. peer is the file > that gets rewritten and file1 is the one that ends up with the flag > cleared. > > Start, nothing shared: > > peer size 4096 -> [100] > file1 size 12288 -> [200] [201] [202] > donor1 size 4096 -> [300] > donor2 size 4096 -> [400] > > 1. FICLONERANGE peer's block into file1's middle slot: > > peer size 4096 -> [100] > file1 size 12288 -> [200] [100] [202] reflink flag set > ^^^^^ both files own block 100 > > 2. XFS_IOC_EXCHANGE_RANGE file1 against donor1, file1_offset 0, > file2_offset 8192, length 0, TO_EOF. One block is exchanged, and the sizes > are exchanged with it: > > file1 size 4096 -> [200] [100] [300] reflink flag set > ^^^^^^^^^^^ still owned, now past the size > donor1 size 12288 -> [202] > > file1 says it is one block long while it still owns three. Nothing was > unmapped. file2_offset is block 2, so this exchange cannot clear a flag. > > 3. XFS_IOC_EXCHANGE_RANGE file1 against donor2, both offsets 0, length 4096. > By i_disk_size this is two one-block files exchanging everything, so > xmi_can_exchange_reflink_flags() agrees and the flag is cleared: > > file1 size 4096 -> [400] [100] [300] reflink flag CLEARED > donor2 size 4096 -> [200] > > Block 100 is still shared with peer. > > 4. pwrite(file1, 4096, 4096), which is file1's second slot, block 100. > xfs_is_cow_inode(file1) is false, so no CoW: > > peer size 4096 -> [100] same block, new contents > > Block 100 is the only one that matters here. PoC below, it reads a file > called peer in the current directory. The filesystem needs reflink and > exchange-range, which mkfs.xfs 7.x turns on by default. So yes, TO_EOF is screwing up the file size even if it's actually exchanging the data block correctly. The following is the intended behavior of XFS_EXCHANGE_RANGE_TO_EOF, as described by its manpage: "Ignore the length parameter. All bytes in file1_fd from file1_offset to EOF are moved to file2_fd, and file2's size is set to (file2_offset+(file1_length-file1_offset)). Meanwhile, all bytes in file2 from file2_offset to EOF are moved to file1 and file1's size is set to (file1_offset+(file2_length-file2_offset))." In other words, if we only exchange a portion of the file data, then we should only exchange the portion of the file size that relates to the exchanged data. Unfortunately, the code swaps the ondisk file size and the incore file sizes, which is incorrect if the caller passes in a nonzero offset. Having confused the file sizes, the PoC uses a second exchange-range call on what looks like a non-reflinked single-block file. Because the file size is set incorrectly, exchange-range thinks it's exchanging the full contents of two files and clears the reflink flag on the broken file. That enables an extending write of the broken file to rewrite the shared block that's just past EOF. This has become known colloquially as refluxfs. Once the file sizes are set correctly, the second exchange-range no longer thinks that it's doing a full-contents swap, so it won't clear the reflink flag on either of its file arguments. Norbert followed up to an earlier version of this patch with a second complaint about us not handling the i_disk_size state correctly: > The offsets are bounded against the incore i_size but subtracted from > ---truncated---

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/bridge: Fix unlocked list_del in drm_bridge_add() When re-adding a bridge that was previously removed, drm_bridge_add() drops it from bridge_lingering_list without holding bridge_lock. Both bridge_list and bridge_lingering_list are protected by bridge_lock, as they are concurrently modified by drm_bridge_remove() and __drm_bridge_free(), and walked by the debugfs 'bridges' file. Running the list_empty() test and the list_del_init() outside of the lock may therefore corrupt either list. Perform both under bridge_lock.

Linux

In the Linux kernel, the following vulnerability has been resolved: cgroup/cpuset: Don't access cpuset_cgrp_subsys.root in is_in_v2_mode() After seeing the patch [1] to guard is_in_v2_mode() with RCU, it makes me realize that is_in_v2_mode() may be called in a context where a new cgroup filesystem is being rebound with stale cpuset_cgrp_subsys.root pointer. Avoid this potential UaF situation by adding a new cpuset_v2_mode flag which is set when the cpuset_v2_mode mount option is used. This flag is written into only when cpuset_bind() is being called with a stable cpuset_cgrp_subsys.root value. The is_in_v2_mode() helper is modified to read the new cpuset_v2_mode flag instead of accessing cpuset_cgrp_subsys.root directly. [1] https://lore.kernel.org/lkml/20260929084124.626693-2-arighi@nvidia.com

Linux

In the Linux kernel, the following vulnerability has been resolved: gpio: exar: initialize the ID before registering its cleanup devm_add_action_or_reset() invokes the cleanup action immediately when registering the action fails. gpio_exar_probe() currently stores the allocated ID in exar_gpio only after registering the action, so that failure makes exar_devm_ida_free() release ID 0. If the allocated ID is nonzero, this leaks that ID and can free ID 0 belonging to another device. Store the ID before registering the cleanup action so both the immediate failure path and normal device teardown release the correct allocation.

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/vc4: Free the BO cache size list on teardown vc4_bo_cache_destroy() purges the cached BOs but never frees the size list, leaking it on every unbind.

