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authorDaniel Baumann <daniel.baumann@progress-linux.org>2024-04-07 18:49:45 +0000
committerDaniel Baumann <daniel.baumann@progress-linux.org>2024-04-07 18:49:45 +0000
commit2c3c1048746a4622d8c89a29670120dc8fab93c4 (patch)
tree848558de17fb3008cdf4d861b01ac7781903ce39 /Documentation/x86
parentInitial commit. (diff)
downloadlinux-2c3c1048746a4622d8c89a29670120dc8fab93c4.tar.xz
linux-2c3c1048746a4622d8c89a29670120dc8fab93c4.zip
Adding upstream version 6.1.76.upstream/6.1.76upstream
Signed-off-by: Daniel Baumann <daniel.baumann@progress-linux.org>
Diffstat (limited to '')
-rw-r--r--Documentation/x86/amd-memory-encryption.rst133
-rw-r--r--Documentation/x86/amd_hsmp.rst86
-rw-r--r--Documentation/x86/boot.rst1442
-rw-r--r--Documentation/x86/booting-dt.rst21
-rw-r--r--Documentation/x86/buslock.rst126
-rw-r--r--Documentation/x86/cpuinfo.rst154
-rw-r--r--Documentation/x86/earlyprintk.rst151
-rw-r--r--Documentation/x86/elf_auxvec.rst53
-rw-r--r--Documentation/x86/entry_64.rst110
-rw-r--r--Documentation/x86/exception-tables.rst357
-rw-r--r--Documentation/x86/features.rst3
-rw-r--r--Documentation/x86/i386/IO-APIC.rst123
-rw-r--r--Documentation/x86/i386/index.rst10
-rw-r--r--Documentation/x86/ifs.rst2
-rw-r--r--Documentation/x86/index.rst44
-rw-r--r--Documentation/x86/intel-hfi.rst72
-rw-r--r--Documentation/x86/intel_txt.rst227
-rw-r--r--Documentation/x86/iommu.rst151
-rw-r--r--Documentation/x86/kernel-stacks.rst152
-rw-r--r--Documentation/x86/mds.rst193
-rw-r--r--Documentation/x86/microcode.rst240
-rw-r--r--Documentation/x86/mtrr.rst354
-rw-r--r--Documentation/x86/orc-unwinder.rst182
-rw-r--r--Documentation/x86/pat.rst240
-rw-r--r--Documentation/x86/pti.rst195
-rw-r--r--Documentation/x86/resctrl.rst1304
-rw-r--r--Documentation/x86/sgx.rst302
-rw-r--r--Documentation/x86/sva.rst286
-rw-r--r--Documentation/x86/tdx.rst218
-rw-r--r--Documentation/x86/tlb.rst83
-rw-r--r--Documentation/x86/topology.rst234
-rw-r--r--Documentation/x86/tsx_async_abort.rst117
-rw-r--r--Documentation/x86/usb-legacy-support.rst50
-rw-r--r--Documentation/x86/x86_64/5level-paging.rst67
-rw-r--r--Documentation/x86/x86_64/boot-options.rst319
-rw-r--r--Documentation/x86/x86_64/cpu-hotplug-spec.rst24
-rw-r--r--Documentation/x86/x86_64/fake-numa-for-cpusets.rst78
-rw-r--r--Documentation/x86/x86_64/fsgs.rst199
-rw-r--r--Documentation/x86/x86_64/index.rst17
-rw-r--r--Documentation/x86/x86_64/machinecheck.rst33
-rw-r--r--Documentation/x86/x86_64/mm.rst157
-rw-r--r--Documentation/x86/x86_64/uefi.rst58
-rw-r--r--Documentation/x86/xstate.rst74
-rw-r--r--Documentation/x86/zero-page.rst47
44 files changed, 8488 insertions, 0 deletions
diff --git a/Documentation/x86/amd-memory-encryption.rst b/Documentation/x86/amd-memory-encryption.rst
new file mode 100644
index 000000000..934310ce7
--- /dev/null
+++ b/Documentation/x86/amd-memory-encryption.rst
@@ -0,0 +1,133 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=====================
+AMD Memory Encryption
+=====================
+
+Secure Memory Encryption (SME) and Secure Encrypted Virtualization (SEV) are
+features found on AMD processors.
+
+SME provides the ability to mark individual pages of memory as encrypted using
+the standard x86 page tables. A page that is marked encrypted will be
+automatically decrypted when read from DRAM and encrypted when written to
+DRAM. SME can therefore be used to protect the contents of DRAM from physical
+attacks on the system.
+
+SEV enables running encrypted virtual machines (VMs) in which the code and data
+of the guest VM are secured so that a decrypted version is available only
+within the VM itself. SEV guest VMs have the concept of private and shared
+memory. Private memory is encrypted with the guest-specific key, while shared
+memory may be encrypted with hypervisor key. When SME is enabled, the hypervisor
+key is the same key which is used in SME.
+
+A page is encrypted when a page table entry has the encryption bit set (see
+below on how to determine its position). The encryption bit can also be
+specified in the cr3 register, allowing the PGD table to be encrypted. Each
+successive level of page tables can also be encrypted by setting the encryption
+bit in the page table entry that points to the next table. This allows the full
+page table hierarchy to be encrypted. Note, this means that just because the
+encryption bit is set in cr3, doesn't imply the full hierarchy is encrypted.
+Each page table entry in the hierarchy needs to have the encryption bit set to
+achieve that. So, theoretically, you could have the encryption bit set in cr3
+so that the PGD is encrypted, but not set the encryption bit in the PGD entry
+for a PUD which results in the PUD pointed to by that entry to not be
+encrypted.
+
+When SEV is enabled, instruction pages and guest page tables are always treated
+as private. All the DMA operations inside the guest must be performed on shared
+memory. Since the memory encryption bit is controlled by the guest OS when it
+is operating in 64-bit or 32-bit PAE mode, in all other modes the SEV hardware
+forces the memory encryption bit to 1.
+
+Support for SME and SEV can be determined through the CPUID instruction. The
+CPUID function 0x8000001f reports information related to SME::
+
+ 0x8000001f[eax]:
+ Bit[0] indicates support for SME
+ Bit[1] indicates support for SEV
+ 0x8000001f[ebx]:
+ Bits[5:0] pagetable bit number used to activate memory
+ encryption
+ Bits[11:6] reduction in physical address space, in bits, when
+ memory encryption is enabled (this only affects
+ system physical addresses, not guest physical
+ addresses)
+
+If support for SME is present, MSR 0xc00100010 (MSR_AMD64_SYSCFG) can be used to
+determine if SME is enabled and/or to enable memory encryption::
+
+ 0xc0010010:
+ Bit[23] 0 = memory encryption features are disabled
+ 1 = memory encryption features are enabled
+
+If SEV is supported, MSR 0xc0010131 (MSR_AMD64_SEV) can be used to determine if
+SEV is active::
+
+ 0xc0010131:
+ Bit[0] 0 = memory encryption is not active
+ 1 = memory encryption is active
+
+Linux relies on BIOS to set this bit if BIOS has determined that the reduction
+in the physical address space as a result of enabling memory encryption (see
+CPUID information above) will not conflict with the address space resource
+requirements for the system. If this bit is not set upon Linux startup then
+Linux itself will not set it and memory encryption will not be possible.
+
+The state of SME in the Linux kernel can be documented as follows:
+
+ - Supported:
+ The CPU supports SME (determined through CPUID instruction).
+
+ - Enabled:
+ Supported and bit 23 of MSR_AMD64_SYSCFG is set.
+
+ - Active:
+ Supported, Enabled and the Linux kernel is actively applying
+ the encryption bit to page table entries (the SME mask in the
+ kernel is non-zero).
+
+SME can also be enabled and activated in the BIOS. If SME is enabled and
+activated in the BIOS, then all memory accesses will be encrypted and it will
+not be necessary to activate the Linux memory encryption support. If the BIOS
+merely enables SME (sets bit 23 of the MSR_AMD64_SYSCFG), then Linux can activate
+memory encryption by default (CONFIG_AMD_MEM_ENCRYPT_ACTIVE_BY_DEFAULT=y) or
+by supplying mem_encrypt=on on the kernel command line. However, if BIOS does
+not enable SME, then Linux will not be able to activate memory encryption, even
+if configured to do so by default or the mem_encrypt=on command line parameter
+is specified.
+
+Secure Nested Paging (SNP)
+==========================
+
+SEV-SNP introduces new features (SEV_FEATURES[1:63]) which can be enabled
+by the hypervisor for security enhancements. Some of these features need
+guest side implementation to function correctly. The below table lists the
+expected guest behavior with various possible scenarios of guest/hypervisor
+SNP feature support.
+
++-----------------+---------------+---------------+------------------+
+| Feature Enabled | Guest needs | Guest has | Guest boot |
+| by the HV | implementation| implementation| behaviour |
++=================+===============+===============+==================+
+| No | No | No | Boot |
+| | | | |
++-----------------+---------------+---------------+------------------+
+| No | Yes | No | Boot |
+| | | | |
++-----------------+---------------+---------------+------------------+
+| No | Yes | Yes | Boot |
+| | | | |
++-----------------+---------------+---------------+------------------+
+| Yes | No | No | Boot with |
+| | | | feature enabled |
++-----------------+---------------+---------------+------------------+
+| Yes | Yes | No | Graceful boot |
+| | | | failure |
++-----------------+---------------+---------------+------------------+
+| Yes | Yes | Yes | Boot with |
+| | | | feature enabled |
++-----------------+---------------+---------------+------------------+
+
+More details in AMD64 APM[1] Vol 2: 15.34.10 SEV_STATUS MSR
+
+[1] https://www.amd.com/system/files/TechDocs/40332.pdf
diff --git a/Documentation/x86/amd_hsmp.rst b/Documentation/x86/amd_hsmp.rst
new file mode 100644
index 000000000..440e4b645
--- /dev/null
+++ b/Documentation/x86/amd_hsmp.rst
@@ -0,0 +1,86 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+============================================
+AMD HSMP interface
+============================================
+
+Newer Fam19h EPYC server line of processors from AMD support system
+management functionality via HSMP (Host System Management Port).
+
+The Host System Management Port (HSMP) is an interface to provide
+OS-level software with access to system management functions via a
+set of mailbox registers.
+
+More details on the interface can be found in chapter
+"7 Host System Management Port (HSMP)" of the family/model PPR
+Eg: https://www.amd.com/system/files/TechDocs/55898_B1_pub_0.50.zip
+
+HSMP interface is supported on EPYC server CPU models only.
+
+
+HSMP device
+============================================
+
+amd_hsmp driver under the drivers/platforms/x86/ creates miscdevice
+/dev/hsmp to let user space programs run hsmp mailbox commands.
+
+$ ls -al /dev/hsmp
+crw-r--r-- 1 root root 10, 123 Jan 21 21:41 /dev/hsmp
+
+Characteristics of the dev node:
+ * Write mode is used for running set/configure commands
+ * Read mode is used for running get/status monitor commands
+
+Access restrictions:
+ * Only root user is allowed to open the file in write mode.
+ * The file can be opened in read mode by all the users.
+
+In-kernel integration:
+ * Other subsystems in the kernel can use the exported transport
+ function hsmp_send_message().
+ * Locking across callers is taken care by the driver.
+
+
+An example
+==========
+
+To access hsmp device from a C program.
+First, you need to include the headers::
+
+ #include <linux/amd_hsmp.h>
+
+Which defines the supported messages/message IDs.
+
+Next thing, open the device file, as follows::
+
+ int file;
+
+ file = open("/dev/hsmp", O_RDWR);
+ if (file < 0) {
+ /* ERROR HANDLING; you can check errno to see what went wrong */
+ exit(1);
+ }
+
+The following IOCTL is defined:
+
+``ioctl(file, HSMP_IOCTL_CMD, struct hsmp_message *msg)``
+ The argument is a pointer to a::
+
+ struct hsmp_message {
+ __u32 msg_id; /* Message ID */
+ __u16 num_args; /* Number of input argument words in message */
+ __u16 response_sz; /* Number of expected output/response words */
+ __u32 args[HSMP_MAX_MSG_LEN]; /* argument/response buffer */
+ __u16 sock_ind; /* socket number */
+ };
+
+The ioctl would return a non-zero on failure; you can read errno to see
+what happened. The transaction returns 0 on success.
+
+More details on the interface and message definitions can be found in chapter
+"7 Host System Management Port (HSMP)" of the respective family/model PPR
+eg: https://www.amd.com/system/files/TechDocs/55898_B1_pub_0.50.zip
+
+User space C-APIs are made available by linking against the esmi library,
+which is provided by the E-SMS project https://developer.amd.com/e-sms/.
+See: https://github.com/amd/esmi_ib_library
diff --git a/Documentation/x86/boot.rst b/Documentation/x86/boot.rst
new file mode 100644
index 000000000..894a19897
--- /dev/null
+++ b/Documentation/x86/boot.rst
@@ -0,0 +1,1442 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===========================
+The Linux/x86 Boot Protocol
+===========================
+
+On the x86 platform, the Linux kernel uses a rather complicated boot
+convention. This has evolved partially due to historical aspects, as
+well as the desire in the early days to have the kernel itself be a
+bootable image, the complicated PC memory model and due to changed
+expectations in the PC industry caused by the effective demise of
+real-mode DOS as a mainstream operating system.
+
+Currently, the following versions of the Linux/x86 boot protocol exist.
+
+============= ============================================================
+Old kernels zImage/Image support only. Some very early kernels
+ may not even support a command line.
+
+Protocol 2.00 (Kernel 1.3.73) Added bzImage and initrd support, as
+ well as a formalized way to communicate between the
+ boot loader and the kernel. setup.S made relocatable,
+ although the traditional setup area still assumed
+ writable.
+
+Protocol 2.01 (Kernel 1.3.76) Added a heap overrun warning.
+
+Protocol 2.02 (Kernel 2.4.0-test3-pre3) New command line protocol.
+ Lower the conventional memory ceiling. No overwrite
+ of the traditional setup area, thus making booting
+ safe for systems which use the EBDA from SMM or 32-bit
+ BIOS entry points. zImage deprecated but still
+ supported.
+
+Protocol 2.03 (Kernel 2.4.18-pre1) Explicitly makes the highest possible
+ initrd address available to the bootloader.
+
+Protocol 2.04 (Kernel 2.6.14) Extend the syssize field to four bytes.
+
+Protocol 2.05 (Kernel 2.6.20) Make protected mode kernel relocatable.
+ Introduce relocatable_kernel and kernel_alignment fields.
+
+Protocol 2.06 (Kernel 2.6.22) Added a field that contains the size of
+ the boot command line.
+
+Protocol 2.07 (Kernel 2.6.24) Added paravirtualised boot protocol.
+ Introduced hardware_subarch and hardware_subarch_data
+ and KEEP_SEGMENTS flag in load_flags.
+
+Protocol 2.08 (Kernel 2.6.26) Added crc32 checksum and ELF format
+ payload. Introduced payload_offset and payload_length
+ fields to aid in locating the payload.
+
+Protocol 2.09 (Kernel 2.6.26) Added a field of 64-bit physical
+ pointer to single linked list of struct setup_data.
+
+Protocol 2.10 (Kernel 2.6.31) Added a protocol for relaxed alignment
+ beyond the kernel_alignment added, new init_size and
+ pref_address fields. Added extended boot loader IDs.
+
+Protocol 2.11 (Kernel 3.6) Added a field for offset of EFI handover
+ protocol entry point.
+
+Protocol 2.12 (Kernel 3.8) Added the xloadflags field and extension fields
+ to struct boot_params for loading bzImage and ramdisk
+ above 4G in 64bit.
+
+Protocol 2.13 (Kernel 3.14) Support 32- and 64-bit flags being set in
+ xloadflags to support booting a 64-bit kernel from 32-bit
+ EFI
+
+Protocol 2.14 BURNT BY INCORRECT COMMIT
+ ae7e1238e68f2a472a125673ab506d49158c1889
+ (x86/boot: Add ACPI RSDP address to setup_header)
+ DO NOT USE!!! ASSUME SAME AS 2.13.
+
+Protocol 2.15 (Kernel 5.5) Added the kernel_info and kernel_info.setup_type_max.
+============= ============================================================
+
+.. note::
+ The protocol version number should be changed only if the setup header
+ is changed. There is no need to update the version number if boot_params
+ or kernel_info are changed. Additionally, it is recommended to use
+ xloadflags (in this case the protocol version number should not be
+ updated either) or kernel_info to communicate supported Linux kernel
+ features to the boot loader. Due to very limited space available in
+ the original setup header every update to it should be considered
+ with great care. Starting from the protocol 2.15 the primary way to
+ communicate things to the boot loader is the kernel_info.
+
+
+Memory Layout
+=============
+
+The traditional memory map for the kernel loader, used for Image or
+zImage kernels, typically looks like::
+
+ | |
+ 0A0000 +------------------------+
+ | Reserved for BIOS | Do not use. Reserved for BIOS EBDA.
+ 09A000 +------------------------+
+ | Command line |
+ | Stack/heap | For use by the kernel real-mode code.
+ 098000 +------------------------+
+ | Kernel setup | The kernel real-mode code.
+ 090200 +------------------------+
+ | Kernel boot sector | The kernel legacy boot sector.
+ 090000 +------------------------+
+ | Protected-mode kernel | The bulk of the kernel image.
+ 010000 +------------------------+
+ | Boot loader | <- Boot sector entry point 0000:7C00
+ 001000 +------------------------+
+ | Reserved for MBR/BIOS |
+ 000800 +------------------------+
+ | Typically used by MBR |
+ 000600 +------------------------+
+ | BIOS use only |
+ 000000 +------------------------+
+
+When using bzImage, the protected-mode kernel was relocated to
+0x100000 ("high memory"), and the kernel real-mode block (boot sector,
+setup, and stack/heap) was made relocatable to any address between
+0x10000 and end of low memory. Unfortunately, in protocols 2.00 and
+2.01 the 0x90000+ memory range is still used internally by the kernel;
+the 2.02 protocol resolves that problem.
+
+It is desirable to keep the "memory ceiling" -- the highest point in
+low memory touched by the boot loader -- as low as possible, since
+some newer BIOSes have begun to allocate some rather large amounts of
+memory, called the Extended BIOS Data Area, near the top of low
+memory. The boot loader should use the "INT 12h" BIOS call to verify
+how much low memory is available.
+
+Unfortunately, if INT 12h reports that the amount of memory is too
+low, there is usually nothing the boot loader can do but to report an
+error to the user. The boot loader should therefore be designed to
+take up as little space in low memory as it reasonably can. For
+zImage or old bzImage kernels, which need data written into the
+0x90000 segment, the boot loader should make sure not to use memory
+above the 0x9A000 point; too many BIOSes will break above that point.
+
+For a modern bzImage kernel with boot protocol version >= 2.02, a
+memory layout like the following is suggested::
+
+ ~ ~
+ | Protected-mode kernel |
+ 100000 +------------------------+
+ | I/O memory hole |
+ 0A0000 +------------------------+
+ | Reserved for BIOS | Leave as much as possible unused
+ ~ ~
+ | Command line | (Can also be below the X+10000 mark)
+ X+10000 +------------------------+
+ | Stack/heap | For use by the kernel real-mode code.
+ X+08000 +------------------------+
+ | Kernel setup | The kernel real-mode code.
+ | Kernel boot sector | The kernel legacy boot sector.
+ X +------------------------+
+ | Boot loader | <- Boot sector entry point 0000:7C00
+ 001000 +------------------------+
+ | Reserved for MBR/BIOS |
+ 000800 +------------------------+
+ | Typically used by MBR |
+ 000600 +------------------------+
+ | BIOS use only |
+ 000000 +------------------------+
+
+ ... where the address X is as low as the design of the boot loader permits.
+
+
+The Real-Mode Kernel Header
+===========================
+
+In the following text, and anywhere in the kernel boot sequence, "a
+sector" refers to 512 bytes. It is independent of the actual sector
+size of the underlying medium.
+
+The first step in loading a Linux kernel should be to load the
+real-mode code (boot sector and setup code) and then examine the
+following header at offset 0x01f1. The real-mode code can total up to
+32K, although the boot loader may choose to load only the first two
+sectors (1K) and then examine the bootup sector size.
+
+The header looks like:
+
+=========== ======== ===================== ============================================
+Offset/Size Proto Name Meaning
+=========== ======== ===================== ============================================
+01F1/1 ALL(1) setup_sects The size of the setup in sectors
+01F2/2 ALL root_flags If set, the root is mounted readonly
+01F4/4 2.04+(2) syssize The size of the 32-bit code in 16-byte paras
+01F8/2 ALL ram_size DO NOT USE - for bootsect.S use only
+01FA/2 ALL vid_mode Video mode control
+01FC/2 ALL root_dev Default root device number
+01FE/2 ALL boot_flag 0xAA55 magic number
+0200/2 2.00+ jump Jump instruction
+0202/4 2.00+ header Magic signature "HdrS"
+0206/2 2.00+ version Boot protocol version supported
+0208/4 2.00+ realmode_swtch Boot loader hook (see below)
+020C/2 2.00+ start_sys_seg The load-low segment (0x1000) (obsolete)
+020E/2 2.00+ kernel_version Pointer to kernel version string
+0210/1 2.00+ type_of_loader Boot loader identifier
+0211/1 2.00+ loadflags Boot protocol option flags
+0212/2 2.00+ setup_move_size Move to high memory size (used with hooks)
+0214/4 2.00+ code32_start Boot loader hook (see below)
+0218/4 2.00+ ramdisk_image initrd load address (set by boot loader)
+021C/4 2.00+ ramdisk_size initrd size (set by boot loader)
+0220/4 2.00+ bootsect_kludge DO NOT USE - for bootsect.S use only
+0224/2 2.01+ heap_end_ptr Free memory after setup end
+0226/1 2.02+(3) ext_loader_ver Extended boot loader version
+0227/1 2.02+(3) ext_loader_type Extended boot loader ID
+0228/4 2.02+ cmd_line_ptr 32-bit pointer to the kernel command line
+022C/4 2.03+ initrd_addr_max Highest legal initrd address
+0230/4 2.05+ kernel_alignment Physical addr alignment required for kernel
+0234/1 2.05+ relocatable_kernel Whether kernel is relocatable or not
+0235/1 2.10+ min_alignment Minimum alignment, as a power of two
+0236/2 2.12+ xloadflags Boot protocol option flags
+0238/4 2.06+ cmdline_size Maximum size of the kernel command line
+023C/4 2.07+ hardware_subarch Hardware subarchitecture
+0240/8 2.07+ hardware_subarch_data Subarchitecture-specific data
+0248/4 2.08+ payload_offset Offset of kernel payload
+024C/4 2.08+ payload_length Length of kernel payload
+0250/8 2.09+ setup_data 64-bit physical pointer to linked list
+ of struct setup_data
+0258/8 2.10+ pref_address Preferred loading address
+0260/4 2.10+ init_size Linear memory required during initialization
+0264/4 2.11+ handover_offset Offset of handover entry point
+0268/4 2.15+ kernel_info_offset Offset of the kernel_info
+=========== ======== ===================== ============================================
+
+.. note::
+ (1) For backwards compatibility, if the setup_sects field contains 0, the
+ real value is 4.
+
+ (2) For boot protocol prior to 2.04, the upper two bytes of the syssize
+ field are unusable, which means the size of a bzImage kernel
+ cannot be determined.