Linux

In the Linux kernel, the following vulnerability has been resolved: net/sched: em_text: reject unterminated algo before autoload This is a follow-up to commit 9c572a83037a ("net/sched: fix potential stack infoleak in em_text_dump()"), which zero-initialised the dump-side struct. Sashiko pointed at a pre-existing bug that exposed the change-side input-validation gap open. em_text_change() casts the raw netlink attribute payload to struct tcf_em_text and passes conf->algo, a 16-byte array with no enforced NUL terminator, to textsearch_prepare(). On the TS_AUTOLOAD retry textsearch_prepare() calls request_module("ts_%s", algo); vsnprintf() walks %s until it finds a NUL, so an algo[] with no NUL reads past the array into the adjacent struct fields and payload and feeds those bytes to the usermode helper command line. This is an out-of-bounds read and a kernel memory disclosure. Reject the attribute when no NUL exists within TC_EM_TEXT_ALGOSIZ bytes, matching the check xt_string has had since 3ab720881b6e. Conditions to recreate the bug: - CAP_NET_ADMIN on the target netns (namespace-local via unshare -Urn is enough) - install clsact on an interface, then add a basic filter with one text ematch whose algo[] is 16 bytes with no NUL (e.g. "A"*16) and pattern_len at least 1 - the add fails with -ENOENT and the kernel invokes modprobe with a module name made of "ts_" plus the 16 bytes and the adjacent payload

Linux

In the Linux kernel, the following vulnerability has been resolved: amt: send the relay's General Query directly from the receive path An skb queued in a qdisc can outlive the tunnel it references through a raw pointer in skb->cb. For example, with igmp_qrv set to 1 on the relay a tunnel lives for 135s, so a netem delay of 180s on the amt device outlives it; when the tunnel expires and is freed, the subsequent dequeue triggers a use-after-free in amt_dev_xmit(). BUG: KASAN: slab-use-after-free in amt_dev_xmit+0x2763/0x2e20 Call Trace: amt_dev_xmit+0x2763/0x2e20 [drivers/net/amt.c:1262] dev_hard_start_xmit+0x22f/0x620 sch_direct_xmit+0x12e/0xac0 netem_dequeue+0x333/0xc50 net_tx_action+0x35c/0xa60 amt_send_igmp_gq() and amt_send_mld_gq() are only called from amt_request_handler(), inside the rcu_read_lock_bh() section of amt_rcv(). amt_request_handler() already has the tunnel the query is for: it found or created it inside that section. Queuing the query with dev_queue_xmit() only leads back into amt_dev_xmit(), which strips the Ethernet header and calls amt_send_membership_query() for that tunnel. Make that call directly from the two senders instead, the same way amt_send_advertisement() transmits from the receive path. The query never waits in a qdisc, the tunnel is only dereferenced inside the RCU section that found or created it, and nothing is stored in skb->cb, so no lookup or refcount is needed. Remove the query branch of amt_dev_xmit(), amt_skb_cb() and struct amt_skb_cb, which have no users left. Behaviour changes: - The relay's own General Queries no longer pass through the amt device's egress path: its qdisc, tc egress (clsact/tcx), the netfilter egress hook and packet taps. They are still visible as UDP on the underlay. - A query that is sent successfully is no longer counted as tx_dropped. The old query branch left through the unlock label, which counted every sent query as dropped. - A query that reaches amt_dev_xmit() on a relay from elsewhere, such as a userspace querier, is now dropped at the IGMP/MLD type switch. Before, it trusted whatever skb->cb held, and a NULL tunnel hit the WARN_ON(1).

Linux

In the Linux kernel, the following vulnerability has been resolved: cxgb4/ch_ktls: disable softirqs around tid_list erases chcr_ktls_cpl_act_open_rpl() inserts into tid_list from the receive path using xa_insert_bh(). The array is shared by the adapter's connections and initialized with XA_FLAGS_LOCK_BH. However, chcr_ktls_dev_del() and the chcr_ktls_dev_add() error path use xa_erase(), which takes the XArray lock without disabling softirqs. When cleanup runs with softirqs enabled, a receive softirq for another connection on the same CPU can interrupt the erase and spin on the lock held by the interrupted task. Use xa_erase_bh() at both sites to match the insertion path.

Linux

In the Linux kernel, the following vulnerability has been resolved: bonding: fix the program leak when XDP is replaced We need to release the reference on the old BPF prog, whether we are going to no-prog, or replacing the prog with another. Spotted by AI while running tests for the propagation series (leaked programs were accumulating). Selftest will come in the net-next series to avoid conflicts.

Linux

In the Linux kernel, the following vulnerability has been resolved: ice: Restore Ordered MMIO Writes for Tx Doorbells The DQL accounting state which tracks the number of bytes queued for the NIC to transmit, namely dql->num_queued, is updated by the ICE driver when it invokes __netdev_tx_sent_queue(). Subsequently the driver updates the hardware queue tail allowing the NIC to begin processing the work queue. The queue accounting update should be ordered before the NIC begins transmitting descriptors. The transmit path currently updates the hardware doorbell using writel_relaxed(), which (on arm64) does not provide the same ordering guarantees between writes to normal memory and writes to MMIO registers that writel() does. This introduces a potential race where the NIC begins transmitting descriptors before the dql->num_queued update is globally visible. If the NIC finishes before the update is observed by dql_completed(), it detects an invalid state where more bytes have been completed than have been queued. When this happens the following BUG_ON is triggered. BUG_ON(count > num_queued - dql->num_completed); This has been observed and manifests as the following crash (note that the trace has been trimmed) in an environment with sustained network load that uses an Intel Corporation Ethernet Controller E810-XXV for SFP (rev 02) on an arm64 machine. Replacing writel_relaxed() with writel() in a test kernel eliminated the crash in the user's workload, which previously reproduced the issue reliably. kernel BUG at lib/dynamic_queue_limits.c:99 Internal error: Oops - BUG: 00000000f2000800 [#1] SMP pc : dql_completed+0x268/0x2a0 lr : ice_clean_tx_irq+0x1d4/0x620 [ice] Call trace: dql_completed+0x268/0x2a0 (P) ice_napi_poll+0x94/0x520 [ice] __napi_poll+0x48/0x3f0 net_rx_action+0x194/0x420 This restores the behaviour prior to commit ccde82e90946 ("ice: add E830 Earliest TxTime First Offload support"), which changed the notification mechanism from writel() to writel_relaxed().

Linux

In the Linux kernel, the following vulnerability has been resolved: nvmet: use the local P2P device on metadata allocation failure nvmet_req_alloc_p2pmem_sgls() assigns req->p2p_dev only after the data and metadata SGL allocations succeed. If the data allocation succeeds but the metadata allocation fails, out_free_sg passes req->p2p_dev to pci_p2pmem_free_sgl() before that assignment has taken place. Pass the local p2p_dev argument instead, so cleanup uses the same PCI device as the successful data allocation.

Linux

In the Linux kernel, the following vulnerability has been resolved: btrfs: always return -EIOCBQUEUED after btrfs_uring_read_extent_endio If all bios finish before btrfs_encoded_read_regular_fill_pages() returns, it calls btrfs_uring_read_extent_endio() and previously returned the I/O status. A negative errno then made btrfs_uring_read_extent() unlock and free while btrfs_uring_read_finished() did the same again. Return -EIOCBQUEUED so only the deferred path cleans up.

Linux

In the Linux kernel, the following vulnerability has been resolved: btrfs: free iov when btrfs_uring_read_extent() fails After btrfs_uring_read_extent(), the caller always jumped to out_acct. That skips kfree(data->iov), which is only correct for -EIOCBQUEUED where the deferred path owns the iov. On failure, fall through to out_free instead.

Linux

In the Linux kernel, the following vulnerability has been resolved: btrfs: unlock inode and extent in caller when io_uring read extent fails btrfs_uring_read_extent() runs only after btrfs_encoded_read() has taken the inode shared lock and the extent lock. On failure it used to unlock in out_fail, and a pages-array allocation failure returned -ENOMEM without unlocking at all. Unlock in the caller instead on all failure returns, matching the copy_to_user() error path. The deferred -EIOCBQUEUED path still unlocks in btrfs_uring_read_finished().

Linux

In the Linux kernel, the following vulnerability has been resolved: btrfs: don't stash io_uring encoded data across -EAGAIN Returning -EAGAIN while leaving btrfs_uring_encoded_data in the cmd PDU leaks if the request is cancelled or the ring exits before reissue. io_uring does not free driver PDU allocations on cleanup. Write: io_queue_sqe() always issues with IO_URING_F_NONBLOCK first, so return -EAGAIN before allocating and free data on every exit. Read: free on nowait -EAGAIN too; only -EIOCBQUEUED keeps the allocation for btrfs_uring_read_finished().