+
+ (3) Ignored, but safe to set, for boot protocols 2.02-2.09.
+
+If the "HdrS" (0x53726448) magic number is not found at offset 0x202,
+the boot protocol version is "old". Loading an old kernel, the
+following parameters should be assumed::
+
+ Image type = zImage
+ initrd not supported
+ Real-mode kernel must be located at 0x90000.
+
+Otherwise, the "version" field contains the protocol version,
+e.g. protocol version 2.01 will contain 0x0201 in this field. When
+setting fields in the header, you must make sure only to set fields
+supported by the protocol version in use.
+
+
+Details of Header Fields
+========================
+
+For each field, some are information from the kernel to the bootloader
+("read"), some are expected to be filled out by the bootloader
+("write"), and some are expected to be read and modified by the
+bootloader ("modify").
+
+All general purpose boot loaders should write the fields marked
+(obligatory). Boot loaders who want to load the kernel at a
+nonstandard address should fill in the fields marked (reloc); other
+boot loaders can ignore those fields.
+
+The byte order of all fields is littleendian (this is x86, after all.)
+
+============ ===========
+Field name: setup_sects
+Type: read
+Offset/size: 0x1f1/1
+Protocol: ALL
+============ ===========
+
+ The size of the setup code in 512-byte sectors. If this field is
+ 0, the real value is 4. The real-mode code consists of the boot
+ sector (always one 512-byte sector) plus the setup code.
+
+============ =================
+Field name: root_flags
+Type: modify (optional)
+Offset/size: 0x1f2/2
+Protocol: ALL
+============ =================
+
+ If this field is nonzero, the root defaults to readonly. The use of
+ this field is deprecated; use the "ro" or "rw" options on the
+ command line instead.
+
+============ ===============================================
+Field name: syssize
+Type: read
+Offset/size: 0x1f4/4 (protocol 2.04+) 0x1f4/2 (protocol ALL)
+Protocol: 2.04+
+============ ===============================================
+
+ The size of the protected-mode code in units of 16-byte paragraphs.
+ For protocol versions older than 2.04 this field is only two bytes
+ wide, and therefore cannot be trusted for the size of a kernel if
+ the LOAD_HIGH flag is set.
+
+============ ===============
+Field name: ram_size
+Type: kernel internal
+Offset/size: 0x1f8/2
+Protocol: ALL
+============ ===============
+
+ This field is obsolete.
+
+============ ===================
+Field name: vid_mode
+Type: modify (obligatory)
+Offset/size: 0x1fa/2
+============ ===================
+
+ Please see the section on SPECIAL COMMAND LINE OPTIONS.
+
+============ =================
+Field name: root_dev
+Type: modify (optional)
+Offset/size: 0x1fc/2
+Protocol: ALL
+============ =================
+
+ The default root device device number. The use of this field is
+ deprecated, use the "root=" option on the command line instead.
+
+============ =========
+Field name: boot_flag
+Type: read
+Offset/size: 0x1fe/2
+Protocol: ALL
+============ =========
+
+ Contains 0xAA55. This is the closest thing old Linux kernels have
+ to a magic number.
+
+============ =======
+Field name: jump
+Type: read
+Offset/size: 0x200/2
+Protocol: 2.00+
+============ =======
+
+ Contains an x86 jump instruction, 0xEB followed by a signed offset
+ relative to byte 0x202. This can be used to determine the size of
+ the header.
+
+============ =======
+Field name: header
+Type: read
+Offset/size: 0x202/4
+Protocol: 2.00+
+============ =======
+
+ Contains the magic number "HdrS" (0x53726448).
+
+============ =======
+Field name: version
+Type: read
+Offset/size: 0x206/2
+Protocol: 2.00+
+============ =======
+
+ Contains the boot protocol version, in (major << 8)+minor format,
+ e.g. 0x0204 for version 2.04, and 0x0a11 for a hypothetical version
+ 10.17.
+
+============ =================
+Field name: realmode_swtch
+Type: modify (optional)
+Offset/size: 0x208/4
+Protocol: 2.00+
+============ =================
+
+ Boot loader hook (see ADVANCED BOOT LOADER HOOKS below.)
+
+============ =============
+Field name: start_sys_seg
+Type: read
+Offset/size: 0x20c/2
+Protocol: 2.00+
+============ =============
+
+ The load low segment (0x1000). Obsolete.
+
+============ ==============
+Field name: kernel_version
+Type: read
+Offset/size: 0x20e/2
+Protocol: 2.00+
+============ ==============
+
+ If set to a nonzero value, contains a pointer to a NUL-terminated
+ human-readable kernel version number string, less 0x200. This can
+ be used to display the kernel version to the user. This value
+ should be less than (0x200*setup_sects).
+
+ For example, if this value is set to 0x1c00, the kernel version
+ number string can be found at offset 0x1e00 in the kernel file.
+ This is a valid value if and only if the "setup_sects" field
+ contains the value 15 or higher, as::
+
+ 0x1c00 < 15*0x200 (= 0x1e00) but
+ 0x1c00 >= 14*0x200 (= 0x1c00)
+
+ 0x1c00 >> 9 = 14, So the minimum value for setup_secs is 15.
+
+============ ==================
+Field name: type_of_loader
+Type: write (obligatory)
+Offset/size: 0x210/1
+Protocol: 2.00+
+============ ==================
+
+ If your boot loader has an assigned id (see table below), enter
+ 0xTV here, where T is an identifier for the boot loader and V is
+ a version number. Otherwise, enter 0xFF here.
+
+ For boot loader IDs above T = 0xD, write T = 0xE to this field and
+ write the extended ID minus 0x10 to the ext_loader_type field.
+ Similarly, the ext_loader_ver field can be used to provide more than
+ four bits for the bootloader version.
+
+ For example, for T = 0x15, V = 0x234, write::
+
+ type_of_loader <- 0xE4
+ ext_loader_type <- 0x05
+ ext_loader_ver <- 0x23
+
+ Assigned boot loader ids (hexadecimal):
+
+ == =======================================
+ 0 LILO
+ (0x00 reserved for pre-2.00 bootloader)
+ 1 Loadlin
+ 2 bootsect-loader
+ (0x20, all other values reserved)
+ 3 Syslinux
+ 4 Etherboot/gPXE/iPXE
+ 5 ELILO
+ 7 GRUB
+ 8 U-Boot
+ 9 Xen
+ A Gujin
+ B Qemu
+ C Arcturus Networks uCbootloader
+ D kexec-tools
+ E Extended (see ext_loader_type)
+ F Special (0xFF = undefined)
+ 10 Reserved
+ 11 Minimal Linux Bootloader
+ <http://sebastian-plotz.blogspot.de>
+ 12 OVMF UEFI virtualization stack
+ == =======================================
+
+ Please contact <hpa@zytor.com> if you need a bootloader ID value assigned.
+
+============ ===================
+Field name: loadflags
+Type: modify (obligatory)
+Offset/size: 0x211/1
+Protocol: 2.00+
+============ ===================
+
+ This field is a bitmask.
+
+ Bit 0 (read): LOADED_HIGH
+
+ - If 0, the protected-mode code is loaded at 0x10000.
+ - If 1, the protected-mode code is loaded at 0x100000.
+
+ Bit 1 (kernel internal): KASLR_FLAG
+
+ - Used internally by the compressed kernel to communicate
+ KASLR status to kernel proper.
+
+ - If 1, KASLR enabled.
+ - If 0, KASLR disabled.
+
+ Bit 5 (write): QUIET_FLAG
+
+ - If 0, print early messages.
+ - If 1, suppress early messages.
+
+ This requests to the kernel (decompressor and early
+ kernel) to not write early messages that require
+ accessing the display hardware directly.
+
+ Bit 6 (obsolete): KEEP_SEGMENTS
+
+ Protocol: 2.07+
+
+ - This flag is obsolete.
+
+ Bit 7 (write): CAN_USE_HEAP
+
+ Set this bit to 1 to indicate that the value entered in the
+ heap_end_ptr is valid. If this field is clear, some setup code
+ functionality will be disabled.
+
+
+============ ===================
+Field name: setup_move_size
+Type: modify (obligatory)
+Offset/size: 0x212/2
+Protocol: 2.00-2.01
+============ ===================
+
+ When using protocol 2.00 or 2.01, if the real mode kernel is not
+ loaded at 0x90000, it gets moved there later in the loading
+ sequence. Fill in this field if you want additional data (such as
+ the kernel command line) moved in addition to the real-mode kernel
+ itself.
+
+ The unit is bytes starting with the beginning of the boot sector.
+
+ This field is can be ignored when the protocol is 2.02 or higher, or
+ if the real-mode code is loaded at 0x90000.
+
+============ ========================
+Field name: code32_start
+Type: modify (optional, reloc)
+Offset/size: 0x214/4
+Protocol: 2.00+
+============ ========================
+
+ The address to jump to in protected mode. This defaults to the load
+ address of the kernel, and can be used by the boot loader to
+ determine the proper load address.
+
+ This field can be modified for two purposes:
+
+ 1. as a boot loader hook (see Advanced Boot Loader Hooks below.)
+
+ 2. if a bootloader which does not install a hook loads a
+ relocatable kernel at a nonstandard address it will have to modify
+ this field to point to the load address.
+
+============ ==================
+Field name: ramdisk_image
+Type: write (obligatory)
+Offset/size: 0x218/4
+Protocol: 2.00+
+============ ==================
+
+ The 32-bit linear address of the initial ramdisk or ramfs. Leave at
+ zero if there is no initial ramdisk/ramfs.
+
+============ ==================
+Field name: ramdisk_size
+Type: write (obligatory)
+Offset/size: 0x21c/4
+Protocol: 2.00+
+============ ==================
+
+ Size of the initial ramdisk or ramfs. Leave at zero if there is no
+ initial ramdisk/ramfs.
+
+============ ===============
+Field name: bootsect_kludge
+Type: kernel internal
+Offset/size: 0x220/4
+Protocol: 2.00+
+============ ===============
+
+ This field is obsolete.
+
+============ ==================
+Field name: heap_end_ptr
+Type: write (obligatory)
+Offset/size: 0x224/2
+Protocol: 2.01+
+============ ==================
+
+ Set this field to the offset (from the beginning of the real-mode
+ code) of the end of the setup stack/heap, minus 0x0200.
+
+============ ================
+Field name: ext_loader_ver
+Type: write (optional)
+Offset/size: 0x226/1
+Protocol: 2.02+
+============ ================
+
+ This field is used as an extension of the version number in the
+ type_of_loader field. The total version number is considered to be
+ (type_of_loader & 0x0f) + (ext_loader_ver << 4).
+
+ The use of this field is boot loader specific. If not written, it
+ is zero.
+
+ Kernels prior to 2.6.31 did not recognize this field, but it is safe
+ to write for protocol version 2.02 or higher.
+
+============ =====================================================
+Field name: ext_loader_type
+Type: write (obligatory if (type_of_loader & 0xf0) == 0xe0)
+Offset/size: 0x227/1
+Protocol: 2.02+
+============ =====================================================
+
+ This field is used as an extension of the type number in
+ type_of_loader field. If the type in type_of_loader is 0xE, then
+ the actual type is (ext_loader_type + 0x10).
+
+ This field is ignored if the type in type_of_loader is not 0xE.
+
+ Kernels prior to 2.6.31 did not recognize this field, but it is safe
+ to write for protocol version 2.02 or higher.
+
+============ ==================
+Field name: cmd_line_ptr
+Type: write (obligatory)
+Offset/size: 0x228/4
+Protocol: 2.02+
+============ ==================
+
+ Set this field to the linear address of the kernel command line.
+ The kernel command line can be located anywhere between the end of
+ the setup heap and 0xA0000; it does not have to be located in the
+ same 64K segment as the real-mode code itself.
+
+ Fill in this field even if your boot loader does not support a
+ command line, in which case you can point this to an empty string
+ (or better yet, to the string "auto".) If this field is left at
+ zero, the kernel will assume that your boot loader does not support
+ the 2.02+ protocol.
+
+============ ===============
+Field name: initrd_addr_max
+Type: read
+Offset/size: 0x22c/4
+Protocol: 2.03+
+============ ===============
+
+ The maximum address that may be occupied by the initial
+ ramdisk/ramfs contents. For boot protocols 2.02 or earlier, this
+ field is not present, and the maximum address is 0x37FFFFFF. (This
+ address is defined as the address of the highest safe byte, so if
+ your ramdisk is exactly 131072 bytes long and this field is
+ 0x37FFFFFF, you can start your ramdisk at 0x37FE0000.)
+
+============ ============================
+Field name: kernel_alignment
+Type: read/modify (reloc)
+Offset/size: 0x230/4
+Protocol: 2.05+ (read), 2.10+ (modify)
+============ ============================
+
+ Alignment unit required by the kernel (if relocatable_kernel is
+ true.) A relocatable kernel that is loaded at an alignment
+ incompatible with the value in this field will be realigned during
+ kernel initialization.
+
+ Starting with protocol version 2.10, this reflects the kernel
+ alignment preferred for optimal performance; it is possible for the
+ loader to modify this field to permit a lesser alignment. See the
+ min_alignment and pref_address field below.
+
+============ ==================
+Field name: relocatable_kernel
+Type: read (reloc)
+Offset/size: 0x234/1
+Protocol: 2.05+
+============ ==================
+
+ If this field is nonzero, the protected-mode part of the kernel can
+ be loaded at any address that satisfies the kernel_alignment field.
+ After loading, the boot loader must set the code32_start field to
+ point to the loaded code, or to a boot loader hook.
+
+============ =============
+Field name: min_alignment
+Type: read (reloc)
+Offset/size: 0x235/1
+Protocol: 2.10+
+============ =============
+
+ This field, if nonzero, indicates as a power of two the minimum
+ alignment required, as opposed to preferred, by the kernel to boot.
+ If a boot loader makes use of this field, it should update the
+ kernel_alignment field with the alignment unit desired; typically::
+
+ kernel_alignment = 1 << min_alignment
+
+ There may be a considerable performance cost with an excessively
+ misaligned kernel. Therefore, a loader should typically try each
+ power-of-two alignment from kernel_alignment down to this alignment.
+
+============ ==========
+Field name: xloadflags
+Type: read
+Offset/size: 0x236/2
+Protocol: 2.12+
+============ ==========
+
+ This field is a bitmask.
+
+ Bit 0 (read): XLF_KERNEL_64
+
+ - If 1, this kernel has the legacy 64-bit entry point at 0x200.
+
+ Bit 1 (read): XLF_CAN_BE_LOADED_ABOVE_4G
+
+ - If 1, kernel/boot_params/cmdline/ramdisk can be above 4G.
+
+ Bit 2 (read): XLF_EFI_HANDOVER_32
+
+ - If 1, the kernel supports the 32-bit EFI handoff entry point
+ given at handover_offset.
+
+ Bit 3 (read): XLF_EFI_HANDOVER_64
+
+ - If 1, the kernel supports the 64-bit EFI handoff entry point
+ given at handover_offset + 0x200.
+
+ Bit 4 (read): XLF_EFI_KEXEC
+
+ - If 1, the kernel supports kexec EFI boot with EFI runtime support.
+
+
+============ ============
+Field name: cmdline_size
+Type: read
+Offset/size: 0x238/4
+Protocol: 2.06+
+============ ============
+
+ The maximum size of the command line without the terminating
+ zero. This means that the command line can contain at most
+ cmdline_size characters. With protocol version 2.05 and earlier, the
+ maximum size was 255.
+
+============ ====================================
+Field name: hardware_subarch
+Type: write (optional, defaults to x86/PC)
+Offset/size: 0x23c/4
+Protocol: 2.07+
+============ ====================================
+
+ In a paravirtualized environment the hardware low level architectural
+ pieces such as interrupt handling, page table handling, and
+ accessing process control registers needs to be done differently.
+
+ This field allows the bootloader to inform the kernel we are in one
+ one of those environments.
+
+ ========== ==============================
+ 0x00000000 The default x86/PC environment
+ 0x00000001 lguest
+ 0x00000002 Xen
+ 0x00000003 Moorestown MID
+ 0x00000004 CE4100 TV Platform
+ ========== ==============================
+
+============ =========================
+Field name: hardware_subarch_data
+Type: write (subarch-dependent)
+Offset/size: 0x240/8
+Protocol: 2.07+
+============ =========================
+
+ A pointer to data that is specific to hardware subarch
+ This field is currently unused for the default x86/PC environment,
+ do not modify.
+
+============ ==============
+Field name: payload_offset
+Type: read
+Offset/size: 0x248/4
+Protocol: 2.08+
+============ ==============
+
+ If non-zero then this field contains the offset from the beginning
+ of the protected-mode code to the payload.
+
+ The payload may be compressed. The format of both the compressed and
+ uncompressed data should be determined using the standard magic
+ numbers. The currently supported compression formats are gzip
+ (magic numbers 1F 8B or 1F 9E), bzip2 (magic number 42 5A), LZMA
+ (magic number 5D 00), XZ (magic number FD 37), LZ4 (magic number
+ 02 21) and ZSTD (magic number 28 B5). The uncompressed payload is
+ currently always ELF (magic number 7F 45 4C 46).
+
+============ ==============
+Field name: payload_length
+Type: read
+Offset/size: 0x24c/4
+Protocol: 2.08+
+============ ==============
+
+ The length of the payload.
+
+============ ===============
+Field name: setup_data
+Type: write (special)
+Offset/size: 0x250/8
+Protocol: 2.09+
+============ ===============
+
+ The 64-bit physical pointer to NULL terminated single linked list of
+ struct setup_data. This is used to define a more extensible boot
+ parameters passing mechanism. The definition of struct setup_data is
+ as follow::
+
+ struct setup_data {
+ u64 next;
+ u32 type;
+ u32 len;
+ u8 data[0];
+ };
+
+ Where, the next is a 64-bit physical pointer to the next node of
+ linked list, the next field of the last node is 0; the type is used
+ to identify the contents of data; the len is the length of data
+ field; the data holds the real payload.
+
+ This list may be modified at a number of points during the bootup
+ process. Therefore, when modifying this list one should always make
+ sure to consider the case where the linked list already contains
+ entries.
+
+ The setup_data is a bit awkward to use for extremely large data objects,
+ both because the setup_data header has to be adjacent to the data object
+ and because it has a 32-bit length field. However, it is important that
+ intermediate stages of the boot process have a way to identify which
+ chunks of memory are occupied by kernel data.
+
+ Thus setup_indirect struct and SETUP_INDIRECT type were introduced in
+ protocol 2.15::
+
+ struct setup_indirect {
+ __u32 type;
+ __u32 reserved; /* Reserved, must be set to zero. */
+ __u64 len;
+ __u64 addr;
+ };
+
+ The type member is a SETUP_INDIRECT | SETUP_* type. However, it cannot be
+ SETUP_INDIRECT itself since making the setup_indirect a tree structure
+ could require a lot of stack space in something that needs to parse it
+ and stack space can be limited in boot contexts.
+
+ Let's give an example how to point to SETUP_E820_EXT data using setup_indirect.
+ In this case setup_data and setup_indirect will look like this::
+
+ struct setup_data {
+ __u64 next = 0 or <addr_of_next_setup_data_struct>;
+ __u32 type = SETUP_INDIRECT;
+ __u32 len = sizeof(setup_indirect);
+ __u8 data[sizeof(setup_indirect)] = struct setup_indirect {
+ __u32 type = SETUP_INDIRECT | SETUP_E820_EXT;
+ __u32 reserved = 0;
+ __u64 len = <len_of_SETUP_E820_EXT_data>;
+ __u64 addr = <addr_of_SETUP_E820_EXT_data>;
+ }
+ }
+
+.. note::
+ SETUP_INDIRECT | SETUP_NONE objects cannot be properly distinguished
+ from SETUP_INDIRECT itself. So, this kind of objects cannot be provided
+ by the bootloaders.
+
+============ ============
+Field name: pref_address
+Type: read (reloc)
+Offset/size: 0x258/8
+Protocol: 2.10+
+============ ============
+
+ This field, if nonzero, represents a preferred load address for the
+ kernel. A relocating bootloader should attempt to load at this
+ address if possible.
+
+ A non-relocatable kernel will unconditionally move itself and to run
+ at this address.
+
+============ =======
+Field name: init_size
+Type: read
+Offset/size: 0x260/4
+============ =======
+
+ This field indicates the amount of linear contiguous memory starting
+ at the kernel runtime start address that the kernel needs before it
+ is capable of examining its memory map. This is not the same thing
+ as the total amount of memory the kernel needs to boot, but it can
+ be used by a relocating boot loader to help select a safe load
+ address for the kernel.
+
+ The kernel runtime start address is determined by the following algorithm::
+
+ if (relocatable_kernel)
+ runtime_start = align_up(load_address, kernel_alignment)
+ else
+ runtime_start = pref_address
+
+============ ===============
+Field name: handover_offset
+Type: read
+Offset/size: 0x264/4
+============ ===============
+
+ This field is the offset from the beginning of the kernel image to
+ the EFI handover protocol entry point. Boot loaders using the EFI
+ handover protocol to boot the kernel should jump to this offset.
+
+ See EFI HANDOVER PROTOCOL below for more details.
+
+============ ==================
+Field name: kernel_info_offset
+Type: read
+Offset/size: 0x268/4
+Protocol: 2.15+
+============ ==================
+
+ This field is the offset from the beginning of the kernel image to the
+ kernel_info. The kernel_info structure is embedded in the Linux image
+ in the uncompressed protected mode region.
+
+
+The kernel_info
+===============
+
+The relationships between the headers are analogous to the various data
+sections:
+
+ setup_header = .data
+ boot_params/setup_data = .bss
+
+What is missing from the above list? That's right:
+
+ kernel_info = .rodata
+
+We have been (ab)using .data for things that could go into .rodata or .bss for
+a long time, for lack of alternatives and -- especially early on -- inertia.
+Also, the BIOS stub is responsible for creating boot_params, so it isn't
+available to a BIOS-based loader (setup_data is, though).
+
+setup_header is permanently limited to 144 bytes due to the reach of the
+2-byte jump field, which doubles as a length field for the structure, combined
+with the size of the "hole" in struct boot_params that a protected-mode loader
+or the BIOS stub has to copy it into. It is currently 119 bytes long, which
+leaves us with 25 very precious bytes. This isn't something that can be fixed
+without revising the boot protocol entirely, breaking backwards compatibility.
+
+boot_params proper is limited to 4096 bytes, but can be arbitrarily extended
+by adding setup_data entries. It cannot be used to communicate properties of
+the kernel image, because it is .bss and has no image-provided content.
+
+kernel_info solves this by providing an extensible place for information about
+the kernel image. It is readonly, because the kernel cannot rely on a
+bootloader copying its contents anywhere, but that is OK; if it becomes
+necessary it can still contain data items that an enabled bootloader would be
+expected to copy into a setup_data chunk.