Linux

In the Linux kernel, the following vulnerability has been resolved: btrfs: fix xattr replace when multiple xattrs are packed in the same item If we have a btrfs_dir_item item that packs multiple xattrs and then we replace the value of one of them (with the setxattr(2) family of syscalls) with another value of a different size, we end up not having a fully initialized btrfs_dir_item, resulting in a corruption that the tree checker will detect at extent buffer writeback time. This is because in btrfs_setxattr() when we find a btrfs_dir_item with multiple xattrs (due to the crc32c hash of their name being the same) we delete one of the xattr items (btrfs_dir_item) and then insert a new one, but the deletion and insertion results in shifting existing data in the leaf and therefore when the new value of a xattr has a different size, the new btrfs_dir_item is placed in a leaf section that was not initialized and we only copy the value's data and set the value's length in the new btrfs_dir_item, without setting the name, the name's length, the key (which must be all zeroes for xattrs), flags (BTRFS_FT_XATTR) and transaction ID. The following script reproduces the issue: $ cat test.sh #!/bin/bash DEV=/dev/sdi MNT=/mnt/sdi mkfs.btrfs -f $DEV mount $DEV $MNT touch $MNT/testfile # Add two xattrs that, on btrfs, have the same hash (crc32c) for their # name and therefore are packed into the same btrfs_dir_item. setfattr -n user.foobar -v 123 $MNT/testfile setfattr -n user.WvG1c1Td -v qwerty $MNT/testfile # Verify the xattrs are present. echo "xattrs before:" getfattr --absolute-names --dump $MNT/testfile # Now replace the value of the foobar xattr with a significantly larger # value. setfattr -n user.foobar -v abcdefghijklmnopqrstuvwxyz $MNT/testfile # Check the xattrs have the expected values. echo "xattrs after:" getfattr --absolute-names --dump $MNT/testfile umount $MNT Running it: $ ./test.sh (...) xattrs before: # file: /mnt/sdi/testfile user.WvG1c1Td="qwerty" user.foobar="123" xattrs after: # file: /mnt/sdi/testfile user.WvG1c1Td="qwerty" So the "user.foobar" xattr is missing and there was a transaction abort when unmounting the fs with the following traces in dmesg: $ dmesg [869800.159271] BTRFS warning (device sdi): access to eb bytenr 30474240 len 16384 out of range start 16015 len 25964 [869800.159293] ------------[ cut here ]------------ [869800.159296] WARNING: fs/btrfs/extent_io.c:4408 at report_eb_range+0x44/0x60 [btrfs], CPU#8: getfattr/2605179 [869800.168961] Modules linked in: btrfs dm_thin_pool (...) [869800.190527] CPU: 8 UID: 0 PID: 2605179 Comm: getfattr Tainted: G W 7.3.0-rc3-btrfs-next-244+ #1 PREEMPT(full) [869800.193726] Tainted: [W]=WARN [869800.194502] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS rel-1.16.2-0-gea1b7a073390-prebuilt.qemu.org 04/01/2014 [869800.197576] RIP: 0010:report_eb_range+0x44/0x60 [btrfs] [869800.198985] Code: 48 8b 7b 18 (...) [869800.203531] RSP: 0018:ffffce4541937d58 EFLAGS: 00010246 [869800.204609] RAX: 0000000000000000 RBX: ffff8dde054b8738 RCX: 0000000000000000 [869800.206137] RDX: 0000000000000000 RSI: 0000000000000001 RDI: ffffffffc04c92a0 [869800.207645] RBP: 0000000000003e8f R08: 0000000000000000 R09: 3fffffffffefffff [869800.226175] R10: ffffce4541937a88 R11: 0000000000000003 R12: 000000000000656c [869800.227594] R13: ffff8dde14be800f R14: 000000000000656c R15: 0000000000000069 [869800.229097] FS: 00007f59023bb780(0000) GS:ffff8de5788ed000(0000) knlGS:0000000000000000 [869800.231137] CS: 0010 DS: 0000 ES: 0000 CR0: 0000000080050033 [869800.232321] CR2: 0000558dc45e6a78 CR3: 0000000765b16004 CR4: 0000000000370ef0 [869800.233783] Call Trace: [869800.234330] <TASK> [869800.234791] read_extent_buffer+0x4d/0x100 [btrfs] [869800.235898] btrfs_listxattr+0x199/0x240 [btrfs] [869800.236912] vfs_listxattr+0x51/0 ---truncated---

Linux

In the Linux kernel, the following vulnerability has been resolved: wifi: mt76: account non-AQL frames per peer rather than per link The driver increments wcid->non_aql_packets when a frame goes to the hardware queue. The completion path decrements the counter again. For an MLO peer, the hardware reports the completion on the link that it used for the transmission. This link can differ from the link that the driver queued the frame on. In that case, the increment and the decrement go to different link wcids, and the count drifts upwards until it blocks the station. Keep the count on the wcid of the primary link, like the other peer-wide state.

Linux

In the Linux kernel, the following vulnerability has been resolved: wifi: mt76: mt7915: disable rx napi when removing device Unloading the mt7915e module with a MT7916 triggers multiples WARN in __netif_napi_del_locked() and in page_pool_disable_direct_recycling() because the driver does not disable the napi before destroying it. This is troublesome since on MT7916 it is required to unload the module and reinsert it with a different enable_6ghz parameter to change the frequency. The system generally becomes unstable after reinserting the module. Fix it by disabling rx napi inside mt7915e itself, even through napi rx is enabled in the mt76 generic code. Disabling it in the mt76 generic code was previously attempted and reverted because napi was disabled too late and rx handlers could access freed structures. Note that the tx side is enabled and disabled by the mt76 generic code. It is disabled in mt7915_stop_hardware() since, unlike mt7915_unregister_device(), is it also called in the probe path if an error happens after rx napi is enabled.

Linux

In the Linux kernel, the following vulnerability has been resolved: ASoC: adau1372: allow sleepy powerdown GPIOs gpiod_set_value() emits a WARN when the GPIO provider indicates that it can sleep, as is typical for I2C controlled GPIO expanders. Since the powerdown GPIO is only manipulated in process context, use the _cansleep variant instead. This lets the driver use sleepy GPIOs to power off the chip.

Linux

In the Linux kernel, the following vulnerability has been resolved: flow_dissector: avoid u16 truncation of skb->len when computing thoff __skb_flow_dissect() computes key_control->thoff using: key_control->thoff = min_t(u16, nhoff, skb ? skb->len : hlen); min_t(u16, ...) casts both arguments to u16 before comparing them. Whenever skb->len (or hlen) modulo 65536 is smaller than nhoff, the truncated length wins and thoff is set to a bogus small value, even when the dissection succeeded. For example, an IPv4/TCP frame with skb->len == 65540 and nhoff == 34 gets thoff == 4, so callers such as skb_probe_transport_header() point the transport header inside the Ethernet header. Such skbs are not exotic: - At the time of commit d0c081b49137 ("flow_dissector: properly cap thoff field"), AF_PACKET with PACKET_VNET_HDR could already build GSO skbs larger than 64KB (MTU checks are skipped for GSO, and alloc_skb_with_frags() accepted up to MAX_SKB_FRAGS (17) order-0 pages on top of the linear part). - BIG TCP now makes skbs larger than 64KB common. - The following patch makes tun_get_user() dissect IFF_TAP frames before eth_type_trans() pulls the Ethernet header, so a GSO frame carrying a 65522..65535 byte L3 packet will be dissected with skb->len in [65536, 65549]. Compare as u32 instead. If the resulting offset cannot be represented in the u16 key_control->thoff, cap it to U16_MAX and report the dissection as failed rather than silently returning a wrong transport offset. thoff is still set on failure, as some callers (such as eth_get_headlen()) use it regardless of the return value.

Linux

In the Linux kernel, the following vulnerability has been resolved: net: always dissect GSO packets in __virtio_net_hdr_to_skb() Commit 9e8db5913264 ("net: avoid false positives in untrusted gso validation") added a '&& skb->network_header' check before flow-dissecting GSO packets without VIRTIO_NET_HDR_F_NEEDS_CSUM in __virtio_net_hdr_to_skb(), because some callers (such as tun_get_user(), tun_xdp_one(), virtnet_receive_done(), and raw_verify_header()) called virtio_net_hdr_*_to_skb() before initializing skb->network_header and skb->dev. Because __alloc_skb() and __build_skb_around() zero-initialize skb->network_header to 0 (unlike mac_header and transport_header which are initialized to ~0U), those four callers always had skb->network_header == 0 and bypassed flow dissection in __virtio_net_hdr_to_skb(). More generally, skb->network_header is an offset from skb->head (where 0 is also a valid offset whenever skb_headroom(skb) is 0), not a boolean flag. Whenever the 'if (gso_type && skb->network_header)' branch was skipped, the fallback 'else if (gso_type)' only pulled nh_min_len + thlen (40 bytes for TCPv4) without dissecting the packet, without validating ip_proto or n_proto, and without setting skb->transport_header. If the packet has a malformed network header, it is not rejected and a subsequent skb_probe_transport_header() also fails, leaving skb->transport_header at ~0U (0xffff). Similarly, if an IPv4 packet carries IP options (ihl > 5) or an IPv6 packet carries extension headers, pulling only nh_min_len + thlen can leave the TCP header outside skb->head. In both cases, tcp_hdrlen(skb) in skb_gso_transport_seglen() reads out-of-bounds: BUG: KASAN: slab-out-of-bounds in skb_gso_transport_seglen Read of size 2 by task poc/133 skb_gso_transport_seglen (net/core/gso.c:155) skb_gso_validate_mac_len (net/core/gso.c:270) tbf_enqueue (net/sched/sch_tbf.c:260) dev_qdisc_enqueue (net/core/dev.c:4227) __dev_queue_xmit (net/core/dev.c:4884) In addition, checking virtio_net_hdr_match_proto() only inside 'if (!skb->protocol)' before flow dissection both skipped validation when skb->protocol was pre-set by the caller and rejected VLAN-tagged frames whose outer L2 protocol is ETH_P_8021Q or ETH_P_8021AD. Fix this by: 1. Initializing skb->dev and skb->network_header (plus skb->protocol for IFF_TUN) before virtio_net_hdr_*_to_skb() in tun_get_user(), tun_xdp_one(), virtnet_receive_done(), and raw_verify_header(). In tun_get_user(), drop the redundant skb_reset_mac_header(skb) in the IFF_TUN case since __virtio_net_hdr_to_skb() unconditionally resets mac_header. 2. Removing '&& skb->network_header' and the unvalidated 'else if (gso_type)' fallback in __virtio_net_hdr_to_skb() so all GSO packets without VIRTIO_NET_HDR_F_NEEDS_CSUM are flow-dissected, have their transport header pulled into linear data, and have skb->transport_header set. 3. Moving the virtio_net_hdr_match_proto() check to after skb_flow_dissect_flow_keys_basic(), validating keys.basic.n_proto against hdr_gso_type.

Linux

In the Linux kernel, the following vulnerability has been resolved: devlink: fix devlink_rel reference leak when notify work is pending devlink_rel_nested_in_notify_work_schedule() takes a reference on the devlink_rel for the notify work and then queues the work, ignoring the return value of schedule_delayed_work(). If the work is already pending, nothing new is queued, the work runs only once and drops only one reference, so the extra one is leaked together with the devlink_rel and its index in devlink_rels. This is easy to hit. devl_register() of a nested instance queues the work, and if the instance is unregistered before the work has run, devlink_rel_put() queues it again while it is still pending. The work also keeps rescheduling itself for as long as the parent devlink lock cannot be taken, which widens the window. Registering and unregistering a nested devlink instance 100 times while holding the parent lock leaks all 100 devlink_rel objects. The reschedule path in devlink_rel_nested_in_notify_work() has the same problem: if the work was queued again while it was running, the reference it holds is never dropped. Drop the reference in both places when the work was already pending. The pending work holds its own reference, so this can not be the last one.