+
+All kernel_info data should be part of this structure. Fixed size data have to
+be put before kernel_info_var_len_data label. Variable size data have to be put
+after kernel_info_var_len_data label. Each chunk of variable size data has to
+be prefixed with header/magic and its size, e.g.::
+
+ kernel_info:
+ .ascii "LToP" /* Header, Linux top (structure). */
+ .long kernel_info_var_len_data - kernel_info
+ .long kernel_info_end - kernel_info
+ .long 0x01234567 /* Some fixed size data for the bootloaders. */
+ kernel_info_var_len_data:
+ example_struct: /* Some variable size data for the bootloaders. */
+ .ascii "0123" /* Header/Magic. */
+ .long example_struct_end - example_struct
+ .ascii "Struct"
+ .long 0x89012345
+ example_struct_end:
+ example_strings: /* Some variable size data for the bootloaders. */
+ .ascii "ABCD" /* Header/Magic. */
+ .long example_strings_end - example_strings
+ .asciz "String_0"
+ .asciz "String_1"
+ example_strings_end:
+ kernel_info_end:
+
+This way the kernel_info is self-contained blob.
+
+.. note::
+ Each variable size data header/magic can be any 4-character string,
+ without \0 at the end of the string, which does not collide with
+ existing variable length data headers/magics.
+
+
+Details of the kernel_info Fields
+=================================
+
+============ ========
+Field name: header
+Offset/size: 0x0000/4
+============ ========
+
+ Contains the magic number "LToP" (0x506f544c).
+
+============ ========
+Field name: size
+Offset/size: 0x0004/4
+============ ========
+
+ This field contains the size of the kernel_info including kernel_info.header.
+ It does not count kernel_info.kernel_info_var_len_data size. This field should be
+ used by the bootloaders to detect supported fixed size fields in the kernel_info
+ and beginning of kernel_info.kernel_info_var_len_data.
+
+============ ========
+Field name: size_total
+Offset/size: 0x0008/4
+============ ========
+
+ This field contains the size of the kernel_info including kernel_info.header
+ and kernel_info.kernel_info_var_len_data.
+
+============ ==============
+Field name: setup_type_max
+Offset/size: 0x000c/4
+============ ==============
+
+ This field contains maximal allowed type for setup_data and setup_indirect structs.
+
+
+The Image Checksum
+==================
+
+From boot protocol version 2.08 onwards the CRC-32 is calculated over
+the entire file using the characteristic polynomial 0x04C11DB7 and an
+initial remainder of 0xffffffff. The checksum is appended to the
+file; therefore the CRC of the file up to the limit specified in the
+syssize field of the header is always 0.
+
+
+The Kernel Command Line
+=======================
+
+The kernel command line has become an important way for the boot
+loader to communicate with the kernel. Some of its options are also
+relevant to the boot loader itself, see "special command line options"
+below.
+
+The kernel command line is a null-terminated string. The maximum
+length can be retrieved from the field cmdline_size. Before protocol
+version 2.06, the maximum was 255 characters. A string that is too
+long will be automatically truncated by the kernel.
+
+If the boot protocol version is 2.02 or later, the address of the
+kernel command line is given by the header field cmd_line_ptr (see
+above.) This address can be anywhere between the end of the setup
+heap and 0xA0000.
+
+If the protocol version is *not* 2.02 or higher, the kernel
+command line is entered using the following protocol:
+
+ - At offset 0x0020 (word), "cmd_line_magic", enter the magic
+ number 0xA33F.
+
+ - At offset 0x0022 (word), "cmd_line_offset", enter the offset
+ of the kernel command line (relative to the start of the
+ real-mode kernel).
+
+ - The kernel command line *must* be within the memory region
+ covered by setup_move_size, so you may need to adjust this
+ field.
+
+
+Memory Layout of The Real-Mode Code
+===================================
+
+The real-mode code requires a stack/heap to be set up, as well as
+memory allocated for the kernel command line. This needs to be done
+in the real-mode accessible memory in bottom megabyte.
+
+It should be noted that modern machines often have a sizable Extended
+BIOS Data Area (EBDA). As a result, it is advisable to use as little
+of the low megabyte as possible.
+
+Unfortunately, under the following circumstances the 0x90000 memory
+segment has to be used:
+
+ - When loading a zImage kernel ((loadflags & 0x01) == 0).
+ - When loading a 2.01 or earlier boot protocol kernel.
+
+.. note::
+ For the 2.00 and 2.01 boot protocols, the real-mode code
+ can be loaded at another address, but it is internally
+ relocated to 0x90000. For the "old" protocol, the
+ real-mode code must be loaded at 0x90000.
+
+When loading at 0x90000, avoid using memory above 0x9a000.
+
+For boot protocol 2.02 or higher, the command line does not have to be
+located in the same 64K segment as the real-mode setup code; it is
+thus permitted to give the stack/heap the full 64K segment and locate
+the command line above it.
+
+The kernel command line should not be located below the real-mode
+code, nor should it be located in high memory.
+
+
+Sample Boot Configuartion
+=========================
+
+As a sample configuration, assume the following layout of the real
+mode segment.
+
+ When loading below 0x90000, use the entire segment:
+
+ ============= ===================
+ 0x0000-0x7fff Real mode kernel
+ 0x8000-0xdfff Stack and heap
+ 0xe000-0xffff Kernel command line
+ ============= ===================
+
+ When loading at 0x90000 OR the protocol version is 2.01 or earlier:
+
+ ============= ===================
+ 0x0000-0x7fff Real mode kernel
+ 0x8000-0x97ff Stack and heap
+ 0x9800-0x9fff Kernel command line
+ ============= ===================
+
+Such a boot loader should enter the following fields in the header::
+
+ unsigned long base_ptr; /* base address for real-mode segment */
+
+ if ( setup_sects == 0 ) {
+ setup_sects = 4;
+ }
+
+ if ( protocol >= 0x0200 ) {
+ type_of_loader = <type code>;
+ if ( loading_initrd ) {
+ ramdisk_image = <initrd_address>;
+ ramdisk_size = <initrd_size>;
+ }
+
+ if ( protocol >= 0x0202 && loadflags & 0x01 )
+ heap_end = 0xe000;
+ else
+ heap_end = 0x9800;
+
+ if ( protocol >= 0x0201 ) {
+ heap_end_ptr = heap_end - 0x200;
+ loadflags |= 0x80; /* CAN_USE_HEAP */
+ }
+
+ if ( protocol >= 0x0202 ) {
+ cmd_line_ptr = base_ptr + heap_end;
+ strcpy(cmd_line_ptr, cmdline);
+ } else {
+ cmd_line_magic = 0xA33F;
+ cmd_line_offset = heap_end;
+ setup_move_size = heap_end + strlen(cmdline)+1;
+ strcpy(base_ptr+cmd_line_offset, cmdline);
+ }
+ } else {
+ /* Very old kernel */
+
+ heap_end = 0x9800;
+
+ cmd_line_magic = 0xA33F;
+ cmd_line_offset = heap_end;
+
+ /* A very old kernel MUST have its real-mode code
+ loaded at 0x90000 */
+
+ if ( base_ptr != 0x90000 ) {
+ /* Copy the real-mode kernel */
+ memcpy(0x90000, base_ptr, (setup_sects+1)*512);
+ base_ptr = 0x90000; /* Relocated */
+ }
+
+ strcpy(0x90000+cmd_line_offset, cmdline);
+
+ /* It is recommended to clear memory up to the 32K mark */
+ memset(0x90000 + (setup_sects+1)*512, 0,
+ (64-(setup_sects+1))*512);
+ }
+
+
+Loading The Rest of The Kernel
+==============================
+
+The 32-bit (non-real-mode) kernel starts at offset (setup_sects+1)*512
+in the kernel file (again, if setup_sects == 0 the real value is 4.)
+It should be loaded at address 0x10000 for Image/zImage kernels and
+0x100000 for bzImage kernels.
+
+The kernel is a bzImage kernel if the protocol >= 2.00 and the 0x01
+bit (LOAD_HIGH) in the loadflags field is set::
+
+ is_bzImage = (protocol >= 0x0200) && (loadflags & 0x01);
+ load_address = is_bzImage ? 0x100000 : 0x10000;
+
+Note that Image/zImage kernels can be up to 512K in size, and thus use
+the entire 0x10000-0x90000 range of memory. This means it is pretty
+much a requirement for these kernels to load the real-mode part at
+0x90000. bzImage kernels allow much more flexibility.
+
+Special Command Line Options
+============================
+
+If the command line provided by the boot loader is entered by the
+user, the user may expect the following command line options to work.
+They should normally not be deleted from the kernel command line even
+though not all of them are actually meaningful to the kernel. Boot
+loader authors who need additional command line options for the boot
+loader itself should get them registered in
+Documentation/admin-guide/kernel-parameters.rst to make sure they will not
+conflict with actual kernel options now or in the future.
+
+ vga=<mode>
+ <mode> here is either an integer (in C notation, either
+ decimal, octal, or hexadecimal) or one of the strings
+ "normal" (meaning 0xFFFF), "ext" (meaning 0xFFFE) or "ask"
+ (meaning 0xFFFD). This value should be entered into the
+ vid_mode field, as it is used by the kernel before the command
+ line is parsed.
+
+ mem=<size>
+ <size> is an integer in C notation optionally followed by
+ (case insensitive) K, M, G, T, P or E (meaning << 10, << 20,
+ << 30, << 40, << 50 or << 60). This specifies the end of
+ memory to the kernel. This affects the possible placement of
+ an initrd, since an initrd should be placed near end of
+ memory. Note that this is an option to *both* the kernel and
+ the bootloader!
+
+ initrd=<file>
+ An initrd should be loaded. The meaning of <file> is
+ obviously bootloader-dependent, and some boot loaders
+ (e.g. LILO) do not have such a command.
+
+In addition, some boot loaders add the following options to the
+user-specified command line:
+
+ BOOT_IMAGE=<file>
+ The boot image which was loaded. Again, the meaning of <file>
+ is obviously bootloader-dependent.
+
+ auto
+ The kernel was booted without explicit user intervention.
+
+If these options are added by the boot loader, it is highly
+recommended that they are located *first*, before the user-specified
+or configuration-specified command line. Otherwise, "init=/bin/sh"
+gets confused by the "auto" option.
+
+
+Running the Kernel
+==================
+
+The kernel is started by jumping to the kernel entry point, which is
+located at *segment* offset 0x20 from the start of the real mode
+kernel. This means that if you loaded your real-mode kernel code at
+0x90000, the kernel entry point is 9020:0000.
+
+At entry, ds = es = ss should point to the start of the real-mode
+kernel code (0x9000 if the code is loaded at 0x90000), sp should be
+set up properly, normally pointing to the top of the heap, and
+interrupts should be disabled. Furthermore, to guard against bugs in
+the kernel, it is recommended that the boot loader sets fs = gs = ds =
+es = ss.
+
+In our example from above, we would do::
+
+ /* Note: in the case of the "old" kernel protocol, base_ptr must
+ be == 0x90000 at this point; see the previous sample code */
+
+ seg = base_ptr >> 4;
+
+ cli(); /* Enter with interrupts disabled! */
+
+ /* Set up the real-mode kernel stack */
+ _SS = seg;
+ _SP = heap_end;
+
+ _DS = _ES = _FS = _GS = seg;
+ jmp_far(seg+0x20, 0); /* Run the kernel */
+
+If your boot sector accesses a floppy drive, it is recommended to
+switch off the floppy motor before running the kernel, since the
+kernel boot leaves interrupts off and thus the motor will not be
+switched off, especially if the loaded kernel has the floppy driver as
+a demand-loaded module!
+
+
+Advanced Boot Loader Hooks
+==========================
+
+If the boot loader runs in a particularly hostile environment (such as
+LOADLIN, which runs under DOS) it may be impossible to follow the
+standard memory location requirements. Such a boot loader may use the
+following hooks that, if set, are invoked by the kernel at the
+appropriate time. The use of these hooks should probably be
+considered an absolutely last resort!
+
+IMPORTANT: All the hooks are required to preserve %esp, %ebp, %esi and
+%edi across invocation.
+
+ realmode_swtch:
+ A 16-bit real mode far subroutine invoked immediately before
+ entering protected mode. The default routine disables NMI, so
+ your routine should probably do so, too.
+
+ code32_start:
+ A 32-bit flat-mode routine *jumped* to immediately after the
+ transition to protected mode, but before the kernel is
+ uncompressed. No segments, except CS, are guaranteed to be
+ set up (current kernels do, but older ones do not); you should
+ set them up to BOOT_DS (0x18) yourself.
+
+ After completing your hook, you should jump to the address
+ that was in this field before your boot loader overwrote it
+ (relocated, if appropriate.)
+
+
+32-bit Boot Protocol
+====================
+
+For machine with some new BIOS other than legacy BIOS, such as EFI,
+LinuxBIOS, etc, and kexec, the 16-bit real mode setup code in kernel
+based on legacy BIOS can not be used, so a 32-bit boot protocol needs
+to be defined.
+
+In 32-bit boot protocol, the first step in loading a Linux kernel
+should be to setup the boot parameters (struct boot_params,
+traditionally known as "zero page"). The memory for struct boot_params
+should be allocated and initialized to all zero. Then the setup header
+from offset 0x01f1 of kernel image on should be loaded into struct
+boot_params and examined. The end of setup header can be calculated as
+follow::
+
+ 0x0202 + byte value at offset 0x0201
+
+In addition to read/modify/write the setup header of the struct
+boot_params as that of 16-bit boot protocol, the boot loader should
+also fill the additional fields of the struct boot_params as
+described in chapter Documentation/x86/zero-page.rst.
+
+After setting up the struct boot_params, the boot loader can load the
+32/64-bit kernel in the same way as that of 16-bit boot protocol.
+
+In 32-bit boot protocol, the kernel is started by jumping to the
+32-bit kernel entry point, which is the start address of loaded
+32/64-bit kernel.
+
+At entry, the CPU must be in 32-bit protected mode with paging
+disabled; a GDT must be loaded with the descriptors for selectors
+__BOOT_CS(0x10) and __BOOT_DS(0x18); both descriptors must be 4G flat
+segment; __BOOT_CS must have execute/read permission, and __BOOT_DS
+must have read/write permission; CS must be __BOOT_CS and DS, ES, SS
+must be __BOOT_DS; interrupt must be disabled; %esi must hold the base
+address of the struct boot_params; %ebp, %edi and %ebx must be zero.
+
+64-bit Boot Protocol
+====================
+
+For machine with 64bit cpus and 64bit kernel, we could use 64bit bootloader
+and we need a 64-bit boot protocol.
+
+In 64-bit boot protocol, the first step in loading a Linux kernel
+should be to setup the boot parameters (struct boot_params,
+traditionally known as "zero page"). The memory for struct boot_params
+could be allocated anywhere (even above 4G) and initialized to all zero.
+Then, the setup header at offset 0x01f1 of kernel image on should be
+loaded into struct boot_params and examined. The end of setup header
+can be calculated as follows::
+
+ 0x0202 + byte value at offset 0x0201
+
+In addition to read/modify/write the setup header of the struct
+boot_params as that of 16-bit boot protocol, the boot loader should
+also fill the additional fields of the struct boot_params as described
+in chapter Documentation/x86/zero-page.rst.
+
+After setting up the struct boot_params, the boot loader can load
+64-bit kernel in the same way as that of 16-bit boot protocol, but
+kernel could be loaded above 4G.
+
+In 64-bit boot protocol, the kernel is started by jumping to the
+64-bit kernel entry point, which is the start address of loaded
+64-bit kernel plus 0x200.
+
+At entry, the CPU must be in 64-bit mode with paging enabled.
+The range with setup_header.init_size from start address of loaded
+kernel and zero page and command line buffer get ident mapping;
+a GDT must be loaded with the descriptors for selectors
+__BOOT_CS(0x10) and __BOOT_DS(0x18); both descriptors must be 4G flat
+segment; __BOOT_CS must have execute/read permission, and __BOOT_DS
+must have read/write permission; CS must be __BOOT_CS and DS, ES, SS
+must be __BOOT_DS; interrupt must be disabled; %rsi must hold the base
+address of the struct boot_params.
+
+EFI Handover Protocol (deprecated)
+==================================
+
+This protocol allows boot loaders to defer initialisation to the EFI
+boot stub. The boot loader is required to load the kernel/initrd(s)
+from the boot media and jump to the EFI handover protocol entry point
+which is hdr->handover_offset bytes from the beginning of
+startup_{32,64}.
+
+The boot loader MUST respect the kernel's PE/COFF metadata when it comes
+to section alignment, the memory footprint of the executable image beyond
+the size of the file itself, and any other aspect of the PE/COFF header
+that may affect correct operation of the image as a PE/COFF binary in the
+execution context provided by the EFI firmware.
+
+The function prototype for the handover entry point looks like this::
+
+ efi_main(void *handle, efi_system_table_t *table, struct boot_params *bp)
+
+'handle' is the EFI image handle passed to the boot loader by the EFI
+firmware, 'table' is the EFI system table - these are the first two
+arguments of the "handoff state" as described in section 2.3 of the
+UEFI specification. 'bp' is the boot loader-allocated boot params.
+
+The boot loader *must* fill out the following fields in bp::
+
+ - hdr.cmd_line_ptr
+ - hdr.ramdisk_image (if applicable)
+ - hdr.ramdisk_size (if applicable)
+
+All other fields should be zero.
+
+NOTE: The EFI Handover Protocol is deprecated in favour of the ordinary PE/COFF
+ entry point, combined with the LINUX_EFI_INITRD_MEDIA_GUID based initrd
+ loading protocol (refer to [0] for an example of the bootloader side of
+ this), which removes the need for any knowledge on the part of the EFI
+ bootloader regarding the internal representation of boot_params or any
+ requirements/limitations regarding the placement of the command line
+ and ramdisk in memory, or the placement of the kernel image itself.
+
+[0] https://github.com/u-boot/u-boot/commit/ec80b4735a593961fe701cc3a5d717d4739b0fd0
diff --git a/Documentation/x86/booting-dt.rst b/Documentation/x86/booting-dt.rst
new file mode 100644
index 000000000..965a37407
--- /dev/null
+++ b/Documentation/x86/booting-dt.rst
@@ -0,0 +1,21 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+DeviceTree Booting
+------------------
+
+ There is one single 32bit entry point to the kernel at code32_start,
+ the decompressor (the real mode entry point goes to the same 32bit
+ entry point once it switched into protected mode). That entry point
+ supports one calling convention which is documented in
+ Documentation/x86/boot.rst
+ The physical pointer to the device-tree block is passed via setup_data
+ which requires at least boot protocol 2.09.
+ The type filed is defined as
+
+ #define SETUP_DTB 2
+
+ This device-tree is used as an extension to the "boot page". As such it
+ does not parse / consider data which is already covered by the boot
+ page. This includes memory size, reserved ranges, command line arguments
+ or initrd address. It simply holds information which can not be retrieved
+ otherwise like interrupt routing or a list of devices behind an I2C bus.
diff --git a/Documentation/x86/buslock.rst b/Documentation/x86/buslock.rst
new file mode 100644
index 000000000..7c051e714
--- /dev/null
+++ b/Documentation/x86/buslock.rst
@@ -0,0 +1,126 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+.. include:: <isonum.txt>
+
+===============================
+Bus lock detection and handling
+===============================
+
+:Copyright: |copy| 2021 Intel Corporation
+:Authors: - Fenghua Yu <fenghua.yu@intel.com>
+ - Tony Luck <tony.luck@intel.com>
+
+Problem
+=======
+
+A split lock is any atomic operation whose operand crosses two cache lines.
+Since the operand spans two cache lines and the operation must be atomic,
+the system locks the bus while the CPU accesses the two cache lines.
+
+A bus lock is acquired through either split locked access to writeback (WB)
+memory or any locked access to non-WB memory. This is typically thousands of
+cycles slower than an atomic operation within a cache line. It also disrupts
+performance on other cores and brings the whole system to its knees.
+
+Detection
+=========
+
+Intel processors may support either or both of the following hardware
+mechanisms to detect split locks and bus locks.
+
+#AC exception for split lock detection
+--------------------------------------
+
+Beginning with the Tremont Atom CPU split lock operations may raise an
+Alignment Check (#AC) exception when a split lock operation is attemped.
+
+#DB exception for bus lock detection
+------------------------------------
+
+Some CPUs have the ability to notify the kernel by an #DB trap after a user
+instruction acquires a bus lock and is executed. This allows the kernel to
+terminate the application or to enforce throttling.
+
+Software handling
+=================
+
+The kernel #AC and #DB handlers handle bus lock based on the kernel
+parameter "split_lock_detect". Here is a summary of different options:
+
++------------------+----------------------------+-----------------------+
+|split_lock_detect=|#AC for split lock |#DB for bus lock |
++------------------+----------------------------+-----------------------+
+|off |Do nothing |Do nothing |
++------------------+----------------------------+-----------------------+
+|warn |Kernel OOPs |Warn once per task and |
+|(default) |Warn once per task and |and continues to run. |
+| |disable future checking | |
+| |When both features are | |
+| |supported, warn in #AC | |
++------------------+----------------------------+-----------------------+
+|fatal |Kernel OOPs |Send SIGBUS to user. |
+| |Send SIGBUS to user | |
+| |When both features are | |
+| |supported, fatal in #AC | |
++------------------+----------------------------+-----------------------+
+|ratelimit:N |Do nothing |Limit bus lock rate to |
+|(0 < N <= 1000) | |N bus locks per second |
+| | |system wide and warn on|
+| | |bus locks. |
++------------------+----------------------------+-----------------------+
+
+Usages
+======
+
+Detecting and handling bus lock may find usages in various areas:
+
+It is critical for real time system designers who build consolidated real
+time systems. These systems run hard real time code on some cores and run
+"untrusted" user processes on other cores. The hard real time cannot afford
+to have any bus lock from the untrusted processes to hurt real time
+performance. To date the designers have been unable to deploy these
+solutions as they have no way to prevent the "untrusted" user code from
+generating split lock and bus lock to block the hard real time code to
+access memory during bus locking.
+
+It's also useful for general computing to prevent guests or user
+applications from slowing down the overall system by executing instructions
+with bus lock.
+
+
+Guidance
+========
+off
+---
+
+Disable checking for split lock and bus lock. This option can be useful if
+there are legacy applications that trigger these events at a low rate so
+that mitigation is not needed.
+
+warn
+----
+
+A warning is emitted when a bus lock is detected which allows to identify
+the offending application. This is the default behavior.
+
+fatal
+-----
+
+In this case, the bus lock is not tolerated and the process is killed.
+
+ratelimit
+---------
+
+A system wide bus lock rate limit N is specified where 0 < N <= 1000. This
+allows a bus lock rate up to N bus locks per second. When the bus lock rate
+is exceeded then any task which is caught via the buslock #DB exception is
+throttled by enforced sleeps until the rate goes under the limit again.
+
+This is an effective mitigation in cases where a minimal impact can be
+tolerated, but an eventual Denial of Service attack has to be prevented. It
+allows to identify the offending processes and analyze whether they are
+malicious or just badly written.
+
+Selecting a rate limit of 1000 allows the bus to be locked for up to about
+seven million cycles each second (assuming 7000 cycles for each bus
+lock). On a 2 GHz processor that would be about 0.35% system slowdown.
diff --git a/Documentation/x86/cpuinfo.rst b/Documentation/x86/cpuinfo.rst
new file mode 100644
index 000000000..08246e8ac
--- /dev/null
+++ b/Documentation/x86/cpuinfo.rst
@@ -0,0 +1,154 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=================
+x86 Feature Flags
+=================
+
+Introduction
+============
+
+On x86, flags appearing in /proc/cpuinfo have an X86_FEATURE definition
+in arch/x86/include/asm/cpufeatures.h. If the kernel cares about a feature
+or KVM want to expose the feature to a KVM guest, it can and should have
+an X86_FEATURE_* defined. These flags represent hardware features as
+well as software features.