Linux

In the Linux kernel, the following vulnerability has been resolved: pfcp: make sure the SEID is linear before reading it pfcp_encap_recv() only makes sure that the UDP header and the 4 byte PFCP header are in the linear area. When the S flag is set, pfcp_session_recv() then reads the 8 byte SEID that follows, which is not covered by the pskb_may_pull() check. A short PFCP packet with the S flag set, or one whose session header lies in a fragment, therefore makes pfcp_session_recv() read beyond the end of the packet data, and whatever it finds there is stored in the tunnel metadata that flower later classifies on. Pull up to the end of the SEID before reading it, and drop packets that are too short to contain it. Reload the header pointer afterwards since pskb_may_pull() may reallocate the skb head. Use offsetofend() rather than sizeof(struct pfcphdr_session): the structure is not packed, so its size is 16 bytes because of the alignment of the __be64 member, while the header on the wire is only 12 bytes, and valid session messages would be dropped.

Linux

In the Linux kernel, the following vulnerability has been resolved: riscv: ptrace: Zero-initialize regset buffers before copyin user_regset_copyin() only copies `count` bytes, so a PTRACE_SETREGSET request with a short iov_len leaves the rest of the stack buffer uninitialized. In riscv_vr_set() that garbage is validated and copied into the target's vstate, and in riscv_cfi_set() it can be stored as the target's shadow stack pointer. Zero both buffers, as the get path already does.

Linux

In the Linux kernel, the following vulnerability has been resolved: flow_dissector: avoid u16 truncation of hlen in bpf_flow_dissect() bpf_flow_dissect() sanitizes the offsets returned by the BPF program with clamp_t(u16, ..., hlen). If hlen is bigger than 65535, the upper bound is truncated before the comparison. For instance, with hlen == 65536, both nhoff and thoff are forced to zero, even if the program returned sane values. This is the same class of bug as the one fixed in commit 99bda1ecbd56 ("flow_dissector: avoid u16 truncation of skb->len when computing thoff"). That fix does not cover the BPF path, because __skb_flow_dissect() returns early when a BPF flow dissector handles the packet. struct bpf_flow_keys stores nhoff and thoff as u16, so cap hlen to U16_MAX before clamping. Found by sashiko while reviewing the commit above.

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/imagination: Fix reference and vm_bo handling in remap() When a map overlaps part of an existing mapping, pvr_vm_gpuva_remap() splits the mapping into prev/next parts covering what the request did not take, instead of creating something new. It gets two things wrong. - A GEM reference is taken for each part but it's never needed and never dropped, so it leaks one reference per split (remap-next-in-2m in tests/imagination/pvr_vm_map.c): CRITICAL: 33558528 bytes of shmem still held after close - The parts still belong to the original object being split, but pvr_vm_gpuva_remap() links them to ctx->gpuvm_bo, the new object's vm_bo: prev_va --obj--> BO_A BO_B <--obj-- vm_bo (ctx->gpuvm_bo) \___________link___________/ (mismatch: BO_A != BO_B) drm_gpuva_link() catches the mismatch: WARNING: drivers/gpu/drm/drm_gpuvm.c:2108 at drm_gpuva_link+0x2ec/0x310 drm_WARN_ON(obj != vm_bo->obj) Call trace: drm_gpuva_link pvr_vm_gpuva_remap __drm_gpuvm_sm_map pvr_vm_map pvr_ioctl_vm_map Link them to op->remap.unmap->va->vm_bo instead. The locking of the GPUVA lists of the other objects touched by a split is not addressed here; it is handled by switching the GPUVM to immediate mode.

Linux

In the Linux kernel, the following vulnerability has been resolved: mlxsw: spectrum_flower: Fix port range register leak in tmplt_create() mlxsw_sp_flower_tmplt_create() parses a flow_cls_offload template into a stack-local struct mlxsw_sp_acl_rule_info purely to compute rulei.values.elusage. Parsing can acquire port range registers via mlxsw_sp_flower_parse_ports_range(), but since this rulei never goes through mlxsw_sp_acl_rulei_destroy(), those registers were never released, including through chain template deletion. Factor out of mlxsw_sp_acl_rulei_destroy() the code to actually release the necessary resources and call from mlxsw_sp_flower_tmplt_create() to plug the leak. The issue was found during a review of Wentao Liang's patch referenced below.