+
+If users want to know if a feature is available on a given system, they
+try to find the flag in /proc/cpuinfo. If a given flag is present, it
+means that the kernel supports it and is currently making it available.
+If such flag represents a hardware feature, it also means that the
+hardware supports it.
+
+If the expected flag does not appear in /proc/cpuinfo, things are murkier.
+Users need to find out the reason why the flag is missing and find the way
+how to enable it, which is not always easy. There are several factors that
+can explain missing flags: the expected feature failed to enable, the feature
+is missing in hardware, platform firmware did not enable it, the feature is
+disabled at build or run time, an old kernel is in use, or the kernel does
+not support the feature and thus has not enabled it. In general, /proc/cpuinfo
+shows features which the kernel supports. For a full list of CPUID flags
+which the CPU supports, use tools/arch/x86/kcpuid.
+
+How are feature flags created?
+==============================
+
+a: Feature flags can be derived from the contents of CPUID leaves.
+------------------------------------------------------------------
+These feature definitions are organized mirroring the layout of CPUID
+leaves and grouped in words with offsets as mapped in enum cpuid_leafs
+in cpufeatures.h (see arch/x86/include/asm/cpufeatures.h for details).
+If a feature is defined with a X86_FEATURE_<name> definition in
+cpufeatures.h, and if it is detected at run time, the flags will be
+displayed accordingly in /proc/cpuinfo. For example, the flag "avx2"
+comes from X86_FEATURE_AVX2 in cpufeatures.h.
+
+b: Flags can be from scattered CPUID-based features.
+----------------------------------------------------
+Hardware features enumerated in sparsely populated CPUID leaves get
+software-defined values. Still, CPUID needs to be queried to determine
+if a given feature is present. This is done in init_scattered_cpuid_features().
+For instance, X86_FEATURE_CQM_LLC is defined as 11*32 + 0 and its presence is
+checked at runtime in the respective CPUID leaf [EAX=f, ECX=0] bit EDX[1].
+
+The intent of scattering CPUID leaves is to not bloat struct
+cpuinfo_x86.x86_capability[] unnecessarily. For instance, the CPUID leaf
+[EAX=7, ECX=0] has 30 features and is dense, but the CPUID leaf [EAX=7, EAX=1]
+has only one feature and would waste 31 bits of space in the x86_capability[]
+array. Since there is a struct cpuinfo_x86 for each possible CPU, the wasted
+memory is not trivial.
+
+c: Flags can be created synthetically under certain conditions for hardware features.
+-------------------------------------------------------------------------------------
+Examples of conditions include whether certain features are present in
+MSR_IA32_CORE_CAPS or specific CPU models are identified. If the needed
+conditions are met, the features are enabled by the set_cpu_cap or
+setup_force_cpu_cap macros. For example, if bit 5 is set in MSR_IA32_CORE_CAPS,
+the feature X86_FEATURE_SPLIT_LOCK_DETECT will be enabled and
+"split_lock_detect" will be displayed. The flag "ring3mwait" will be
+displayed only when running on INTEL_FAM6_XEON_PHI_[KNL|KNM] processors.
+
+d: Flags can represent purely software features.
+------------------------------------------------
+These flags do not represent hardware features. Instead, they represent a
+software feature implemented in the kernel. For example, Kernel Page Table
+Isolation is purely software feature and its feature flag X86_FEATURE_PTI is
+also defined in cpufeatures.h.
+
+Naming of Flags
+===============
+
+The script arch/x86/kernel/cpu/mkcapflags.sh processes the
+#define X86_FEATURE_<name> from cpufeatures.h and generates the
+x86_cap/bug_flags[] arrays in kernel/cpu/capflags.c. The names in the
+resulting x86_cap/bug_flags[] are used to populate /proc/cpuinfo. The naming
+of flags in the x86_cap/bug_flags[] are as follows:
+
+a: The name of the flag is from the string in X86_FEATURE_<name> by default.
+----------------------------------------------------------------------------
+By default, the flag <name> in /proc/cpuinfo is extracted from the respective
+X86_FEATURE_<name> in cpufeatures.h. For example, the flag "avx2" is from
+X86_FEATURE_AVX2.
+
+b: The naming can be overridden.
+--------------------------------
+If the comment on the line for the #define X86_FEATURE_* starts with a
+double-quote character (""), the string inside the double-quote characters
+will be the name of the flags. For example, the flag "sse4_1" comes from
+the comment "sse4_1" following the X86_FEATURE_XMM4_1 definition.
+
+There are situations in which overriding the displayed name of the flag is
+needed. For instance, /proc/cpuinfo is a userspace interface and must remain
+constant. If, for some reason, the naming of X86_FEATURE_<name> changes, one
+shall override the new naming with the name already used in /proc/cpuinfo.
+
+c: The naming override can be "", which means it will not appear in /proc/cpuinfo.
+----------------------------------------------------------------------------------
+The feature shall be omitted from /proc/cpuinfo if it does not make sense for
+the feature to be exposed to userspace. For example, X86_FEATURE_ALWAYS is
+defined in cpufeatures.h but that flag is an internal kernel feature used
+in the alternative runtime patching functionality. So, its name is overridden
+with "". Its flag will not appear in /proc/cpuinfo.
+
+Flags are missing when one or more of these happen
+==================================================
+
+a: The hardware does not enumerate support for it.
+--------------------------------------------------
+For example, when a new kernel is running on old hardware or the feature is
+not enabled by boot firmware. Even if the hardware is new, there might be a
+problem enabling the feature at run time, the flag will not be displayed.
+
+b: The kernel does not know about the flag.
+-------------------------------------------
+For example, when an old kernel is running on new hardware.
+
+c: The kernel disabled support for it at compile-time.
+------------------------------------------------------
+For example, if 5-level-paging is not enabled when building (i.e.,
+CONFIG_X86_5LEVEL is not selected) the flag "la57" will not show up [#f1]_.
+Even though the feature will still be detected via CPUID, the kernel disables
+it by clearing via setup_clear_cpu_cap(X86_FEATURE_LA57).
+
+d: The feature is disabled at boot-time.
+----------------------------------------
+A feature can be disabled either using a command-line parameter or because
+it failed to be enabled. The command-line parameter clearcpuid= can be used
+to disable features using the feature number as defined in
+/arch/x86/include/asm/cpufeatures.h. For instance, User Mode Instruction
+Protection can be disabled using clearcpuid=514. The number 514 is calculated
+from #define X86_FEATURE_UMIP (16*32 + 2).
+
+In addition, there exists a variety of custom command-line parameters that
+disable specific features. The list of parameters includes, but is not limited
+to, nofsgsbase, nosgx, noxsave, etc. 5-level paging can also be disabled using
+"no5lvl".
+
+e: The feature was known to be non-functional.
+----------------------------------------------
+The feature was known to be non-functional because a dependency was
+missing at runtime. For example, AVX flags will not show up if XSAVE feature
+is disabled since they depend on XSAVE feature. Another example would be broken
+CPUs and them missing microcode patches. Due to that, the kernel decides not to
+enable a feature.
+
+.. [#f1] 5-level paging uses linear address of 57 bits.
diff --git a/Documentation/x86/earlyprintk.rst b/Documentation/x86/earlyprintk.rst
new file mode 100644
index 000000000..51ef11e8f
--- /dev/null
+++ b/Documentation/x86/earlyprintk.rst
@@ -0,0 +1,151 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+============
+Early Printk
+============
+
+Mini-HOWTO for using the earlyprintk=dbgp boot option with a
+USB2 Debug port key and a debug cable, on x86 systems.
+
+You need two computers, the 'USB debug key' special gadget and
+two USB cables, connected like this::
+
+ [host/target] <-------> [USB debug key] <-------> [client/console]
+
+Hardware requirements
+=====================
+
+ a) Host/target system needs to have USB debug port capability.
+
+ You can check this capability by looking at a 'Debug port' bit in
+ the lspci -vvv output::
+
+ # lspci -vvv
+ ...
+ 00:1d.7 USB Controller: Intel Corporation 82801H (ICH8 Family) USB2 EHCI Controller #1 (rev 03) (prog-if 20 [EHCI])
+ Subsystem: Lenovo ThinkPad T61
+ Control: I/O- Mem+ BusMaster+ SpecCycle- MemWINV- VGASnoop- ParErr- Stepping- SERR+ FastB2B- DisINTx-
+ Status: Cap+ 66MHz- UDF- FastB2B+ ParErr- DEVSEL=medium >TAbort- <TAbort- <MAbort- >SERR- <PERR- INTx-
+ Latency: 0
+ Interrupt: pin D routed to IRQ 19
+ Region 0: Memory at fe227000 (32-bit, non-prefetchable) [size=1K]
+ Capabilities: [50] Power Management version 2
+ Flags: PMEClk- DSI- D1- D2- AuxCurrent=375mA PME(D0+,D1-,D2-,D3hot+,D3cold+)
+ Status: D0 PME-Enable- DSel=0 DScale=0 PME+
+ Capabilities: [58] Debug port: BAR=1 offset=00a0
+ ^^^^^^^^^^^ <==================== [ HERE ]
+ Kernel driver in use: ehci_hcd
+ Kernel modules: ehci-hcd
+ ...
+
+ .. note::
+ If your system does not list a debug port capability then you probably
+ won't be able to use the USB debug key.
+
+ b) You also need a NetChip USB debug cable/key:
+
+ http://www.plxtech.com/products/NET2000/NET20DC/default.asp
+
+ This is a small blue plastic connector with two USB connections;
+ it draws power from its USB connections.
+
+ c) You need a second client/console system with a high speed USB 2.0 port.
+
+ d) The NetChip device must be plugged directly into the physical
+ debug port on the "host/target" system. You cannot use a USB hub in
+ between the physical debug port and the "host/target" system.
+
+ The EHCI debug controller is bound to a specific physical USB
+ port and the NetChip device will only work as an early printk
+ device in this port. The EHCI host controllers are electrically
+ wired such that the EHCI debug controller is hooked up to the
+ first physical port and there is no way to change this via software.
+ You can find the physical port through experimentation by trying
+ each physical port on the system and rebooting. Or you can try
+ and use lsusb or look at the kernel info messages emitted by the
+ usb stack when you plug a usb device into various ports on the
+ "host/target" system.
+
+ Some hardware vendors do not expose the usb debug port with a
+ physical connector and if you find such a device send a complaint
+ to the hardware vendor, because there is no reason not to wire
+ this port into one of the physically accessible ports.
+
+ e) It is also important to note, that many versions of the NetChip
+ device require the "client/console" system to be plugged into the
+ right hand side of the device (with the product logo facing up and
+ readable left to right). The reason being is that the 5 volt
+ power supply is taken from only one side of the device and it
+ must be the side that does not get rebooted.
+
+Software requirements
+=====================
+
+ a) On the host/target system:
+
+ You need to enable the following kernel config option::
+
+ CONFIG_EARLY_PRINTK_DBGP=y
+
+ And you need to add the boot command line: "earlyprintk=dbgp".
+
+ .. note::
+ If you are using Grub, append it to the 'kernel' line in
+ /etc/grub.conf. If you are using Grub2 on a BIOS firmware system,
+ append it to the 'linux' line in /boot/grub2/grub.cfg. If you are
+ using Grub2 on an EFI firmware system, append it to the 'linux'
+ or 'linuxefi' line in /boot/grub2/grub.cfg or
+ /boot/efi/EFI/<distro>/grub.cfg.
+
+ On systems with more than one EHCI debug controller you must
+ specify the correct EHCI debug controller number. The ordering
+ comes from the PCI bus enumeration of the EHCI controllers. The
+ default with no number argument is "0" or the first EHCI debug
+ controller. To use the second EHCI debug controller, you would
+ use the command line: "earlyprintk=dbgp1"
+
+ .. note::
+ normally earlyprintk console gets turned off once the
+ regular console is alive - use "earlyprintk=dbgp,keep" to keep
+ this channel open beyond early bootup. This can be useful for
+ debugging crashes under Xorg, etc.
+
+ b) On the client/console system:
+
+ You should enable the following kernel config option::
+
+ CONFIG_USB_SERIAL_DEBUG=y
+
+ On the next bootup with the modified kernel you should
+ get a /dev/ttyUSBx device(s).
+
+ Now this channel of kernel messages is ready to be used: start
+ your favorite terminal emulator (minicom, etc.) and set
+ it up to use /dev/ttyUSB0 - or use a raw 'cat /dev/ttyUSBx' to
+ see the raw output.
+
+ c) On Nvidia Southbridge based systems: the kernel will try to probe
+ and find out which port has a debug device connected.
+
+Testing
+=======
+
+You can test the output by using earlyprintk=dbgp,keep and provoking
+kernel messages on the host/target system. You can provoke a harmless
+kernel message by for example doing::
+
+ echo h > /proc/sysrq-trigger
+
+On the host/target system you should see this help line in "dmesg" output::
+
+ SysRq : HELP : loglevel(0-9) reBoot Crashdump terminate-all-tasks(E) memory-full-oom-kill(F) kill-all-tasks(I) saK show-backtrace-all-active-cpus(L) show-memory-usage(M) nice-all-RT-tasks(N) powerOff show-registers(P) show-all-timers(Q) unRaw Sync show-task-states(T) Unmount show-blocked-tasks(W) dump-ftrace-buffer(Z)
+
+On the client/console system do::
+
+ cat /dev/ttyUSB0
+
+And you should see the help line above displayed shortly after you've
+provoked it on the host system.
+
+If it does not work then please ask about it on the linux-kernel@vger.kernel.org
+mailing list or contact the x86 maintainers.
diff --git a/Documentation/x86/elf_auxvec.rst b/Documentation/x86/elf_auxvec.rst
new file mode 100644
index 000000000..18e474471
--- /dev/null
+++ b/Documentation/x86/elf_auxvec.rst
@@ -0,0 +1,53 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==================================
+x86-specific ELF Auxiliary Vectors
+==================================
+
+This document describes the semantics of the x86 auxiliary vectors.
+
+Introduction
+============
+
+ELF Auxiliary vectors enable the kernel to efficiently provide
+configuration-specific parameters to userspace. In this example, a program
+allocates an alternate stack based on the kernel-provided size::
+
+ #include <sys/auxv.h>
+ #include <elf.h>
+ #include <signal.h>
+ #include <stdlib.h>
+ #include <assert.h>
+ #include <err.h>
+
+ #ifndef AT_MINSIGSTKSZ
+ #define AT_MINSIGSTKSZ 51
+ #endif
+
+ ....
+ stack_t ss;
+
+ ss.ss_sp = malloc(ss.ss_size);
+ assert(ss.ss_sp);
+
+ ss.ss_size = getauxval(AT_MINSIGSTKSZ) + SIGSTKSZ;
+ ss.ss_flags = 0;
+
+ if (sigaltstack(&ss, NULL))
+ err(1, "sigaltstack");
+
+
+The exposed auxiliary vectors
+=============================
+
+AT_SYSINFO is used for locating the vsyscall entry point. It is not
+exported on 64-bit mode.
+
+AT_SYSINFO_EHDR is the start address of the page containing the vDSO.
+
+AT_MINSIGSTKSZ denotes the minimum stack size required by the kernel to
+deliver a signal to user-space. AT_MINSIGSTKSZ comprehends the space
+consumed by the kernel to accommodate the user context for the current
+hardware configuration. It does not comprehend subsequent user-space stack
+consumption, which must be added by the user. (e.g. Above, user-space adds
+SIGSTKSZ to AT_MINSIGSTKSZ.)
diff --git a/Documentation/x86/entry_64.rst b/Documentation/x86/entry_64.rst
new file mode 100644
index 000000000..0afdce3c0
--- /dev/null
+++ b/Documentation/x86/entry_64.rst
@@ -0,0 +1,110 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==============
+Kernel Entries
+==============
+
+This file documents some of the kernel entries in
+arch/x86/entry/entry_64.S. A lot of this explanation is adapted from
+an email from Ingo Molnar:
+
+https://lore.kernel.org/r/20110529191055.GC9835%40elte.hu
+
+The x86 architecture has quite a few different ways to jump into
+kernel code. Most of these entry points are registered in
+arch/x86/kernel/traps.c and implemented in arch/x86/entry/entry_64.S
+for 64-bit, arch/x86/entry/entry_32.S for 32-bit and finally
+arch/x86/entry/entry_64_compat.S which implements the 32-bit compatibility
+syscall entry points and thus provides for 32-bit processes the
+ability to execute syscalls when running on 64-bit kernels.
+
+The IDT vector assignments are listed in arch/x86/include/asm/irq_vectors.h.
+
+Some of these entries are:
+
+ - system_call: syscall instruction from 64-bit code.
+
+ - entry_INT80_compat: int 0x80 from 32-bit or 64-bit code; compat syscall
+ either way.
+
+ - entry_INT80_compat, ia32_sysenter: syscall and sysenter from 32-bit
+ code
+
+ - interrupt: An array of entries. Every IDT vector that doesn't
+ explicitly point somewhere else gets set to the corresponding
+ value in interrupts. These point to a whole array of
+ magically-generated functions that make their way to common_interrupt()
+ with the interrupt number as a parameter.
+
+ - APIC interrupts: Various special-purpose interrupts for things
+ like TLB shootdown.
+
+ - Architecturally-defined exceptions like divide_error.
+
+There are a few complexities here. The different x86-64 entries
+have different calling conventions. The syscall and sysenter
+instructions have their own peculiar calling conventions. Some of
+the IDT entries push an error code onto the stack; others don't.
+IDT entries using the IST alternative stack mechanism need their own
+magic to get the stack frames right. (You can find some
+documentation in the AMD APM, Volume 2, Chapter 8 and the Intel SDM,
+Volume 3, Chapter 6.)
+
+Dealing with the swapgs instruction is especially tricky. Swapgs
+toggles whether gs is the kernel gs or the user gs. The swapgs
+instruction is rather fragile: it must nest perfectly and only in
+single depth, it should only be used if entering from user mode to
+kernel mode and then when returning to user-space, and precisely
+so. If we mess that up even slightly, we crash.
+
+So when we have a secondary entry, already in kernel mode, we *must
+not* use SWAPGS blindly - nor must we forget doing a SWAPGS when it's
+not switched/swapped yet.
+
+Now, there's a secondary complication: there's a cheap way to test
+which mode the CPU is in and an expensive way.
+
+The cheap way is to pick this info off the entry frame on the kernel
+stack, from the CS of the ptregs area of the kernel stack::
+
+ xorl %ebx,%ebx
+ testl $3,CS+8(%rsp)
+ je error_kernelspace
+ SWAPGS
+
+The expensive (paranoid) way is to read back the MSR_GS_BASE value
+(which is what SWAPGS modifies)::
+
+ movl $1,%ebx
+ movl $MSR_GS_BASE,%ecx
+ rdmsr
+ testl %edx,%edx
+ js 1f /* negative -> in kernel */
+ SWAPGS
+ xorl %ebx,%ebx
+ 1: ret
+
+If we are at an interrupt or user-trap/gate-alike boundary then we can
+use the faster check: the stack will be a reliable indicator of
+whether SWAPGS was already done: if we see that we are a secondary
+entry interrupting kernel mode execution, then we know that the GS
+base has already been switched. If it says that we interrupted
+user-space execution then we must do the SWAPGS.
+
+But if we are in an NMI/MCE/DEBUG/whatever super-atomic entry context,
+which might have triggered right after a normal entry wrote CS to the
+stack but before we executed SWAPGS, then the only safe way to check
+for GS is the slower method: the RDMSR.
+
+Therefore, super-atomic entries (except NMI, which is handled separately)
+must use idtentry with paranoid=1 to handle gsbase correctly. This
+triggers three main behavior changes:
+
+ - Interrupt entry will use the slower gsbase check.
+ - Interrupt entry from user mode will switch off the IST stack.
+ - Interrupt exit to kernel mode will not attempt to reschedule.
+
+We try to only use IST entries and the paranoid entry code for vectors
+that absolutely need the more expensive check for the GS base - and we
+generate all 'normal' entry points with the regular (faster) paranoid=0
+variant.
diff --git a/Documentation/x86/exception-tables.rst b/Documentation/x86/exception-tables.rst
new file mode 100644
index 000000000..efde1fef4
--- /dev/null
+++ b/Documentation/x86/exception-tables.rst
@@ -0,0 +1,357 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===============================
+Kernel level exception handling
+===============================
+
+Commentary by Joerg Pommnitz <joerg@raleigh.ibm.com>
+
+When a process runs in kernel mode, it often has to access user
+mode memory whose address has been passed by an untrusted program.
+To protect itself the kernel has to verify this address.
+
+In older versions of Linux this was done with the
+int verify_area(int type, const void * addr, unsigned long size)
+function (which has since been replaced by access_ok()).
+
+This function verified that the memory area starting at address
+'addr' and of size 'size' was accessible for the operation specified
+in type (read or write). To do this, verify_read had to look up the
+virtual memory area (vma) that contained the address addr. In the
+normal case (correctly working program), this test was successful.
+It only failed for a few buggy programs. In some kernel profiling
+tests, this normally unneeded verification used up a considerable
+amount of time.
+
+To overcome this situation, Linus decided to let the virtual memory
+hardware present in every Linux-capable CPU handle this test.
+
+How does this work?
+
+Whenever the kernel tries to access an address that is currently not
+accessible, the CPU generates a page fault exception and calls the
+page fault handler::
+
+ void exc_page_fault(struct pt_regs *regs, unsigned long error_code)
+
+in arch/x86/mm/fault.c. The parameters on the stack are set up by
+the low level assembly glue in arch/x86/entry/entry_32.S. The parameter
+regs is a pointer to the saved registers on the stack, error_code
+contains a reason code for the exception.
+
+exc_page_fault() first obtains the inaccessible address from the CPU
+control register CR2. If the address is within the virtual address
+space of the process, the fault probably occurred, because the page
+was not swapped in, write protected or something similar. However,
+we are interested in the other case: the address is not valid, there
+is no vma that contains this address. In this case, the kernel jumps
+to the bad_area label.
+
+There it uses the address of the instruction that caused the exception
+(i.e. regs->eip) to find an address where the execution can continue
+(fixup). If this search is successful, the fault handler modifies the
+return address (again regs->eip) and returns. The execution will
+continue at the address in fixup.
+
+Where does fixup point to?
+
+Since we jump to the contents of fixup, fixup obviously points
+to executable code. This code is hidden inside the user access macros.
+I have picked the get_user() macro defined in arch/x86/include/asm/uaccess.h
+as an example. The definition is somewhat hard to follow, so let's peek at
+the code generated by the preprocessor and the compiler. I selected
+the get_user() call in drivers/char/sysrq.c for a detailed examination.