Linux

In the Linux kernel, the following vulnerability has been resolved: vsock: Fix memory leak in vmci_transport_recv_dgram_cb() During the closure of a datagram socket, the vmci_transport_recv_dgram_cb() function may be called, which will add the packets to the socket's backlog; then, after the receive queue is cleared, the __release_sock() function will move the packet back from the socket's backlog to the receive queue, which will lead to a memory leak. sock_close __sock_release __vsock_release // take ownership by user-space lock_sock_nested sock_set_flag(sk, SOCK_DEAD) vmci_transport_release vmci_dispatch_dgs vmci_datagram_invoke_guest_handler vmci_transport_recv_dgram_cb sk_receive_skb if (!sock_owned_by_user(sk)) [...] (!) else if sk_add_backlog() [...] (!) skb_queue_purge(&sk->sk_receive_queue) release_sock __release_sock // move all packets from the backlog // to the receive queue (!) sk_backlog_rcv vsock_queue_rcv_skb [...] __sock_queue_rcv_skb if (!sock_flag(sk, SOCK_DEAD)) // Since the socket is "dead", the function will not // be called sk->sk_data_ready(sk) sock_release_ownership(sk); sock_put Add a receive queue cleanup in the socket destructor to fix this. syzkaller report: unreferenced object 0xffff8880117bfb80 (size 240): comm "irq/56-vmw_vmci", pid 205, jiffies 4295064448 (age 73.992s) hex dump (first 32 bytes): 78 dc 26 1e 80 88 ff ff 78 dc 26 1e 80 88 ff ff x.&.....x.&..... 00 00 00 00 00 00 00 00 c0 da 26 1e 80 88 ff ff ..........&..... backtrace: [<ffffffff82be8d97>] __alloc_skb+0x287/0x330 net/core/skbuff.c:505 [<ffffffffa114bcbd>] vmci_transport_recv_dgram_cb+0xbd/0x210 [vmw_vsock_vmci_transport] [<ffffffffa05072c6>] vmci_datagram_invoke_guest_handler+0x366/0x450 [vmw_vmci] [<ffffffffa050ad75>] vmci_dispatch_dgs+0x235/0x490 [vmw_vmci] [<ffffffffa050b025>] vmci_interrupt+0x55/0x290 [vmw_vmci] [<ffffffff8137ad2b>] irq_thread_fn+0x8b/0x1a0 kernel/irq/manage.c:1205 [<ffffffff8137c7ae>] irq_thread+0x28e/0x530 kernel/irq/manage.c:1314 [<ffffffff81253dd6>] kthread+0x2e6/0x3a0 kernel/kthread.c:376 [<ffffffff81004392>] ret_from_fork+0x22/0x30 arch/x86/entry/entry_64.S:295 Found by InfoTeCS on behalf of Linux Verification Center (linuxtesting.org) with Syzkaller.

Linux

In the Linux kernel, the following vulnerability has been resolved: drm/vc4: Disable the V3D interrupt across runtime suspend vc4_irq_disable() masks the V3D interrupt sources and then calls synchronize_irq() before the V3D is powered down. However, by itself, this is not enough to quiesce the interrupt handler. synchronize_irq() waits for handlers that have already set IRQD_IRQ_INPROGRESS, and for irqchips reporting IRQCHIP_STATE_ACTIVE. A GIC interrupt chip is able to mark an interrupt active as soon as a CPU acknowledges it, so there these checks cover the whole dispatch path. On RPi 0-3, however, the interrupt controller is ARMCTRL, which has no active state. A CPU that has read the hwirq out of the pending register but has not yet reached handle_level_irq() stays invisible to synchronize_irq(). During a power transition, vc4_irq() may therefore run after the power domain is off, where every V3D register read will return 0xdeadbeef. 0xdeadbeef has FLDONE, FRDONE and OUTOMEM set. This problem doesn't trigger NULL pointer dereference issues only because the functions vc4_irq_finish_bin_job() and vc4_irq_finish_render_job() return early when there is no job pointer, but OUTOMEM is still able to schedule vc4_overflow_mem_work() after vc4_irq_disable() has already cancelled it. Therefore, fix the spurious interrupts by disabling the interrupt line across the PM transition. To preserve the enable/disable balance, move vc4_irq_install() and request the IRQ line disabled before the first runtime resume.

Linux

In the Linux kernel, the following vulnerability has been resolved: HID: steam: Rearrange teardown sequence This fixes a narrow window during the teardown sequence where callbacks could still be scheduled during cleanup that would then have a dangling pointer to the now-freed steam struct. This also puts work canceling for rumble and mode switch in steam_unregister, as that shouldn't persist while the client hdev is open.

Linux

In the Linux kernel, the following vulnerability has been resolved: HID: steam: Refactor registration This refactors and simplifies the registration/unregistration flow. Since we now only perform registration when the client isn't opened anymore, the logic for handling that can be removed, and the rest of the function streamlined. We also remove the previous assumption that we have a serial number to show we're registered, replacing it with a single purpose boolean. In a previous refactor the code for unregistering a battery if later registration steps failed was accidentally left out. As a result, a lingering power_supply object could get left over after the steam object was torn down. This is also fixed here.

Linux

In the Linux kernel, the following vulnerability has been resolved: HID: steam: Clean up locking This cleans up several issues with locking behavior, including RCU accesses not being guarded behind a lock.

Linux

In the Linux kernel, the following vulnerability has been resolved: sched_ext: Build the cid tables privately and publish them with RCU The cid tables are visible to the cid kfuncs while being modified: the first enable publishes the global pointers before filling them, ops.init_cids() overrides rewrite them in place, and re-enables rebuild them in place. A racing TRACING or SYSCALL program can read unfilled entries, including uninitialized memory in the kmalloc'd tables, or torn topo updates. Tie the tables' lifetimes to the root sched instead: each root enable builds a fresh set privately and publishes the per-table __rcu globals once the layout is final, and root disable unpublishes and RCU-frees the set. A non-NULL global is now always a fully built table which stays valid for the reader's RCU read section, and lookups stay two loads. Kfuncs treat NULL as no-mapping, also after the scheduler exits instead of reporting the stale last mapping. The cid kfuncs are available whether the root scheduler is cid-form or cpu-form, the latter to allow gradual migration to cids. Every root therefore builds and publishes a default mapping. Every reader must either be gated on scheduler liveness or NULL-check inside an RCU read section. Fix the two kfuncs that were neither: scx_bpf_this_cid() read the table with no RCU or preemption protection and scx_bpf_task_cid() relied on KF_RCU, which doesn't put a sleepable program in an RCU read section. The hotplug callbacks are instead serialized by retiring the tables inside the cpus_read_lock() section that clears scx_root. v2: Document why every root builds the tables (desc + cid.c comment).