+
+The original code in sysrq.c line 587::
+
+ get_user(c, buf);
+
+The preprocessor output (edited to become somewhat readable)::
+
+ (
+ {
+ long __gu_err = - 14 , __gu_val = 0;
+ const __typeof__(*( ( buf ) )) *__gu_addr = ((buf));
+ if (((((0 + current_set[0])->tss.segment) == 0x18 ) ||
+ (((sizeof(*(buf))) <= 0xC0000000UL) &&
+ ((unsigned long)(__gu_addr ) <= 0xC0000000UL - (sizeof(*(buf)))))))
+ do {
+ __gu_err = 0;
+ switch ((sizeof(*(buf)))) {
+ case 1:
+ __asm__ __volatile__(
+ "1: mov" "b" " %2,%" "b" "1\n"
+ "2:\n"
+ ".section .fixup,\"ax\"\n"
+ "3: movl %3,%0\n"
+ " xor" "b" " %" "b" "1,%" "b" "1\n"
+ " jmp 2b\n"
+ ".section __ex_table,\"a\"\n"
+ " .align 4\n"
+ " .long 1b,3b\n"
+ ".text" : "=r"(__gu_err), "=q" (__gu_val): "m"((*(struct __large_struct *)
+ ( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err )) ;
+ break;
+ case 2:
+ __asm__ __volatile__(
+ "1: mov" "w" " %2,%" "w" "1\n"
+ "2:\n"
+ ".section .fixup,\"ax\"\n"
+ "3: movl %3,%0\n"
+ " xor" "w" " %" "w" "1,%" "w" "1\n"
+ " jmp 2b\n"
+ ".section __ex_table,\"a\"\n"
+ " .align 4\n"
+ " .long 1b,3b\n"
+ ".text" : "=r"(__gu_err), "=r" (__gu_val) : "m"((*(struct __large_struct *)
+ ( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err ));
+ break;
+ case 4:
+ __asm__ __volatile__(
+ "1: mov" "l" " %2,%" "" "1\n"
+ "2:\n"
+ ".section .fixup,\"ax\"\n"
+ "3: movl %3,%0\n"
+ " xor" "l" " %" "" "1,%" "" "1\n"
+ " jmp 2b\n"
+ ".section __ex_table,\"a\"\n"
+ " .align 4\n" " .long 1b,3b\n"
+ ".text" : "=r"(__gu_err), "=r" (__gu_val) : "m"((*(struct __large_struct *)
+ ( __gu_addr )) ), "i"(- 14 ), "0"(__gu_err));
+ break;
+ default:
+ (__gu_val) = __get_user_bad();
+ }
+ } while (0) ;
+ ((c)) = (__typeof__(*((buf))))__gu_val;
+ __gu_err;
+ }
+ );
+
+WOW! Black GCC/assembly magic. This is impossible to follow, so let's
+see what code gcc generates::
+
+ > xorl %edx,%edx
+ > movl current_set,%eax
+ > cmpl $24,788(%eax)
+ > je .L1424
+ > cmpl $-1073741825,64(%esp)
+ > ja .L1423
+ > .L1424:
+ > movl %edx,%eax
+ > movl 64(%esp),%ebx
+ > #APP
+ > 1: movb (%ebx),%dl /* this is the actual user access */
+ > 2:
+ > .section .fixup,"ax"
+ > 3: movl $-14,%eax
+ > xorb %dl,%dl
+ > jmp 2b
+ > .section __ex_table,"a"
+ > .align 4
+ > .long 1b,3b
+ > .text
+ > #NO_APP
+ > .L1423:
+ > movzbl %dl,%esi
+
+The optimizer does a good job and gives us something we can actually
+understand. Can we? The actual user access is quite obvious. Thanks
+to the unified address space we can just access the address in user
+memory. But what does the .section stuff do?????
+
+To understand this we have to look at the final kernel::
+
+ > objdump --section-headers vmlinux
+ >
+ > vmlinux: file format elf32-i386
+ >
+ > Sections:
+ > Idx Name Size VMA LMA File off Algn
+ > 0 .text 00098f40 c0100000 c0100000 00001000 2**4
+ > CONTENTS, ALLOC, LOAD, READONLY, CODE
+ > 1 .fixup 000016bc c0198f40 c0198f40 00099f40 2**0
+ > CONTENTS, ALLOC, LOAD, READONLY, CODE
+ > 2 .rodata 0000f127 c019a5fc c019a5fc 0009b5fc 2**2
+ > CONTENTS, ALLOC, LOAD, READONLY, DATA
+ > 3 __ex_table 000015c0 c01a9724 c01a9724 000aa724 2**2
+ > CONTENTS, ALLOC, LOAD, READONLY, DATA
+ > 4 .data 0000ea58 c01abcf0 c01abcf0 000abcf0 2**4
+ > CONTENTS, ALLOC, LOAD, DATA
+ > 5 .bss 00018e21 c01ba748 c01ba748 000ba748 2**2
+ > ALLOC
+ > 6 .comment 00000ec4 00000000 00000000 000ba748 2**0
+ > CONTENTS, READONLY
+ > 7 .note 00001068 00000ec4 00000ec4 000bb60c 2**0
+ > CONTENTS, READONLY
+
+There are obviously 2 non standard ELF sections in the generated object
+file. But first we want to find out what happened to our code in the
+final kernel executable::
+
+ > objdump --disassemble --section=.text vmlinux
+ >
+ > c017e785 <do_con_write+c1> xorl %edx,%edx
+ > c017e787 <do_con_write+c3> movl 0xc01c7bec,%eax
+ > c017e78c <do_con_write+c8> cmpl $0x18,0x314(%eax)
+ > c017e793 <do_con_write+cf> je c017e79f <do_con_write+db>
+ > c017e795 <do_con_write+d1> cmpl $0xbfffffff,0x40(%esp,1)
+ > c017e79d <do_con_write+d9> ja c017e7a7 <do_con_write+e3>
+ > c017e79f <do_con_write+db> movl %edx,%eax
+ > c017e7a1 <do_con_write+dd> movl 0x40(%esp,1),%ebx
+ > c017e7a5 <do_con_write+e1> movb (%ebx),%dl
+ > c017e7a7 <do_con_write+e3> movzbl %dl,%esi
+
+The whole user memory access is reduced to 10 x86 machine instructions.
+The instructions bracketed in the .section directives are no longer
+in the normal execution path. They are located in a different section
+of the executable file::
+
+ > objdump --disassemble --section=.fixup vmlinux
+ >
+ > c0199ff5 <.fixup+10b5> movl $0xfffffff2,%eax
+ > c0199ffa <.fixup+10ba> xorb %dl,%dl
+ > c0199ffc <.fixup+10bc> jmp c017e7a7 <do_con_write+e3>
+
+And finally::
+
+ > objdump --full-contents --section=__ex_table vmlinux
+ >
+ > c01aa7c4 93c017c0 e09f19c0 97c017c0 99c017c0 ................
+ > c01aa7d4 f6c217c0 e99f19c0 a5e717c0 f59f19c0 ................
+ > c01aa7e4 080a18c0 01a019c0 0a0a18c0 04a019c0 ................
+
+or in human readable byte order::
+
+ > c01aa7c4 c017c093 c0199fe0 c017c097 c017c099 ................
+ > c01aa7d4 c017c2f6 c0199fe9 c017e7a5 c0199ff5 ................
+ ^^^^^^^^^^^^^^^^^
+ this is the interesting part!
+ > c01aa7e4 c0180a08 c019a001 c0180a0a c019a004 ................
+
+What happened? The assembly directives::
+
+ .section .fixup,"ax"
+ .section __ex_table,"a"
+
+told the assembler to move the following code to the specified
+sections in the ELF object file. So the instructions::
+
+ 3: movl $-14,%eax
+ xorb %dl,%dl
+ jmp 2b
+
+ended up in the .fixup section of the object file and the addresses::
+
+ .long 1b,3b
+
+ended up in the __ex_table section of the object file. 1b and 3b
+are local labels. The local label 1b (1b stands for next label 1
+backward) is the address of the instruction that might fault, i.e.
+in our case the address of the label 1 is c017e7a5:
+the original assembly code: > 1: movb (%ebx),%dl
+and linked in vmlinux : > c017e7a5 <do_con_write+e1> movb (%ebx),%dl
+
+The local label 3 (backwards again) is the address of the code to handle
+the fault, in our case the actual value is c0199ff5:
+the original assembly code: > 3: movl $-14,%eax
+and linked in vmlinux : > c0199ff5 <.fixup+10b5> movl $0xfffffff2,%eax
+
+If the fixup was able to handle the exception, control flow may be returned
+to the instruction after the one that triggered the fault, ie. local label 2b.
+
+The assembly code::
+
+ > .section __ex_table,"a"
+ > .align 4
+ > .long 1b,3b
+
+becomes the value pair::
+
+ > c01aa7d4 c017c2f6 c0199fe9 c017e7a5 c0199ff5 ................
+ ^this is ^this is
+ 1b 3b
+
+c017e7a5,c0199ff5 in the exception table of the kernel.
+
+So, what actually happens if a fault from kernel mode with no suitable
+vma occurs?
+
+#. access to invalid address::
+
+ > c017e7a5 <do_con_write+e1> movb (%ebx),%dl
+#. MMU generates exception
+#. CPU calls exc_page_fault()
+#. exc_page_fault() calls do_user_addr_fault()
+#. do_user_addr_fault() calls kernelmode_fixup_or_oops()
+#. kernelmode_fixup_or_oops() calls fixup_exception() (regs->eip == c017e7a5);
+#. fixup_exception() calls search_exception_tables()
+#. search_exception_tables() looks up the address c017e7a5 in the
+ exception table (i.e. the contents of the ELF section __ex_table)
+ and returns the address of the associated fault handle code c0199ff5.
+#. fixup_exception() modifies its own return address to point to the fault
+ handle code and returns.
+#. execution continues in the fault handling code.
+#. a) EAX becomes -EFAULT (== -14)
+ b) DL becomes zero (the value we "read" from user space)
+ c) execution continues at local label 2 (address of the
+ instruction immediately after the faulting user access).
+
+The steps 8a to 8c in a certain way emulate the faulting instruction.
+
+That's it, mostly. If you look at our example, you might ask why
+we set EAX to -EFAULT in the exception handler code. Well, the
+get_user() macro actually returns a value: 0, if the user access was
+successful, -EFAULT on failure. Our original code did not test this
+return value, however the inline assembly code in get_user() tries to
+return -EFAULT. GCC selected EAX to return this value.
+
+NOTE:
+Due to the way that the exception table is built and needs to be ordered,
+only use exceptions for code in the .text section. Any other section
+will cause the exception table to not be sorted correctly, and the
+exceptions will fail.
+
+Things changed when 64-bit support was added to x86 Linux. Rather than
+double the size of the exception table by expanding the two entries
+from 32-bits to 64 bits, a clever trick was used to store addresses
+as relative offsets from the table itself. The assembly code changed
+from::
+
+ .long 1b,3b
+ to:
+ .long (from) - .
+ .long (to) - .
+
+and the C-code that uses these values converts back to absolute addresses
+like this::
+
+ ex_insn_addr(const struct exception_table_entry *x)
+ {
+ return (unsigned long)&x->insn + x->insn;
+ }
+
+In v4.6 the exception table entry was expanded with a new field "handler".
+This is also 32-bits wide and contains a third relative function
+pointer which points to one of:
+
+1) ``int ex_handler_default(const struct exception_table_entry *fixup)``
+ This is legacy case that just jumps to the fixup code
+
+2) ``int ex_handler_fault(const struct exception_table_entry *fixup)``
+ This case provides the fault number of the trap that occurred at
+ entry->insn. It is used to distinguish page faults from machine
+ check.
+
+More functions can easily be added.
+
+CONFIG_BUILDTIME_TABLE_SORT allows the __ex_table section to be sorted post
+link of the kernel image, via a host utility scripts/sorttable. It will set the
+symbol main_extable_sort_needed to 0, avoiding sorting the __ex_table section
+at boot time. With the exception table sorted, at runtime when an exception
+occurs we can quickly lookup the __ex_table entry via binary search.
+
+This is not just a boot time optimization, some architectures require this
+table to be sorted in order to handle exceptions relatively early in the boot
+process. For example, i386 makes use of this form of exception handling before
+paging support is even enabled!
diff --git a/Documentation/x86/features.rst b/Documentation/x86/features.rst
new file mode 100644
index 000000000..b663f1505
--- /dev/null
+++ b/Documentation/x86/features.rst
@@ -0,0 +1,3 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+.. kernel-feat:: $srctree/Documentation/features x86
diff --git a/Documentation/x86/i386/IO-APIC.rst b/Documentation/x86/i386/IO-APIC.rst
new file mode 100644
index 000000000..ce4d8df15
--- /dev/null
+++ b/Documentation/x86/i386/IO-APIC.rst
@@ -0,0 +1,123 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=======
+IO-APIC
+=======
+
+:Author: Ingo Molnar <mingo@kernel.org>
+
+Most (all) Intel-MP compliant SMP boards have the so-called 'IO-APIC',
+which is an enhanced interrupt controller. It enables us to route
+hardware interrupts to multiple CPUs, or to CPU groups. Without an
+IO-APIC, interrupts from hardware will be delivered only to the
+CPU which boots the operating system (usually CPU#0).
+
+Linux supports all variants of compliant SMP boards, including ones with
+multiple IO-APICs. Multiple IO-APICs are used in high-end servers to
+distribute IRQ load further.
+
+There are (a few) known breakages in certain older boards, such bugs are
+usually worked around by the kernel. If your MP-compliant SMP board does
+not boot Linux, then consult the linux-smp mailing list archives first.
+
+If your box boots fine with enabled IO-APIC IRQs, then your
+/proc/interrupts will look like this one::
+
+ hell:~> cat /proc/interrupts
+ CPU0
+ 0: 1360293 IO-APIC-edge timer
+ 1: 4 IO-APIC-edge keyboard
+ 2: 0 XT-PIC cascade
+ 13: 1 XT-PIC fpu
+ 14: 1448 IO-APIC-edge ide0
+ 16: 28232 IO-APIC-level Intel EtherExpress Pro 10/100 Ethernet
+ 17: 51304 IO-APIC-level eth0
+ NMI: 0
+ ERR: 0
+ hell:~>
+
+Some interrupts are still listed as 'XT PIC', but this is not a problem;
+none of those IRQ sources is performance-critical.
+
+
+In the unlikely case that your board does not create a working mp-table,
+you can use the pirq= boot parameter to 'hand-construct' IRQ entries. This
+is non-trivial though and cannot be automated. One sample /etc/lilo.conf
+entry::
+
+ append="pirq=15,11,10"
+
+The actual numbers depend on your system, on your PCI cards and on their
+PCI slot position. Usually PCI slots are 'daisy chained' before they are
+connected to the PCI chipset IRQ routing facility (the incoming PIRQ1-4
+lines)::
+
+ ,-. ,-. ,-. ,-. ,-.
+ PIRQ4 ----| |-. ,-| |-. ,-| |-. ,-| |--------| |
+ |S| \ / |S| \ / |S| \ / |S| |S|
+ PIRQ3 ----|l|-. `/---|l|-. `/---|l|-. `/---|l|--------|l|
+ |o| \/ |o| \/ |o| \/ |o| |o|
+ PIRQ2 ----|t|-./`----|t|-./`----|t|-./`----|t|--------|t|
+ |1| /\ |2| /\ |3| /\ |4| |5|
+ PIRQ1 ----| |- `----| |- `----| |- `----| |--------| |
+ `-' `-' `-' `-' `-'
+
+Every PCI card emits a PCI IRQ, which can be INTA, INTB, INTC or INTD::
+
+ ,-.
+ INTD--| |
+ |S|
+ INTC--|l|
+ |o|
+ INTB--|t|
+ |x|
+ INTA--| |
+ `-'
+
+These INTA-D PCI IRQs are always 'local to the card', their real meaning
+depends on which slot they are in. If you look at the daisy chaining diagram,
+a card in slot4, issuing INTA IRQ, it will end up as a signal on PIRQ4 of
+the PCI chipset. Most cards issue INTA, this creates optimal distribution
+between the PIRQ lines. (distributing IRQ sources properly is not a
+necessity, PCI IRQs can be shared at will, but it's a good for performance
+to have non shared interrupts). Slot5 should be used for videocards, they
+do not use interrupts normally, thus they are not daisy chained either.
+
+so if you have your SCSI card (IRQ11) in Slot1, Tulip card (IRQ9) in
+Slot2, then you'll have to specify this pirq= line::
+
+ append="pirq=11,9"
+
+the following script tries to figure out such a default pirq= line from
+your PCI configuration::
+
+ echo -n pirq=; echo `scanpci | grep T_L | cut -c56-` | sed 's/ /,/g'
+
+note that this script won't work if you have skipped a few slots or if your
+board does not do default daisy-chaining. (or the IO-APIC has the PIRQ pins
+connected in some strange way). E.g. if in the above case you have your SCSI
+card (IRQ11) in Slot3, and have Slot1 empty::
+
+ append="pirq=0,9,11"
+
+[value '0' is a generic 'placeholder', reserved for empty (or non-IRQ emitting)
+slots.]
+
+Generally, it's always possible to find out the correct pirq= settings, just
+permute all IRQ numbers properly ... it will take some time though. An
+'incorrect' pirq line will cause the booting process to hang, or a device
+won't function properly (e.g. if it's inserted as a module).
+
+If you have 2 PCI buses, then you can use up to 8 pirq values, although such
+boards tend to have a good configuration.
+
+Be prepared that it might happen that you need some strange pirq line::
+
+ append="pirq=0,0,0,0,0,0,9,11"
+
+Use smart trial-and-error techniques to find out the correct pirq line ...
+
+Good luck and mail to linux-smp@vger.kernel.org or
+linux-kernel@vger.kernel.org if you have any problems that are not covered
+by this document.
+
diff --git a/Documentation/x86/i386/index.rst b/Documentation/x86/i386/index.rst
new file mode 100644
index 000000000..8747cf5bb
--- /dev/null
+++ b/Documentation/x86/i386/index.rst
@@ -0,0 +1,10 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+============
+i386 Support
+============
+
+.. toctree::
+ :maxdepth: 2
+
+ IO-APIC
diff --git a/Documentation/x86/ifs.rst b/Documentation/x86/ifs.rst
new file mode 100644
index 000000000..97abb696a
--- /dev/null
+++ b/Documentation/x86/ifs.rst
@@ -0,0 +1,2 @@
+.. SPDX-License-Identifier: GPL-2.0
+.. kernel-doc:: drivers/platform/x86/intel/ifs/ifs.h
diff --git a/Documentation/x86/index.rst b/Documentation/x86/index.rst
new file mode 100644
index 000000000..c73d133fd
--- /dev/null
+++ b/Documentation/x86/index.rst
@@ -0,0 +1,44 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==========================
+x86-specific Documentation
+==========================
+
+.. toctree::
+ :maxdepth: 2
+ :numbered:
+
+ boot
+ booting-dt
+ cpuinfo
+ topology
+ exception-tables
+ kernel-stacks
+ entry_64
+ earlyprintk
+ orc-unwinder
+ zero-page
+ tlb
+ mtrr
+ pat
+ intel-hfi
+ iommu
+ intel_txt
+ amd-memory-encryption
+ amd_hsmp
+ tdx
+ pti
+ mds
+ microcode
+ resctrl
+ tsx_async_abort
+ buslock
+ usb-legacy-support
+ i386/index
+ x86_64/index
+ ifs
+ sva
+ sgx
+ features
+ elf_auxvec
+ xstate
diff --git a/Documentation/x86/intel-hfi.rst b/Documentation/x86/intel-hfi.rst
new file mode 100644
index 000000000..49dea58ea
--- /dev/null
+++ b/Documentation/x86/intel-hfi.rst
@@ -0,0 +1,72 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+============================================================
+Hardware-Feedback Interface for scheduling on Intel Hardware
+============================================================
+
+Overview
+--------
+
+Intel has described the Hardware Feedback Interface (HFI) in the Intel 64 and
+IA-32 Architectures Software Developer's Manual (Intel SDM) Volume 3 Section
+14.6 [1]_.
+
+The HFI gives the operating system a performance and energy efficiency
+capability data for each CPU in the system. Linux can use the information from
+the HFI to influence task placement decisions.
+
+The Hardware Feedback Interface
+-------------------------------
+
+The Hardware Feedback Interface provides to the operating system information
+about the performance and energy efficiency of each CPU in the system. Each
+capability is given as a unit-less quantity in the range [0-255]. Higher values
+indicate higher capability. Energy efficiency and performance are reported in
+separate capabilities. Even though on some systems these two metrics may be
+related, they are specified as independent capabilities in the Intel SDM.
+
+These capabilities may change at runtime as a result of changes in the
+operating conditions of the system or the action of external factors. The rate
+at which these capabilities are updated is specific to each processor model. On
+some models, capabilities are set at boot time and never change. On others,
+capabilities may change every tens of milliseconds. For instance, a remote
+mechanism may be used to lower Thermal Design Power. Such change can be
+reflected in the HFI. Likewise, if the system needs to be throttled due to
+excessive heat, the HFI may reflect reduced performance on specific CPUs.
+
+The kernel or a userspace policy daemon can use these capabilities to modify
+task placement decisions. For instance, if either the performance or energy
+capabilities of a given logical processor becomes zero, it is an indication that
+the hardware recommends to the operating system to not schedule any tasks on
+that processor for performance or energy efficiency reasons, respectively.
+
+Implementation details for Linux
+--------------------------------
+
+The infrastructure to handle thermal event interrupts has two parts. In the
+Local Vector Table of a CPU's local APIC, there exists a register for the
+Thermal Monitor Register. This register controls how interrupts are delivered
+to a CPU when the thermal monitor generates and interrupt. Further details
+can be found in the Intel SDM Vol. 3 Section 10.5 [1]_.
+
+The thermal monitor may generate interrupts per CPU or per package. The HFI
+generates package-level interrupts. This monitor is configured and initialized
+via a set of machine-specific registers. Specifically, the HFI interrupt and
+status are controlled via designated bits in the IA32_PACKAGE_THERM_INTERRUPT
+and IA32_PACKAGE_THERM_STATUS registers, respectively. There exists one HFI
+table per package. Further details can be found in the Intel SDM Vol. 3
+Section 14.9 [1]_.
+
+The hardware issues an HFI interrupt after updating the HFI table and is ready
+for the operating system to consume it. CPUs receive such interrupt via the
+thermal entry in the Local APIC's Local Vector Table.
+
+When servicing such interrupt, the HFI driver parses the updated table and
+relays the update to userspace using the thermal notification framework. Given
+that there may be many HFI updates every second, the updates relayed to
+userspace are throttled at a rate of CONFIG_HZ jiffies.
+
+References
+----------
+
+.. [1] https://www.intel.com/sdm
diff --git a/Documentation/x86/intel_txt.rst b/Documentation/x86/intel_txt.rst
new file mode 100644
index 000000000..d83c1a212
--- /dev/null
+++ b/Documentation/x86/intel_txt.rst
@@ -0,0 +1,227 @@
+=====================
+Intel(R) TXT Overview
+=====================
+
+Intel's technology for safer computing, Intel(R) Trusted Execution
+Technology (Intel(R) TXT), defines platform-level enhancements that
+provide the building blocks for creating trusted platforms.
+
+Intel TXT was formerly known by the code name LaGrande Technology (LT).
+
+Intel TXT in Brief:
+
+- Provides dynamic root of trust for measurement (DRTM)
+- Data protection in case of improper shutdown
+- Measurement and verification of launched environment
+
+Intel TXT is part of the vPro(TM) brand and is also available some
+non-vPro systems. It is currently available on desktop systems
+based on the Q35, X38, Q45, and Q43 Express chipsets (e.g. Dell
+Optiplex 755, HP dc7800, etc.) and mobile systems based on the GM45,
+PM45, and GS45 Express chipsets.
+
+For more information, see http://www.intel.com/technology/security/.
+This site also has a link to the Intel TXT MLE Developers Manual,
+which has been updated for the new released platforms.
+
+Intel TXT has been presented at various events over the past few
+years, some of which are:
+
+ - LinuxTAG 2008:
+ http://www.linuxtag.org/2008/en/conf/events/vp-donnerstag.html
+
+ - TRUST2008:
+ http://www.trust-conference.eu/downloads/Keynote-Speakers/
+ 3_David-Grawrock_The-Front-Door-of-Trusted-Computing.pdf
+
+ - IDF, Shanghai:
+ http://www.prcidf.com.cn/index_en.html
+
+ - IDFs 2006, 2007
+ (I'm not sure if/where they are online)
+
+Trusted Boot Project Overview
+=============================
+
+Trusted Boot (tboot) is an open source, pre-kernel/VMM module that
+uses Intel TXT to perform a measured and verified launch of an OS
+kernel/VMM.
+
+It is hosted on SourceForge at http://sourceforge.net/projects/tboot.
+The mercurial source repo is available at http://www.bughost.org/
+repos.hg/tboot.hg.
+
+Tboot currently supports launching Xen (open source VMM/hypervisor
+w/ TXT support since v3.2), and now Linux kernels.
+
+
+Value Proposition for Linux or "Why should you care?"
+=====================================================
+
+While there are many products and technologies that attempt to
+measure or protect the integrity of a running kernel, they all
+assume the kernel is "good" to begin with. The Integrity
+Measurement Architecture (IMA) and Linux Integrity Module interface
+are examples of such solutions.
+
+To get trust in the initial kernel without using Intel TXT, a
+static root of trust must be used. This bases trust in BIOS
+starting at system reset and requires measurement of all code
+executed between system reset through the completion of the kernel
+boot as well as data objects used by that code. In the case of a
+Linux kernel, this means all of BIOS, any option ROMs, the
+bootloader and the boot config. In practice, this is a lot of
+code/data, much of which is subject to change from boot to boot
+(e.g. changing NICs may change option ROMs). Without reference
+hashes, these measurement changes are difficult to assess or
+confirm as benign. This process also does not provide DMA
+protection, memory configuration/alias checks and locks, crash
+protection, or policy support.
+
+By using the hardware-based root of trust that Intel TXT provides,
+many of these issues can be mitigated. Specifically: many
+pre-launch components can be removed from the trust chain, DMA
+protection is provided to all launched components, a large number
+of platform configuration checks are performed and values locked,
+protection is provided for any data in the event of an improper
+shutdown, and there is support for policy-based execution/verification.
+This provides a more stable measurement and a higher assurance of
+system configuration and initial state than would be otherwise
+possible. Since the tboot project is open source, source code for
+almost all parts of the trust chain is available (excepting SMM and
+Intel-provided firmware).
+
+How Does it Work?
+=================
+
+- Tboot is an executable that is launched by the bootloader as
+ the "kernel" (the binary the bootloader executes).
+- It performs all of the work necessary to determine if the
+ platform supports Intel TXT and, if so, executes the GETSEC[SENTER]
+ processor instruction that initiates the dynamic root of trust.
+
+ - If tboot determines that the system does not support Intel TXT
+ or is not configured correctly (e.g. the SINIT AC Module was
+ incorrect), it will directly launch the kernel with no changes
+ to any state.
+ - Tboot will output various information about its progress to the
+ terminal, serial port, and/or an in-memory log; the output
+ locations can be configured with a command line switch.
+
+- The GETSEC[SENTER] instruction will return control to tboot and
+ tboot then verifies certain aspects of the environment (e.g. TPM NV
+ lock, e820 table does not have invalid entries, etc.).
+- It will wake the APs from the special sleep state the GETSEC[SENTER]
+ instruction had put them in and place them into a wait-for-SIPI
+ state.
+
+ - Because the processors will not respond to an INIT or SIPI when
+ in the TXT environment, it is necessary to create a small VT-x
+ guest for the APs. When they run in this guest, they will
+ simply wait for the INIT-SIPI-SIPI sequence, which will cause
+ VMEXITs, and then disable VT and jump to the SIPI vector. This
+ approach seemed like a better choice than having to insert
+ special code into the kernel's MP wakeup sequence.
+
+- Tboot then applies an (optional) user-defined launch policy to
+ verify the kernel and initrd.
+
+ - This policy is rooted in TPM NV and is described in the tboot
+ project. The tboot project also contains code for tools to
+ create and provision the policy.
+ - Policies are completely under user control and if not present
+ then any kernel will be launched.
+ - Policy action is flexible and can include halting on failures
+ or simply logging them and continuing.
+
+- Tboot adjusts the e820 table provided by the bootloader to reserve
+ its own location in memory as well as to reserve certain other
+ TXT-related regions.
+- As part of its launch, tboot DMA protects all of RAM (using the
+ VT-d PMRs). Thus, the kernel must be booted with 'intel_iommu=on'
+ in order to remove this blanket protection and use VT-d's
+ page-level protection.
+- Tboot will populate a shared page with some data about itself and
+ pass this to the Linux kernel as it transfers control.
+
+ - The location of the shared page is passed via the boot_params
+ struct as a physical address.
+
+- The kernel will look for the tboot shared page address and, if it
+ exists, map it.
+- As one of the checks/protections provided by TXT, it makes a copy
+ of the VT-d DMARs in a DMA-protected region of memory and verifies
+ them for correctness. The VT-d code will detect if the kernel was
+ launched with tboot and use this copy instead of the one in the
+ ACPI table.
+- At this point, tboot and TXT are out of the picture until a
+ shutdown (S<n>)
+- In order to put a system into any of the sleep states after a TXT
+ launch, TXT must first be exited. This is to prevent attacks that
+ attempt to crash the system to gain control on reboot and steal
+ data left in memory.
+
+ - The kernel will perform all of its sleep preparation and
+ populate the shared page with the ACPI data needed to put the
+ platform in the desired sleep state.
+ - Then the kernel jumps into tboot via the vector specified in the
+ shared page.
+ - Tboot will clean up the environment and disable TXT, then use the
+ kernel-provided ACPI information to actually place the platform
+ into the desired sleep state.
+ - In the case of S3, tboot will also register itself as the resume
+ vector. This is necessary because it must re-establish the
+ measured environment upon resume. Once the TXT environment
+ has been restored, it will restore the TPM PCRs and then
+ transfer control back to the kernel's S3 resume vector.
+ In order to preserve system integrity across S3, the kernel
+ provides tboot with a set of memory ranges (RAM and RESERVED_KERN
+ in the e820 table, but not any memory that BIOS might alter over
+ the S3 transition) that tboot will calculate a MAC (message
+ authentication code) over and then seal with the TPM. On resume
+ and once the measured environment has been re-established, tboot
+ will re-calculate the MAC and verify it against the sealed value.
+ Tboot's policy determines what happens if the verification fails.
+ Note that the c/s 194 of tboot which has the new MAC code supports
+ this.
+
+That's pretty much it for TXT support.
+
+
+Configuring the System
+======================
+
+This code works with 32bit, 32bit PAE, and 64bit (x86_64) kernels.
+
+In BIOS, the user must enable: TPM, TXT, VT-x, VT-d. Not all BIOSes
+allow these to be individually enabled/disabled and the screens in
+which to find them are BIOS-specific.
+
+grub.conf needs to be modified as follows::
+
+ title Linux 2.6.29-tip w/ tboot
+ root (hd0,0)
+ kernel /tboot.gz logging=serial,vga,memory
+ module /vmlinuz-2.6.29-tip intel_iommu=on ro
+ root=LABEL=/ rhgb console=ttyS0,115200 3
+ module /initrd-2.6.29-tip.img
+ module /Q35_SINIT_17.BIN
+
+The kernel option for enabling Intel TXT support is found under the
+Security top-level menu and is called "Enable Intel(R) Trusted
+Execution Technology (TXT)". It is considered EXPERIMENTAL and
+depends on the generic x86 support (to allow maximum flexibility in
+kernel build options), since the tboot code will detect whether the
+platform actually supports Intel TXT and thus whether any of the
+kernel code is executed.
+
+The Q35_SINIT_17.BIN file is what Intel TXT refers to as an
+Authenticated Code Module. It is specific to the chipset in the
+system and can also be found on the Trusted Boot site. It is an
+(unencrypted) module signed by Intel that is used as part of the
+DRTM process to verify and configure the system. It is signed
+because it operates at a higher privilege level in the system than
+any other macrocode and its correct operation is critical to the
+establishment of the DRTM. The process for determining the correct
+SINIT ACM for a system is documented in the SINIT-guide.txt file
+that is on the tboot SourceForge site under the SINIT ACM downloads.
diff --git a/Documentation/x86/iommu.rst b/Documentation/x86/iommu.rst
new file mode 100644
index 000000000..42c7a6faa
--- /dev/null
+++ b/Documentation/x86/iommu.rst
@@ -0,0 +1,151 @@
+=================
+x86 IOMMU Support
+=================
+
+The architecture specs can be obtained from the below locations.
+
+- Intel: http://www.intel.com/content/dam/www/public/us/en/documents/product-specifications/vt-directed-io-spec.pdf
+- AMD: https://www.amd.com/system/files/TechDocs/48882_IOMMU.pdf
+
+This guide gives a quick cheat sheet for some basic understanding.
+
+Basic stuff
+-----------
+
+ACPI enumerates and lists the different IOMMUs on the platform, and
+device scope relationships between devices and which IOMMU controls
+them.
+
+Some ACPI Keywords:
+
+- DMAR - Intel DMA Remapping table
+- DRHD - Intel DMA Remapping Hardware Unit Definition
+- RMRR - Intel Reserved Memory Region Reporting Structure
+- IVRS - AMD I/O Virtualization Reporting Structure
+- IVDB - AMD I/O Virtualization Definition Block
+- IVHD - AMD I/O Virtualization Hardware Definition
+
+What is Intel RMRR?
+^^^^^^^^^^^^^^^^^^^
+
+There are some devices the BIOS controls, for e.g USB devices to perform
+PS2 emulation. The regions of memory used for these devices are marked
+reserved in the e820 map. When we turn on DMA translation, DMA to those
+regions will fail. Hence BIOS uses RMRR to specify these regions along with
+devices that need to access these regions. OS is expected to setup
+unity mappings for these regions for these devices to access these regions.
+
+What is AMD IVRS?
+^^^^^^^^^^^^^^^^^
+
+The architecture defines an ACPI-compatible data structure called an I/O
+Virtualization Reporting Structure (IVRS) that is used to convey information
+related to I/O virtualization to system software. The IVRS describes the
+configuration and capabilities of the IOMMUs contained in the platform as
+well as information about the devices that each IOMMU virtualizes.
+
+The IVRS provides information about the following:
+
+- IOMMUs present in the platform including their capabilities and proper configuration
+- System I/O topology relevant to each IOMMU
+- Peripheral devices that cannot be otherwise enumerated
+- Memory regions used by SMI/SMM, platform firmware, and platform hardware. These are generally exclusion ranges to be configured by system software.
+
+How is an I/O Virtual Address (IOVA) generated?
+-----------------------------------------------
+
+Well behaved drivers call dma_map_*() calls before sending command to device
+that needs to perform DMA. Once DMA is completed and mapping is no longer
+required, driver performs dma_unmap_*() calls to unmap the region.
+
+Intel Specific Notes
+--------------------
+
+Graphics Problems?
+^^^^^^^^^^^^^^^^^^
+
+If you encounter issues with graphics devices, you can try adding
+option intel_iommu=igfx_off to turn off the integrated graphics engine.
+If this fixes anything, please ensure you file a bug reporting the problem.
+
+Some exceptions to IOVA
+^^^^^^^^^^^^^^^^^^^^^^^
+
+Interrupt ranges are not address translated, (0xfee00000 - 0xfeefffff).
+The same is true for peer to peer transactions. Hence we reserve the
+address from PCI MMIO ranges so they are not allocated for IOVA addresses.
+
+AMD Specific Notes
+------------------
+
+Graphics Problems?
+^^^^^^^^^^^^^^^^^^
+
+If you encounter issues with integrated graphics devices, you can try adding
+option iommu=pt to the kernel command line use a 1:1 mapping for the IOMMU. If
+this fixes anything, please ensure you file a bug reporting the problem.
+
+Fault reporting
+---------------
+When errors are reported, the IOMMU signals via an interrupt. The fault
+reason and device that caused it is printed on the console.
+
+
+Kernel Log Samples
+------------------
+
+Intel Boot Messages
+^^^^^^^^^^^^^^^^^^^
+
+Something like this gets printed indicating presence of DMAR tables
+in ACPI:
+
+::
+
+ ACPI: DMAR (v001 A M I OEMDMAR 0x00000001 MSFT 0x00000097) @ 0x000000007f5b5ef0
+
+When DMAR is being processed and initialized by ACPI, prints DMAR locations
+and any RMRR's processed:
+
+::
+
+ ACPI DMAR:Host address width 36
+ ACPI DMAR:DRHD (flags: 0x00000000)base: 0x00000000fed90000
+ ACPI DMAR:DRHD (flags: 0x00000000)base: 0x00000000fed91000
+ ACPI DMAR:DRHD (flags: 0x00000001)base: 0x00000000fed93000
+ ACPI DMAR:RMRR base: 0x00000000000ed000 end: 0x00000000000effff
+ ACPI DMAR:RMRR base: 0x000000007f600000 end: 0x000000007fffffff
+
+When DMAR is enabled for use, you will notice:
+
+::
+
+ PCI-DMA: Using DMAR IOMMU
+
+Intel Fault reporting
+^^^^^^^^^^^^^^^^^^^^^
+
+::
+
+ DMAR:[DMA Write] Request device [00:02.0] fault addr 6df084000
+ DMAR:[fault reason 05] PTE Write access is not set
+ DMAR:[DMA Write] Request device [00:02.0] fault addr 6df084000
+ DMAR:[fault reason 05] PTE Write access is not set
+
+AMD Boot Messages
+^^^^^^^^^^^^^^^^^
+
+Something like this gets printed indicating presence of the IOMMU:
+
+::
+
+ iommu: Default domain type: Translated
+ iommu: DMA domain TLB invalidation policy: lazy mode
+
+AMD Fault reporting
+^^^^^^^^^^^^^^^^^^^
+
+::
+
+ AMD-Vi: Event logged [IO_PAGE_FAULT domain=0x0007 address=0xffffc02000 flags=0x0000]
+ AMD-Vi: Event logged [IO_PAGE_FAULT device=07:00.0 domain=0x0007 address=0xffffc02000 flags=0x0000]
diff --git a/Documentation/x86/kernel-stacks.rst b/Documentation/x86/kernel-stacks.rst
new file mode 100644
index 000000000..6b0bcf027
--- /dev/null
+++ b/Documentation/x86/kernel-stacks.rst
@@ -0,0 +1,152 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=============
+Kernel Stacks
+=============
+
+Kernel stacks on x86-64 bit
+===========================
+
+Most of the text from Keith Owens, hacked by AK
+
+x86_64 page size (PAGE_SIZE) is 4K.
+
+Like all other architectures, x86_64 has a kernel stack for every
+active thread. These thread stacks are THREAD_SIZE (2*PAGE_SIZE) big.
+These stacks contain useful data as long as a thread is alive or a
+zombie. While the thread is in user space the kernel stack is empty
+except for the thread_info structure at the bottom.
+
+In addition to the per thread stacks, there are specialized stacks
+associated with each CPU. These stacks are only used while the kernel
+is in control on that CPU; when a CPU returns to user space the
+specialized stacks contain no useful data. The main CPU stacks are:
+
+* Interrupt stack. IRQ_STACK_SIZE
+
+ Used for external hardware interrupts. If this is the first external
+ hardware interrupt (i.e. not a nested hardware interrupt) then the
+ kernel switches from the current task to the interrupt stack. Like
+ the split thread and interrupt stacks on i386, this gives more room
+ for kernel interrupt processing without having to increase the size
+ of every per thread stack.
+
+ The interrupt stack is also used when processing a softirq.
+
+Switching to the kernel interrupt stack is done by software based on a
+per CPU interrupt nest counter. This is needed because x86-64 "IST"
+hardware stacks cannot nest without races.
+
+x86_64 also has a feature which is not available on i386, the ability
+to automatically switch to a new stack for designated events such as
+double fault or NMI, which makes it easier to handle these unusual
+events on x86_64. This feature is called the Interrupt Stack Table
+(IST). There can be up to 7 IST entries per CPU. The IST code is an
+index into the Task State Segment (TSS). The IST entries in the TSS
+point to dedicated stacks; each stack can be a different size.
+
+An IST is selected by a non-zero value in the IST field of an
+interrupt-gate descriptor. When an interrupt occurs and the hardware
+loads such a descriptor, the hardware automatically sets the new stack
+pointer based on the IST value, then invokes the interrupt handler. If
+the interrupt came from user mode, then the interrupt handler prologue
+will switch back to the per-thread stack. If software wants to allow
+nested IST interrupts then the handler must adjust the IST values on
+entry to and exit from the interrupt handler. (This is occasionally
+done, e.g. for debug exceptions.)
+
+Events with different IST codes (i.e. with different stacks) can be
+nested. For example, a debug interrupt can safely be interrupted by an
+NMI. arch/x86_64/kernel/entry.S::paranoidentry adjusts the stack
+pointers on entry to and exit from all IST events, in theory allowing
+IST events with the same code to be nested. However in most cases, the
+stack size allocated to an IST assumes no nesting for the same code.
+If that assumption is ever broken then the stacks will become corrupt.
+
+The currently assigned IST stacks are:
+
+* ESTACK_DF. EXCEPTION_STKSZ (PAGE_SIZE).
+
+ Used for interrupt 8 - Double Fault Exception (#DF).
+
+ Invoked when handling one exception causes another exception. Happens
+ when the kernel is very confused (e.g. kernel stack pointer corrupt).
+ Using a separate stack allows the kernel to recover from it well enough
+ in many cases to still output an oops.
+
+* ESTACK_NMI. EXCEPTION_STKSZ (PAGE_SIZE).
+
+ Used for non-maskable interrupts (NMI).
+
+ NMI can be delivered at any time, including when the kernel is in the
+ middle of switching stacks. Using IST for NMI events avoids making
+ assumptions about the previous state of the kernel stack.
+
+* ESTACK_DB. EXCEPTION_STKSZ (PAGE_SIZE).
+
+ Used for hardware debug interrupts (interrupt 1) and for software
+ debug interrupts (INT3).
+
+ When debugging a kernel, debug interrupts (both hardware and
+ software) can occur at any time. Using IST for these interrupts
+ avoids making assumptions about the previous state of the kernel
+ stack.
+
+ To handle nested #DB correctly there exist two instances of DB stacks. On
+ #DB entry the IST stackpointer for #DB is switched to the second instance
+ so a nested #DB starts from a clean stack. The nested #DB switches
+ the IST stackpointer to a guard hole to catch triple nesting.
+
+* ESTACK_MCE. EXCEPTION_STKSZ (PAGE_SIZE).
+
+ Used for interrupt 18 - Machine Check Exception (#MC).
+
+ MCE can be delivered at any time, including when the kernel is in the
+ middle of switching stacks. Using IST for MCE events avoids making
+ assumptions about the previous state of the kernel stack.
+
+For more details see the Intel IA32 or AMD AMD64 architecture manuals.
+
+
+Printing backtraces on x86
+==========================
+
+The question about the '?' preceding function names in an x86 stacktrace
+keeps popping up, here's an indepth explanation. It helps if the reader
+stares at print_context_stack() and the whole machinery in and around
+arch/x86/kernel/dumpstack.c.
+
+Adapted from Ingo's mail, Message-ID: <20150521101614.GA10889@gmail.com>:
+
+We always scan the full kernel stack for return addresses stored on
+the kernel stack(s) [1]_, from stack top to stack bottom, and print out
+anything that 'looks like' a kernel text address.
+
+If it fits into the frame pointer chain, we print it without a question
+mark, knowing that it's part of the real backtrace.
+
+If the address does not fit into our expected frame pointer chain we
+still print it, but we print a '?'. It can mean two things:
+
+ - either the address is not part of the call chain: it's just stale
+ values on the kernel stack, from earlier function calls. This is
+ the common case.
+
+ - or it is part of the call chain, but the frame pointer was not set
+ up properly within the function, so we don't recognize it.
+
+This way we will always print out the real call chain (plus a few more
+entries), regardless of whether the frame pointer was set up correctly
+or not - but in most cases we'll get the call chain right as well. The
+entries printed are strictly in stack order, so you can deduce more
+information from that as well.
+
+The most important property of this method is that we _never_ lose
+information: we always strive to print _all_ addresses on the stack(s)
+that look like kernel text addresses, so if debug information is wrong,
+we still print out the real call chain as well - just with more question
+marks than ideal.
+
+.. [1] For things like IRQ and IST stacks, we also scan those stacks, in
+ the right order, and try to cross from one stack into another
+ reconstructing the call chain. This works most of the time.
diff --git a/Documentation/x86/mds.rst b/Documentation/x86/mds.rst
new file mode 100644
index 000000000..5d4330be2
--- /dev/null
+++ b/Documentation/x86/mds.rst
@@ -0,0 +1,193 @@
+Microarchitectural Data Sampling (MDS) mitigation
+=================================================
+
+.. _mds:
+
+Overview
+--------
+
+Microarchitectural Data Sampling (MDS) is a family of side channel attacks
+on internal buffers in Intel CPUs. The variants are:
+
+ - Microarchitectural Store Buffer Data Sampling (MSBDS) (CVE-2018-12126)
+ - Microarchitectural Fill Buffer Data Sampling (MFBDS) (CVE-2018-12130)
+ - Microarchitectural Load Port Data Sampling (MLPDS) (CVE-2018-12127)
+ - Microarchitectural Data Sampling Uncacheable Memory (MDSUM) (CVE-2019-11091)
+
+MSBDS leaks Store Buffer Entries which can be speculatively forwarded to a
+dependent load (store-to-load forwarding) as an optimization. The forward
+can also happen to a faulting or assisting load operation for a different
+memory address, which can be exploited under certain conditions. Store
+buffers are partitioned between Hyper-Threads so cross thread forwarding is
+not possible. But if a thread enters or exits a sleep state the store
+buffer is repartitioned which can expose data from one thread to the other.
+
+MFBDS leaks Fill Buffer Entries. Fill buffers are used internally to manage
+L1 miss situations and to hold data which is returned or sent in response
+to a memory or I/O operation. Fill buffers can forward data to a load
+operation and also write data to the cache. When the fill buffer is
+deallocated it can retain the stale data of the preceding operations which
+can then be forwarded to a faulting or assisting load operation, which can
+be exploited under certain conditions. Fill buffers are shared between
+Hyper-Threads so cross thread leakage is possible.
+
+MLPDS leaks Load Port Data. Load ports are used to perform load operations
+from memory or I/O. The received data is then forwarded to the register
+file or a subsequent operation. In some implementations the Load Port can
+contain stale data from a previous operation which can be forwarded to
+faulting or assisting loads under certain conditions, which again can be
+exploited eventually. Load ports are shared between Hyper-Threads so cross
+thread leakage is possible.
+
+MDSUM is a special case of MSBDS, MFBDS and MLPDS. An uncacheable load from
+memory that takes a fault or assist can leave data in a microarchitectural
+structure that may later be observed using one of the same methods used by
+MSBDS, MFBDS or MLPDS.