Linux

In the Linux kernel, the following vulnerability has been resolved: ext4: don't enable large folios on encrypted inodes Commit 709f0f1f1bf5 ("ext4: add checks for large folio incompatibilities when BS > PS") disables large folios at mount time if the "encrypt" feature is set. However, enabling that feature on a mounted filesystem circumvents that check, and writing encrypted data results in a crash: BUG: kernel NULL pointer dereference, address: 0000000000000028 Workqueue: ext4-rsv-conversion ext4_end_io_rsv_work RIP: 0010:ext4_finish_bio+0xf7/0x3e0 Call Trace: ext4_release_io_end+0x51/0x110 ext4_end_io_end+0x4c/0xe0 ext4_end_io_rsv_work+0xb3/0x100 Fix this by disabling large folios on encrypted inodes, as is already done for inodes that use data journaling. This affects all kernels between v6.19 and v7.2, but it is not needed in v7.3, where the code was significantly refactored.

Linux

In the Linux kernel, the following vulnerability has been resolved: ASoC: qcom: lpass-cpu: use DAI table index for playback constraints The DAI ID is a hardware port ID and is not necessarily the index of the corresponding entry in variant->dai_driver. When configuring a DAI with LPAIF_I2SCTL_MODE_QUAD01, the probe code currently uses dai_id to index variant->dai_driver. This is incorrect for platforms where DAI IDs are sparse. For example, the IPQ806x MI2S DAI has ID 4 while it is the only entry in the DAI driver table. If that DAI is configured for QUAD01 and probe reaches this branch, the code writes beyond that single entry via dai_driver[4] instead of updating the DAI being processed. Use the loop index i when updating the current DAI's playback channel constraints, while retaining dai_id for indexing the hardware-port specific playback SD-line mode array. This fixes the incorrect DAI table access for platforms where the DAI ID does not match its position in the driver table.

Linux

In the Linux kernel, the following vulnerability has been resolved: x86/mm: Drop unnecessary PMD page copy when freeing On a box with a discrete GPU, lockdep reports a possible deadlock as soon as kswapd shrinks the TTM page pool. The immediate cause is an x86 commit that added an mmap_read_lock() to kernel page protection munging code. The huge vmap code holds the same lock over a GFP_KERNEL allocation, which is a no-no now that reclaim can take it. That allocation is in a page table *free* path and ends up being for dubious purposes[1]. Basically, it tries to avoid hardware setting Accessed=1 in page table entries that are unreachable by the hardware, a non-issue. Remove the PMD copy. Detach the original PMD page at the PUD, flush the mid-level caches, and free the PTE tables straight from the detached PMD page. With no allocation left, the locking issue is gone. Lockdep splat/analysis: WARNING: possible circular locking dependency detected 7.3.0-rc3-f6e7b42bf05b+ #183 Tainted: G U ------------------------------------------------------ kswapd0/269 is trying to acquire lock: ((init_mm).mmap_lock){++++}-{4:4}, at: change_page_attr_set_clr+0x29a/0x4a0 but task is already holding lock: (pool_shrink_rwsem){.+.+}-{4:4}, at: ttm_pool_shrink+0xb2/0x330 [ttm] Chain exists of: (init_mm).mmap_lock --> fs_reclaim --> pool_shrink_rwsem The cycle is built from three edges: 1) pool_shrink_rwsem -> (init_mm).mmap_lock The TTM shrinker restores the caching attribute of every page it frees, while holding pool_shrink_rwsem: ttm_pool_shrink() -> ttm_pool_dispose_list() -> ttm_pool_free_page() -> set_pages_wb() -> change_page_attr_set_clr() [ init_mm mmap read lock ] 2) fs_reclaim -> pool_shrink_rwsem The same shrinker, called from reclaim. 3) (init_mm).mmap_lock -> fs_reclaim ioremap() installing a huge PUD mapping over an existing PMD table: ioremap_page_range() -> vmap_range_noflush() -> vmap_try_huge_pud() [ init_mm mmap read lock ] -> pud_free_pmd_page() -> __get_free_page(GFP_KERNEL) [ enters reclaim ] [ dhansen: Lots of changelog munging/trimming and merged comments from my version of the fix. ]

Linux

In the Linux kernel, the following vulnerability has been resolved: rust: irq: pass RegistrationInner as cookie to request_irq With commit ba268514ea14 ("rust: devres: fix race condition due to nesting"), Devres::new() returns Result<Self> by value instead of initializing in-place via PinInit. Because request_irq() is called inside Devres::new(), registration.inner is still uninitialized when the IRQ is enabled, so accessing registration.inner.device() in the IRQ callback can read uninitialized memory. Fix this by storing the Device and handler pointer in RegistrationInner and passing RegistrationInner as the IRQ cookie after its fields are initialized. This is not needed in mainline because commit 98c63ce4d760 ("rust: irq: make Registration compatible with lifetime-bound drivers") removed Devres from irq::Registration and moved request_irq() after all fields are initialized.