+
+Exposure assumptions
+--------------------
+
+It is assumed that attack code resides in user space or in a guest with one
+exception. The rationale behind this assumption is that the code construct
+needed for exploiting MDS requires:
+
+ - to control the load to trigger a fault or assist
+
+ - to have a disclosure gadget which exposes the speculatively accessed
+ data for consumption through a side channel.
+
+ - to control the pointer through which the disclosure gadget exposes the
+ data
+
+The existence of such a construct in the kernel cannot be excluded with
+100% certainty, but the complexity involved makes it extremly unlikely.
+
+There is one exception, which is untrusted BPF. The functionality of
+untrusted BPF is limited, but it needs to be thoroughly investigated
+whether it can be used to create such a construct.
+
+
+Mitigation strategy
+-------------------
+
+All variants have the same mitigation strategy at least for the single CPU
+thread case (SMT off): Force the CPU to clear the affected buffers.
+
+This is achieved by using the otherwise unused and obsolete VERW
+instruction in combination with a microcode update. The microcode clears
+the affected CPU buffers when the VERW instruction is executed.
+
+For virtualization there are two ways to achieve CPU buffer
+clearing. Either the modified VERW instruction or via the L1D Flush
+command. The latter is issued when L1TF mitigation is enabled so the extra
+VERW can be avoided. If the CPU is not affected by L1TF then VERW needs to
+be issued.
+
+If the VERW instruction with the supplied segment selector argument is
+executed on a CPU without the microcode update there is no side effect
+other than a small number of pointlessly wasted CPU cycles.
+
+This does not protect against cross Hyper-Thread attacks except for MSBDS
+which is only exploitable cross Hyper-thread when one of the Hyper-Threads
+enters a C-state.
+
+The kernel provides a function to invoke the buffer clearing:
+
+ mds_clear_cpu_buffers()
+
+The mitigation is invoked on kernel/userspace, hypervisor/guest and C-state
+(idle) transitions.
+
+As a special quirk to address virtualization scenarios where the host has
+the microcode updated, but the hypervisor does not (yet) expose the
+MD_CLEAR CPUID bit to guests, the kernel issues the VERW instruction in the
+hope that it might actually clear the buffers. The state is reflected
+accordingly.
+
+According to current knowledge additional mitigations inside the kernel
+itself are not required because the necessary gadgets to expose the leaked
+data cannot be controlled in a way which allows exploitation from malicious
+user space or VM guests.
+
+Kernel internal mitigation modes
+--------------------------------
+
+ ======= ============================================================
+ off Mitigation is disabled. Either the CPU is not affected or
+ mds=off is supplied on the kernel command line
+
+ full Mitigation is enabled. CPU is affected and MD_CLEAR is
+ advertised in CPUID.
+
+ vmwerv Mitigation is enabled. CPU is affected and MD_CLEAR is not
+ advertised in CPUID. That is mainly for virtualization
+ scenarios where the host has the updated microcode but the
+ hypervisor does not expose MD_CLEAR in CPUID. It's a best
+ effort approach without guarantee.
+ ======= ============================================================
+
+If the CPU is affected and mds=off is not supplied on the kernel command
+line then the kernel selects the appropriate mitigation mode depending on
+the availability of the MD_CLEAR CPUID bit.
+
+Mitigation points
+-----------------
+
+1. Return to user space
+^^^^^^^^^^^^^^^^^^^^^^^
+
+ When transitioning from kernel to user space the CPU buffers are flushed
+ on affected CPUs when the mitigation is not disabled on the kernel
+ command line. The migitation is enabled through the static key
+ mds_user_clear.
+
+ The mitigation is invoked in prepare_exit_to_usermode() which covers
+ all but one of the kernel to user space transitions. The exception
+ is when we return from a Non Maskable Interrupt (NMI), which is
+ handled directly in do_nmi().
+
+ (The reason that NMI is special is that prepare_exit_to_usermode() can
+ enable IRQs. In NMI context, NMIs are blocked, and we don't want to
+ enable IRQs with NMIs blocked.)
+
+
+2. C-State transition
+^^^^^^^^^^^^^^^^^^^^^
+
+ When a CPU goes idle and enters a C-State the CPU buffers need to be
+ cleared on affected CPUs when SMT is active. This addresses the
+ repartitioning of the store buffer when one of the Hyper-Threads enters
+ a C-State.
+
+ When SMT is inactive, i.e. either the CPU does not support it or all
+ sibling threads are offline CPU buffer clearing is not required.
+
+ The idle clearing is enabled on CPUs which are only affected by MSBDS
+ and not by any other MDS variant. The other MDS variants cannot be
+ protected against cross Hyper-Thread attacks because the Fill Buffer and
+ the Load Ports are shared. So on CPUs affected by other variants, the
+ idle clearing would be a window dressing exercise and is therefore not
+ activated.
+
+ The invocation is controlled by the static key mds_idle_clear which is
+ switched depending on the chosen mitigation mode and the SMT state of
+ the system.
+
+ The buffer clear is only invoked before entering the C-State to prevent
+ that stale data from the idling CPU from spilling to the Hyper-Thread
+ sibling after the store buffer got repartitioned and all entries are
+ available to the non idle sibling.
+
+ When coming out of idle the store buffer is partitioned again so each
+ sibling has half of it available. The back from idle CPU could be then
+ speculatively exposed to contents of the sibling. The buffers are
+ flushed either on exit to user space or on VMENTER so malicious code
+ in user space or the guest cannot speculatively access them.
+
+ The mitigation is hooked into all variants of halt()/mwait(), but does
+ not cover the legacy ACPI IO-Port mechanism because the ACPI idle driver
+ has been superseded by the intel_idle driver around 2010 and is
+ preferred on all affected CPUs which are expected to gain the MD_CLEAR
+ functionality in microcode. Aside of that the IO-Port mechanism is a
+ legacy interface which is only used on older systems which are either
+ not affected or do not receive microcode updates anymore.
diff --git a/Documentation/x86/microcode.rst b/Documentation/x86/microcode.rst
new file mode 100644
index 000000000..b627c6f36
--- /dev/null
+++ b/Documentation/x86/microcode.rst
@@ -0,0 +1,240 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==========================
+The Linux Microcode Loader
+==========================
+
+:Authors: - Fenghua Yu <fenghua.yu@intel.com>
+ - Borislav Petkov <bp@suse.de>
+ - Ashok Raj <ashok.raj@intel.com>
+
+The kernel has a x86 microcode loading facility which is supposed to
+provide microcode loading methods in the OS. Potential use cases are
+updating the microcode on platforms beyond the OEM End-Of-Life support,
+and updating the microcode on long-running systems without rebooting.
+
+The loader supports three loading methods:
+
+Early load microcode
+====================
+
+The kernel can update microcode very early during boot. Loading
+microcode early can fix CPU issues before they are observed during
+kernel boot time.
+
+The microcode is stored in an initrd file. During boot, it is read from
+it and loaded into the CPU cores.
+
+The format of the combined initrd image is microcode in (uncompressed)
+cpio format followed by the (possibly compressed) initrd image. The
+loader parses the combined initrd image during boot.
+
+The microcode files in cpio name space are:
+
+on Intel:
+ kernel/x86/microcode/GenuineIntel.bin
+on AMD :
+ kernel/x86/microcode/AuthenticAMD.bin
+
+During BSP (BootStrapping Processor) boot (pre-SMP), the kernel
+scans the microcode file in the initrd. If microcode matching the
+CPU is found, it will be applied in the BSP and later on in all APs
+(Application Processors).
+
+The loader also saves the matching microcode for the CPU in memory.
+Thus, the cached microcode patch is applied when CPUs resume from a
+sleep state.
+
+Here's a crude example how to prepare an initrd with microcode (this is
+normally done automatically by the distribution, when recreating the
+initrd, so you don't really have to do it yourself. It is documented
+here for future reference only).
+::
+
+ #!/bin/bash
+
+ if [ -z "$1" ]; then
+ echo "You need to supply an initrd file"
+ exit 1
+ fi
+
+ INITRD="$1"
+
+ DSTDIR=kernel/x86/microcode
+ TMPDIR=/tmp/initrd
+
+ rm -rf $TMPDIR
+
+ mkdir $TMPDIR
+ cd $TMPDIR
+ mkdir -p $DSTDIR
+
+ if [ -d /lib/firmware/amd-ucode ]; then
+ cat /lib/firmware/amd-ucode/microcode_amd*.bin > $DSTDIR/AuthenticAMD.bin
+ fi
+
+ if [ -d /lib/firmware/intel-ucode ]; then
+ cat /lib/firmware/intel-ucode/* > $DSTDIR/GenuineIntel.bin
+ fi
+
+ find . | cpio -o -H newc >../ucode.cpio
+ cd ..
+ mv $INITRD $INITRD.orig
+ cat ucode.cpio $INITRD.orig > $INITRD
+
+ rm -rf $TMPDIR
+
+
+The system needs to have the microcode packages installed into
+/lib/firmware or you need to fixup the paths above if yours are
+somewhere else and/or you've downloaded them directly from the processor
+vendor's site.
+
+Late loading
+============
+
+You simply install the microcode packages your distro supplies and
+run::
+
+ # echo 1 > /sys/devices/system/cpu/microcode/reload
+
+as root.
+
+The loading mechanism looks for microcode blobs in
+/lib/firmware/{intel-ucode,amd-ucode}. The default distro installation
+packages already put them there.
+
+Since kernel 5.19, late loading is not enabled by default.
+
+The /dev/cpu/microcode method has been removed in 5.19.
+
+Why is late loading dangerous?
+==============================
+
+Synchronizing all CPUs
+----------------------
+
+The microcode engine which receives the microcode update is shared
+between the two logical threads in a SMT system. Therefore, when
+the update is executed on one SMT thread of the core, the sibling
+"automatically" gets the update.
+
+Since the microcode can "simulate" MSRs too, while the microcode update
+is in progress, those simulated MSRs transiently cease to exist. This
+can result in unpredictable results if the SMT sibling thread happens to
+be in the middle of an access to such an MSR. The usual observation is
+that such MSR accesses cause #GPs to be raised to signal that former are
+not present.
+
+The disappearing MSRs are just one common issue which is being observed.
+Any other instruction that's being patched and gets concurrently
+executed by the other SMT sibling, can also result in similar,
+unpredictable behavior.
+
+To eliminate this case, a stop_machine()-based CPU synchronization was
+introduced as a way to guarantee that all logical CPUs will not execute
+any code but just wait in a spin loop, polling an atomic variable.
+
+While this took care of device or external interrupts, IPIs including
+LVT ones, such as CMCI etc, it cannot address other special interrupts
+that can't be shut off. Those are Machine Check (#MC), System Management
+(#SMI) and Non-Maskable interrupts (#NMI).
+
+Machine Checks
+--------------
+
+Machine Checks (#MC) are non-maskable. There are two kinds of MCEs.
+Fatal un-recoverable MCEs and recoverable MCEs. While un-recoverable
+errors are fatal, recoverable errors can also happen in kernel context
+are also treated as fatal by the kernel.
+
+On certain Intel machines, MCEs are also broadcast to all threads in a
+system. If one thread is in the middle of executing WRMSR, a MCE will be
+taken at the end of the flow. Either way, they will wait for the thread
+performing the wrmsr(0x79) to rendezvous in the MCE handler and shutdown
+eventually if any of the threads in the system fail to check in to the
+MCE rendezvous.
+
+To be paranoid and get predictable behavior, the OS can choose to set
+MCG_STATUS.MCIP. Since MCEs can be at most one in a system, if an
+MCE was signaled, the above condition will promote to a system reset
+automatically. OS can turn off MCIP at the end of the update for that
+core.
+
+System Management Interrupt
+---------------------------
+
+SMIs are also broadcast to all CPUs in the platform. Microcode update
+requests exclusive access to the core before writing to MSR 0x79. So if
+it does happen such that, one thread is in WRMSR flow, and the 2nd got
+an SMI, that thread will be stopped in the first instruction in the SMI
+handler.
+
+Since the secondary thread is stopped in the first instruction in SMI,
+there is very little chance that it would be in the middle of executing
+an instruction being patched. Plus OS has no way to stop SMIs from
+happening.
+
+Non-Maskable Interrupts
+-----------------------
+
+When thread0 of a core is doing the microcode update, if thread1 is
+pulled into NMI, that can cause unpredictable behavior due to the
+reasons above.
+
+OS can choose a variety of methods to avoid running into this situation.
+
+
+Is the microcode suitable for late loading?
+-------------------------------------------
+
+Late loading is done when the system is fully operational and running
+real workloads. Late loading behavior depends on what the base patch on
+the CPU is before upgrading to the new patch.
+
+This is true for Intel CPUs.
+
+Consider, for example, a CPU has patch level 1 and the update is to
+patch level 3.
+
+Between patch1 and patch3, patch2 might have deprecated a software-visible
+feature.
+
+This is unacceptable if software is even potentially using that feature.
+For instance, say MSR_X is no longer available after an update,
+accessing that MSR will cause a #GP fault.
+
+Basically there is no way to declare a new microcode update suitable
+for late-loading. This is another one of the problems that caused late
+loading to be not enabled by default.
+
+Builtin microcode
+=================
+
+The loader supports also loading of a builtin microcode supplied through
+the regular builtin firmware method CONFIG_EXTRA_FIRMWARE. Only 64-bit is
+currently supported.
+
+Here's an example::
+
+ CONFIG_EXTRA_FIRMWARE="intel-ucode/06-3a-09 amd-ucode/microcode_amd_fam15h.bin"
+ CONFIG_EXTRA_FIRMWARE_DIR="/lib/firmware"
+
+This basically means, you have the following tree structure locally::
+
+ /lib/firmware/
+ |-- amd-ucode
+ ...
+ | |-- microcode_amd_fam15h.bin
+ ...
+ |-- intel-ucode
+ ...
+ | |-- 06-3a-09
+ ...
+
+so that the build system can find those files and integrate them into
+the final kernel image. The early loader finds them and applies them.
+
+Needless to say, this method is not the most flexible one because it
+requires rebuilding the kernel each time updated microcode from the CPU
+vendor is available.
diff --git a/Documentation/x86/mtrr.rst b/Documentation/x86/mtrr.rst
new file mode 100644
index 000000000..9f0b18517
--- /dev/null
+++ b/Documentation/x86/mtrr.rst
@@ -0,0 +1,354 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=========================================
+MTRR (Memory Type Range Register) control
+=========================================
+
+:Authors: - Richard Gooch <rgooch@atnf.csiro.au> - 3 Jun 1999
+ - Luis R. Rodriguez <mcgrof@do-not-panic.com> - April 9, 2015
+
+
+Phasing out MTRR use
+====================
+
+MTRR use is replaced on modern x86 hardware with PAT. Direct MTRR use by
+drivers on Linux is now completely phased out, device drivers should use
+arch_phys_wc_add() in combination with ioremap_wc() to make MTRR effective on
+non-PAT systems while a no-op but equally effective on PAT enabled systems.
+
+Even if Linux does not use MTRRs directly, some x86 platform firmware may still
+set up MTRRs early before booting the OS. They do this as some platform
+firmware may still have implemented access to MTRRs which would be controlled
+and handled by the platform firmware directly. An example of platform use of
+MTRRs is through the use of SMI handlers, one case could be for fan control,
+the platform code would need uncachable access to some of its fan control
+registers. Such platform access does not need any Operating System MTRR code in
+place other than mtrr_type_lookup() to ensure any OS specific mapping requests
+are aligned with platform MTRR setup. If MTRRs are only set up by the platform
+firmware code though and the OS does not make any specific MTRR mapping
+requests mtrr_type_lookup() should always return MTRR_TYPE_INVALID.
+
+For details refer to Documentation/x86/pat.rst.
+
+.. tip::
+ On Intel P6 family processors (Pentium Pro, Pentium II and later)
+ the Memory Type Range Registers (MTRRs) may be used to control
+ processor access to memory ranges. This is most useful when you have
+ a video (VGA) card on a PCI or AGP bus. Enabling write-combining
+ allows bus write transfers to be combined into a larger transfer
+ before bursting over the PCI/AGP bus. This can increase performance
+ of image write operations 2.5 times or more.
+
+ The Cyrix 6x86, 6x86MX and M II processors have Address Range
+ Registers (ARRs) which provide a similar functionality to MTRRs. For
+ these, the ARRs are used to emulate the MTRRs.
+
+ The AMD K6-2 (stepping 8 and above) and K6-3 processors have two
+ MTRRs. These are supported. The AMD Athlon family provide 8 Intel
+ style MTRRs.
+
+ The Centaur C6 (WinChip) has 8 MCRs, allowing write-combining. These
+ are supported.
+
+ The VIA Cyrix III and VIA C3 CPUs offer 8 Intel style MTRRs.
+
+ The CONFIG_MTRR option creates a /proc/mtrr file which may be used
+ to manipulate your MTRRs. Typically the X server should use
+ this. This should have a reasonably generic interface so that
+ similar control registers on other processors can be easily
+ supported.
+
+There are two interfaces to /proc/mtrr: one is an ASCII interface
+which allows you to read and write. The other is an ioctl()
+interface. The ASCII interface is meant for administration. The
+ioctl() interface is meant for C programs (i.e. the X server). The
+interfaces are described below, with sample commands and C code.
+
+
+Reading MTRRs from the shell
+============================
+::
+
+ % cat /proc/mtrr
+ reg00: base=0x00000000 ( 0MB), size= 128MB: write-back, count=1
+ reg01: base=0x08000000 ( 128MB), size= 64MB: write-back, count=1
+
+Creating MTRRs from the C-shell::
+
+ # echo "base=0xf8000000 size=0x400000 type=write-combining" >! /proc/mtrr
+
+or if you use bash::
+
+ # echo "base=0xf8000000 size=0x400000 type=write-combining" >| /proc/mtrr
+
+And the result thereof::
+
+ % cat /proc/mtrr
+ reg00: base=0x00000000 ( 0MB), size= 128MB: write-back, count=1
+ reg01: base=0x08000000 ( 128MB), size= 64MB: write-back, count=1
+ reg02: base=0xf8000000 (3968MB), size= 4MB: write-combining, count=1
+
+This is for video RAM at base address 0xf8000000 and size 4 megabytes. To
+find out your base address, you need to look at the output of your X
+server, which tells you where the linear framebuffer address is. A
+typical line that you may get is::
+
+ (--) S3: PCI: 968 rev 0, Linear FB @ 0xf8000000
+
+Note that you should only use the value from the X server, as it may
+move the framebuffer base address, so the only value you can trust is
+that reported by the X server.
+
+To find out the size of your framebuffer (what, you don't actually
+know?), the following line will tell you::
+
+ (--) S3: videoram: 4096k
+
+That's 4 megabytes, which is 0x400000 bytes (in hexadecimal).
+A patch is being written for XFree86 which will make this automatic:
+in other words the X server will manipulate /proc/mtrr using the
+ioctl() interface, so users won't have to do anything. If you use a
+commercial X server, lobby your vendor to add support for MTRRs.
+
+
+Creating overlapping MTRRs
+==========================
+::
+
+ %echo "base=0xfb000000 size=0x1000000 type=write-combining" >/proc/mtrr
+ %echo "base=0xfb000000 size=0x1000 type=uncachable" >/proc/mtrr
+
+And the results::
+
+ % cat /proc/mtrr
+ reg00: base=0x00000000 ( 0MB), size= 64MB: write-back, count=1
+ reg01: base=0xfb000000 (4016MB), size= 16MB: write-combining, count=1
+ reg02: base=0xfb000000 (4016MB), size= 4kB: uncachable, count=1
+
+Some cards (especially Voodoo Graphics boards) need this 4 kB area
+excluded from the beginning of the region because it is used for
+registers.
+
+NOTE: You can only create type=uncachable region, if the first
+region that you created is type=write-combining.
+
+
+Removing MTRRs from the C-shel
+==============================
+::
+
+ % echo "disable=2" >! /proc/mtrr
+
+or using bash::
+
+ % echo "disable=2" >| /proc/mtrr
+
+
+Reading MTRRs from a C program using ioctl()'s
+==============================================
+::
+
+ /* mtrr-show.c
+
+ Source file for mtrr-show (example program to show MTRRs using ioctl()'s)
+
+ Copyright (C) 1997-1998 Richard Gooch
+
+ This program is free software; you can redistribute it and/or modify
+ it under the terms of the GNU General Public License as published by
+ the Free Software Foundation; either version 2 of the License, or
+ (at your option) any later version.
+
+ This program is distributed in the hope that it will be useful,
+ but WITHOUT ANY WARRANTY; without even the implied warranty of
+ MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the
+ GNU General Public License for more details.
+
+ You should have received a copy of the GNU General Public License
+ along with this program; if not, write to the Free Software
+ Foundation, Inc., 675 Mass Ave, Cambridge, MA 02139, USA.
+
+ Richard Gooch may be reached by email at rgooch@atnf.csiro.au
+ The postal address is:
+ Richard Gooch, c/o ATNF, P. O. Box 76, Epping, N.S.W., 2121, Australia.
+ */
+
+ /*
+ This program will use an ioctl() on /proc/mtrr to show the current MTRR
+ settings. This is an alternative to reading /proc/mtrr.
+
+
+ Written by Richard Gooch 17-DEC-1997
+
+ Last updated by Richard Gooch 2-MAY-1998
+
+
+ */
+ #include <stdio.h>
+ #include <stdlib.h>
+ #include <string.h>
+ #include <sys/types.h>
+ #include <sys/stat.h>
+ #include <fcntl.h>
+ #include <sys/ioctl.h>
+ #include <errno.h>
+ #include <asm/mtrr.h>
+
+ #define TRUE 1
+ #define FALSE 0
+ #define ERRSTRING strerror (errno)
+
+ static char *mtrr_strings[MTRR_NUM_TYPES] =
+ {
+ "uncachable", /* 0 */
+ "write-combining", /* 1 */
+ "?", /* 2 */
+ "?", /* 3 */
+ "write-through", /* 4 */
+ "write-protect", /* 5 */
+ "write-back", /* 6 */
+ };
+
+ int main ()
+ {
+ int fd;
+ struct mtrr_gentry gentry;
+
+ if ( ( fd = open ("/proc/mtrr", O_RDONLY, 0) ) == -1 )
+ {
+ if (errno == ENOENT)
+ {
+ fputs ("/proc/mtrr not found: not supported or you don't have a PPro?\n",
+ stderr);
+ exit (1);
+ }
+ fprintf (stderr, "Error opening /proc/mtrr\t%s\n", ERRSTRING);
+ exit (2);
+ }
+ for (gentry.regnum = 0; ioctl (fd, MTRRIOC_GET_ENTRY, &gentry) == 0;
+ ++gentry.regnum)
+ {
+ if (gentry.size < 1)
+ {
+ fprintf (stderr, "Register: %u disabled\n", gentry.regnum);
+ continue;
+ }
+ fprintf (stderr, "Register: %u base: 0x%lx size: 0x%lx type: %s\n",
+ gentry.regnum, gentry.base, gentry.size,
+ mtrr_strings[gentry.type]);
+ }
+ if (errno == EINVAL) exit (0);
+ fprintf (stderr, "Error doing ioctl(2) on /dev/mtrr\t%s\n", ERRSTRING);
+ exit (3);
+ } /* End Function main */
+
+
+Creating MTRRs from a C programme using ioctl()'s
+=================================================
+::
+
+ /* mtrr-add.c
+
+ Source file for mtrr-add (example programme to add an MTRRs using ioctl())
+
+ Copyright (C) 1997-1998 Richard Gooch
+
+ This program is free software; you can redistribute it and/or modify
+ it under the terms of the GNU General Public License as published by
+ the Free Software Foundation; either version 2 of the License, or
+ (at your option) any later version.
+
+ This program is distributed in the hope that it will be useful,
+ but WITHOUT ANY WARRANTY; without even the implied warranty of
+ MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the
+ GNU General Public License for more details.
+
+ You should have received a copy of the GNU General Public License
+ along with this program; if not, write to the Free Software
+ Foundation, Inc., 675 Mass Ave, Cambridge, MA 02139, USA.
+
+ Richard Gooch may be reached by email at rgooch@atnf.csiro.au
+ The postal address is:
+ Richard Gooch, c/o ATNF, P. O. Box 76, Epping, N.S.W., 2121, Australia.
+ */
+
+ /*
+ This programme will use an ioctl() on /proc/mtrr to add an entry. The first
+ available mtrr is used. This is an alternative to writing /proc/mtrr.
+
+
+ Written by Richard Gooch 17-DEC-1997
+
+ Last updated by Richard Gooch 2-MAY-1998
+
+
+ */
+ #include <stdio.h>
+ #include <string.h>
+ #include <stdlib.h>
+ #include <unistd.h>
+ #include <sys/types.h>
+ #include <sys/stat.h>
+ #include <fcntl.h>
+ #include <sys/ioctl.h>
+ #include <errno.h>
+ #include <asm/mtrr.h>
+
+ #define TRUE 1
+ #define FALSE 0
+ #define ERRSTRING strerror (errno)
+
+ static char *mtrr_strings[MTRR_NUM_TYPES] =
+ {
+ "uncachable", /* 0 */
+ "write-combining", /* 1 */
+ "?", /* 2 */
+ "?", /* 3 */
+ "write-through", /* 4 */
+ "write-protect", /* 5 */
+ "write-back", /* 6 */
+ };
+
+ int main (int argc, char **argv)
+ {
+ int fd;
+ struct mtrr_sentry sentry;
+
+ if (argc != 4)
+ {
+ fprintf (stderr, "Usage:\tmtrr-add base size type\n");
+ exit (1);
+ }
+ sentry.base = strtoul (argv[1], NULL, 0);
+ sentry.size = strtoul (argv[2], NULL, 0);
+ for (sentry.type = 0; sentry.type < MTRR_NUM_TYPES; ++sentry.type)
+ {
+ if (strcmp (argv[3], mtrr_strings[sentry.type]) == 0) break;
+ }
+ if (sentry.type >= MTRR_NUM_TYPES)
+ {
+ fprintf (stderr, "Illegal type: \"%s\"\n", argv[3]);
+ exit (2);
+ }
+ if ( ( fd = open ("/proc/mtrr", O_WRONLY, 0) ) == -1 )
+ {
+ if (errno == ENOENT)
+ {
+ fputs ("/proc/mtrr not found: not supported or you don't have a PPro?\n",
+ stderr);
+ exit (3);
+ }
+ fprintf (stderr, "Error opening /proc/mtrr\t%s\n", ERRSTRING);
+ exit (4);
+ }
+ if (ioctl (fd, MTRRIOC_ADD_ENTRY, &sentry) == -1)
+ {
+ fprintf (stderr, "Error doing ioctl(2) on /dev/mtrr\t%s\n", ERRSTRING);
+ exit (5);
+ }
+ fprintf (stderr, "Sleeping for 5 seconds so you can see the new entry\n");
+ sleep (5);
+ close (fd);
+ fputs ("I've just closed /proc/mtrr so now the new entry should be gone\n",
+ stderr);
+ } /* End Function main */
diff --git a/Documentation/x86/orc-unwinder.rst b/Documentation/x86/orc-unwinder.rst
new file mode 100644
index 000000000..cdb257015
--- /dev/null
+++ b/Documentation/x86/orc-unwinder.rst
@@ -0,0 +1,182 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+============
+ORC unwinder
+============
+
+Overview
+========
+
+The kernel CONFIG_UNWINDER_ORC option enables the ORC unwinder, which is
+similar in concept to a DWARF unwinder. The difference is that the
+format of the ORC data is much simpler than DWARF, which in turn allows
+the ORC unwinder to be much simpler and faster.
+
+The ORC data consists of unwind tables which are generated by objtool.
+They contain out-of-band data which is used by the in-kernel ORC
+unwinder. Objtool generates the ORC data by first doing compile-time
+stack metadata validation (CONFIG_STACK_VALIDATION). After analyzing
+all the code paths of a .o file, it determines information about the
+stack state at each instruction address in the file and outputs that
+information to the .orc_unwind and .orc_unwind_ip sections.
+
+The per-object ORC sections are combined at link time and are sorted and
+post-processed at boot time. The unwinder uses the resulting data to
+correlate instruction addresses with their stack states at run time.
+
+
+ORC vs frame pointers
+=====================
+
+With frame pointers enabled, GCC adds instrumentation code to every
+function in the kernel. The kernel's .text size increases by about
+3.2%, resulting in a broad kernel-wide slowdown. Measurements by Mel
+Gorman [1]_ have shown a slowdown of 5-10% for some workloads.
+
+In contrast, the ORC unwinder has no effect on text size or runtime
+performance, because the debuginfo is out of band. So if you disable
+frame pointers and enable the ORC unwinder, you get a nice performance
+improvement across the board, and still have reliable stack traces.
+
+Ingo Molnar says:
+
+ "Note that it's not just a performance improvement, but also an
+ instruction cache locality improvement: 3.2% .text savings almost
+ directly transform into a similarly sized reduction in cache
+ footprint. That can transform to even higher speedups for workloads
+ whose cache locality is borderline."
+
+Another benefit of ORC compared to frame pointers is that it can
+reliably unwind across interrupts and exceptions. Frame pointer based
+unwinds can sometimes skip the caller of the interrupted function, if it
+was a leaf function or if the interrupt hit before the frame pointer was
+saved.
+
+The main disadvantage of the ORC unwinder compared to frame pointers is
+that it needs more memory to store the ORC unwind tables: roughly 2-4MB
+depending on the kernel config.
+
+
+ORC vs DWARF
+============
+
+ORC debuginfo's advantage over DWARF itself is that it's much simpler.
+It gets rid of the complex DWARF CFI state machine and also gets rid of
+the tracking of unnecessary registers. This allows the unwinder to be
+much simpler, meaning fewer bugs, which is especially important for
+mission critical oops code.
+
+The simpler debuginfo format also enables the unwinder to be much faster
+than DWARF, which is important for perf and lockdep. In a basic
+performance test by Jiri Slaby [2]_, the ORC unwinder was about 20x
+faster than an out-of-tree DWARF unwinder. (Note: That measurement was
+taken before some performance tweaks were added, which doubled
+performance, so the speedup over DWARF may be closer to 40x.)
+
+The ORC data format does have a few downsides compared to DWARF. ORC
+unwind tables take up ~50% more RAM (+1.3MB on an x86 defconfig kernel)
+than DWARF-based eh_frame tables.
+
+Another potential downside is that, as GCC evolves, it's conceivable
+that the ORC data may end up being *too* simple to describe the state of
+the stack for certain optimizations. But IMO this is unlikely because
+GCC saves the frame pointer for any unusual stack adjustments it does,
+so I suspect we'll really only ever need to keep track of the stack
+pointer and the frame pointer between call frames. But even if we do
+end up having to track all the registers DWARF tracks, at least we will
+still be able to control the format, e.g. no complex state machines.
+
+
+ORC unwind table generation
+===========================
+
+The ORC data is generated by objtool. With the existing compile-time
+stack metadata validation feature, objtool already follows all code
+paths, and so it already has all the information it needs to be able to
+generate ORC data from scratch. So it's an easy step to go from stack
+validation to ORC data generation.
+
+It should be possible to instead generate the ORC data with a simple
+tool which converts DWARF to ORC data. However, such a solution would
+be incomplete due to the kernel's extensive use of asm, inline asm, and
+special sections like exception tables.
+
+That could be rectified by manually annotating those special code paths
+using GNU assembler .cfi annotations in .S files, and homegrown
+annotations for inline asm in .c files. But asm annotations were tried
+in the past and were found to be unmaintainable. They were often
+incorrect/incomplete and made the code harder to read and keep updated.
+And based on looking at glibc code, annotating inline asm in .c files
+might be even worse.
+
+Objtool still needs a few annotations, but only in code which does
+unusual things to the stack like entry code. And even then, far fewer
+annotations are needed than what DWARF would need, so they're much more
+maintainable than DWARF CFI annotations.
+
+So the advantages of using objtool to generate ORC data are that it
+gives more accurate debuginfo, with very few annotations. It also
+insulates the kernel from toolchain bugs which can be very painful to
+deal with in the kernel since we often have to workaround issues in
+older versions of the toolchain for years.
+
+The downside is that the unwinder now becomes dependent on objtool's
+ability to reverse engineer GCC code flow. If GCC optimizations become
+too complicated for objtool to follow, the ORC data generation might
+stop working or become incomplete. (It's worth noting that livepatch
+already has such a dependency on objtool's ability to follow GCC code
+flow.)
+
+If newer versions of GCC come up with some optimizations which break
+objtool, we may need to revisit the current implementation. Some
+possible solutions would be asking GCC to make the optimizations more
+palatable, or having objtool use DWARF as an additional input, or
+creating a GCC plugin to assist objtool with its analysis. But for now,
+objtool follows GCC code quite well.
+
+
+Unwinder implementation details
+===============================
+
+Objtool generates the ORC data by integrating with the compile-time
+stack metadata validation feature, which is described in detail in
+tools/objtool/Documentation/objtool.txt. After analyzing all
+the code paths of a .o file, it creates an array of orc_entry structs,
+and a parallel array of instruction addresses associated with those
+structs, and writes them to the .orc_unwind and .orc_unwind_ip sections
+respectively.
+
+The ORC data is split into the two arrays for performance reasons, to
+make the searchable part of the data (.orc_unwind_ip) more compact. The
+arrays are sorted in parallel at boot time.
+
+Performance is further improved by the use of a fast lookup table which
+is created at runtime. The fast lookup table associates a given address
+with a range of indices for the .orc_unwind table, so that only a small
+subset of the table needs to be searched.
+
+
+Etymology
+=========
+
+Orcs, fearsome creatures of medieval folklore, are the Dwarves' natural
+enemies. Similarly, the ORC unwinder was created in opposition to the
+complexity and slowness of DWARF.
+
+"Although Orcs rarely consider multiple solutions to a problem, they do
+excel at getting things done because they are creatures of action, not
+thought." [3]_ Similarly, unlike the esoteric DWARF unwinder, the
+veracious ORC unwinder wastes no time or siloconic effort decoding
+variable-length zero-extended unsigned-integer byte-coded
+state-machine-based debug information entries.
+
+Similar to how Orcs frequently unravel the well-intentioned plans of
+their adversaries, the ORC unwinder frequently unravels stacks with
+brutal, unyielding efficiency.
+
+ORC stands for Oops Rewind Capability.
+
+
+.. [1] https://lore.kernel.org/r/20170602104048.jkkzssljsompjdwy@suse.de
+.. [2] https://lore.kernel.org/r/d2ca5435-6386-29b8-db87-7f227c2b713a@suse.cz
+.. [3] http://dustin.wikidot.com/half-orcs-and-orcs
diff --git a/Documentation/x86/pat.rst b/Documentation/x86/pat.rst
new file mode 100644
index 000000000..5d9017710
--- /dev/null
+++ b/Documentation/x86/pat.rst
@@ -0,0 +1,240 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==========================
+PAT (Page Attribute Table)
+==========================
+
+x86 Page Attribute Table (PAT) allows for setting the memory attribute at the
+page level granularity. PAT is complementary to the MTRR settings which allows
+for setting of memory types over physical address ranges. However, PAT is
+more flexible than MTRR due to its capability to set attributes at page level
+and also due to the fact that there are no hardware limitations on number of
+such attribute settings allowed. Added flexibility comes with guidelines for
+not having memory type aliasing for the same physical memory with multiple
+virtual addresses.
+
+PAT allows for different types of memory attributes. The most commonly used
+ones that will be supported at this time are:
+
+=== ==============
+WB Write-back
+UC Uncached
+WC Write-combined
+WT Write-through
+UC- Uncached Minus
+=== ==============
+
+
+PAT APIs
+========
+
+There are many different APIs in the kernel that allows setting of memory
+attributes at the page level. In order to avoid aliasing, these interfaces
+should be used thoughtfully. Below is a table of interfaces available,
+their intended usage and their memory attribute relationships. Internally,
+these APIs use a reserve_memtype()/free_memtype() interface on the physical
+address range to avoid any aliasing.
+
++------------------------+----------+--------------+------------------+
+| API | RAM | ACPI,... | Reserved/Holes |
++------------------------+----------+--------------+------------------+
+| ioremap | -- | UC- | UC- |
++------------------------+----------+--------------+------------------+
+| ioremap_cache | -- | WB | WB |
++------------------------+----------+--------------+------------------+
+| ioremap_uc | -- | UC | UC |
++------------------------+----------+--------------+------------------+
+| ioremap_wc | -- | -- | WC |
++------------------------+----------+--------------+------------------+
+| ioremap_wt | -- | -- | WT |
++------------------------+----------+--------------+------------------+
+| set_memory_uc, | UC- | -- | -- |
+| set_memory_wb | | | |
++------------------------+----------+--------------+------------------+
+| set_memory_wc, | WC | -- | -- |
+| set_memory_wb | | | |
++------------------------+----------+--------------+------------------+
+| set_memory_wt, | WT | -- | -- |
+| set_memory_wb | | | |
++------------------------+----------+--------------+------------------+
+| pci sysfs resource | -- | -- | UC- |
++------------------------+----------+--------------+------------------+
+| pci sysfs resource_wc | -- | -- | WC |
+| is IORESOURCE_PREFETCH | | | |
++------------------------+----------+--------------+------------------+
+| pci proc | -- | -- | UC- |
+| !PCIIOC_WRITE_COMBINE | | | |
++------------------------+----------+--------------+------------------+
+| pci proc | -- | -- | WC |
+| PCIIOC_WRITE_COMBINE | | | |
++------------------------+----------+--------------+------------------+
+| /dev/mem | -- | WB/WC/UC- | WB/WC/UC- |
+| read-write | | | |
++------------------------+----------+--------------+------------------+
+| /dev/mem | -- | UC- | UC- |
+| mmap SYNC flag | | | |
++------------------------+----------+--------------+------------------+
+| /dev/mem | -- | WB/WC/UC- | WB/WC/UC- |
+| mmap !SYNC flag | | | |
+| and | |(from existing| (from existing |
+| any alias to this area | |alias) | alias) |
++------------------------+----------+--------------+------------------+
+| /dev/mem | -- | WB | WB |
+| mmap !SYNC flag | | | |
+| no alias to this area | | | |
+| and | | | |
+| MTRR says WB | | | |
++------------------------+----------+--------------+------------------+
+| /dev/mem | -- | -- | UC- |
+| mmap !SYNC flag | | | |
+| no alias to this area | | | |
+| and | | | |
+| MTRR says !WB | | | |
++------------------------+----------+--------------+------------------+
+
+
+Advanced APIs for drivers
+=========================
+
+A. Exporting pages to users with remap_pfn_range, io_remap_pfn_range,
+vmf_insert_pfn.
+
+Drivers wanting to export some pages to userspace do it by using mmap
+interface and a combination of:
+
+ 1) pgprot_noncached()
+ 2) io_remap_pfn_range() or remap_pfn_range() or vmf_insert_pfn()
+
+With PAT support, a new API pgprot_writecombine is being added. So, drivers can
+continue to use the above sequence, with either pgprot_noncached() or
+pgprot_writecombine() in step 1, followed by step 2.
+
+In addition, step 2 internally tracks the region as UC or WC in memtype
+list in order to ensure no conflicting mapping.
+
+Note that this set of APIs only works with IO (non RAM) regions. If driver
+wants to export a RAM region, it has to do set_memory_uc() or set_memory_wc()
+as step 0 above and also track the usage of those pages and use set_memory_wb()
+before the page is freed to free pool.
+
+MTRR effects on PAT / non-PAT systems
+=====================================
+
+The following table provides the effects of using write-combining MTRRs when
+using ioremap*() calls on x86 for both non-PAT and PAT systems. Ideally
+mtrr_add() usage will be phased out in favor of arch_phys_wc_add() which will
+be a no-op on PAT enabled systems. The region over which a arch_phys_wc_add()
+is made, should already have been ioremapped with WC attributes or PAT entries,
+this can be done by using ioremap_wc() / set_memory_wc(). Devices which
+combine areas of IO memory desired to remain uncacheable with areas where
+write-combining is desirable should consider use of ioremap_uc() followed by
+set_memory_wc() to white-list effective write-combined areas. Such use is
+nevertheless discouraged as the effective memory type is considered
+implementation defined, yet this strategy can be used as last resort on devices
+with size-constrained regions where otherwise MTRR write-combining would
+otherwise not be effective.
+::
+
+ ==== ======= === ========================= =====================
+ MTRR Non-PAT PAT Linux ioremap value Effective memory type
+ ==== ======= === ========================= =====================
+ PAT Non-PAT | PAT
+ |PCD |
+ ||PWT |
+ ||| |
+ WC 000 WB _PAGE_CACHE_MODE_WB WC | WC
+ WC 001 WC _PAGE_CACHE_MODE_WC WC* | WC
+ WC 010 UC- _PAGE_CACHE_MODE_UC_MINUS WC* | UC
+ WC 011 UC _PAGE_CACHE_MODE_UC UC | UC
+ ==== ======= === ========================= =====================
+
+ (*) denotes implementation defined and is discouraged
+
+.. note:: -- in the above table mean "Not suggested usage for the API". Some
+ of the --'s are strictly enforced by the kernel. Some others are not really
+ enforced today, but may be enforced in future.
+
+For ioremap and pci access through /sys or /proc - The actual type returned
+can be more restrictive, in case of any existing aliasing for that address.
+For example: If there is an existing uncached mapping, a new ioremap_wc can
+return uncached mapping in place of write-combine requested.
+
+set_memory_[uc|wc|wt] and set_memory_wb should be used in pairs, where driver
+will first make a region uc, wc or wt and switch it back to wb after use.
+
+Over time writes to /proc/mtrr will be deprecated in favor of using PAT based
+interfaces. Users writing to /proc/mtrr are suggested to use above interfaces.
+
+Drivers should use ioremap_[uc|wc] to access PCI BARs with [uc|wc] access
+types.
+
+Drivers should use set_memory_[uc|wc|wt] to set access type for RAM ranges.
+
+
+PAT debugging
+=============
+
+With CONFIG_DEBUG_FS enabled, PAT memtype list can be examined by::
+
+ # mount -t debugfs debugfs /sys/kernel/debug
+ # cat /sys/kernel/debug/x86/pat_memtype_list
+ PAT memtype list:
+ uncached-minus @ 0x7fadf000-0x7fae0000
+ uncached-minus @ 0x7fb19000-0x7fb1a000
+ uncached-minus @ 0x7fb1a000-0x7fb1b000
+ uncached-minus @ 0x7fb1b000-0x7fb1c000
+ uncached-minus @ 0x7fb1c000-0x7fb1d000
+ uncached-minus @ 0x7fb1d000-0x7fb1e000
+ uncached-minus @ 0x7fb1e000-0x7fb25000
+ uncached-minus @ 0x7fb25000-0x7fb26000
+ uncached-minus @ 0x7fb26000-0x7fb27000
+ uncached-minus @ 0x7fb27000-0x7fb28000
+ uncached-minus @ 0x7fb28000-0x7fb2e000
+ uncached-minus @ 0x7fb2e000-0x7fb2f000
+ uncached-minus @ 0x7fb2f000-0x7fb30000
+ uncached-minus @ 0x7fb31000-0x7fb32000
+ uncached-minus @ 0x80000000-0x90000000
+
+This list shows physical address ranges and various PAT settings used to
+access those physical address ranges.
+
+Another, more verbose way of getting PAT related debug messages is with
+"debugpat" boot parameter. With this parameter, various debug messages are
+printed to dmesg log.
+
+PAT Initialization
+==================
+
+The following table describes how PAT is initialized under various
+configurations. The PAT MSR must be updated by Linux in order to support WC
+and WT attributes. Otherwise, the PAT MSR has the value programmed in it
+by the firmware. Note, Xen enables WC attribute in the PAT MSR for guests.
+
+ ==== ===== ========================== ========= =======
+ MTRR PAT Call Sequence PAT State PAT MSR
+ ==== ===== ========================== ========= =======
+ E E MTRR -> PAT init Enabled OS
+ E D MTRR -> PAT init Disabled -
+ D E MTRR -> PAT disable Disabled BIOS
+ D D MTRR -> PAT disable Disabled -
+ - np/E PAT -> PAT disable Disabled BIOS
+ - np/D PAT -> PAT disable Disabled -
+ E !P/E MTRR -> PAT init Disabled BIOS
+ D !P/E MTRR -> PAT disable Disabled BIOS
+ !M !P/E MTRR stub -> PAT disable Disabled BIOS
+ ==== ===== ========================== ========= =======
+
+ Legend
+
+ ========= =======================================
+ E Feature enabled in CPU
+ D Feature disabled/unsupported in CPU
+ np "nopat" boot option specified
+ !P CONFIG_X86_PAT option unset
+ !M CONFIG_MTRR option unset
+ Enabled PAT state set to enabled
+ Disabled PAT state set to disabled
+ OS PAT initializes PAT MSR with OS setting
+ BIOS PAT keeps PAT MSR with BIOS setting
+ ========= =======================================
+
diff --git a/Documentation/x86/pti.rst b/Documentation/x86/pti.rst
new file mode 100644
index 000000000..4b858a9ba
--- /dev/null
+++ b/Documentation/x86/pti.rst
@@ -0,0 +1,195 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==========================
+Page Table Isolation (PTI)
+==========================
+
+Overview
+========
+
+Page Table Isolation (pti, previously known as KAISER [1]_) is a
+countermeasure against attacks on the shared user/kernel address
+space such as the "Meltdown" approach [2]_.
+
+To mitigate this class of attacks, we create an independent set of
+page tables for use only when running userspace applications. When
+the kernel is entered via syscalls, interrupts or exceptions, the
+page tables are switched to the full "kernel" copy. When the system
+switches back to user mode, the user copy is used again.
+
+The userspace page tables contain only a minimal amount of kernel
+data: only what is needed to enter/exit the kernel such as the
+entry/exit functions themselves and the interrupt descriptor table
+(IDT). There are a few strictly unnecessary things that get mapped
+such as the first C function when entering an interrupt (see
+comments in pti.c).
+
+This approach helps to ensure that side-channel attacks leveraging
+the paging structures do not function when PTI is enabled. It can be
+enabled by setting CONFIG_PAGE_TABLE_ISOLATION=y at compile time.
+Once enabled at compile-time, it can be disabled at boot with the
+'nopti' or 'pti=' kernel parameters (see kernel-parameters.txt).
+
+Page Table Management
+=====================
+
+When PTI is enabled, the kernel manages two sets of page tables.
+The first set is very similar to the single set which is present in
+kernels without PTI. This includes a complete mapping of userspace
+that the kernel can use for things like copy_to_user().
+
+Although _complete_, the user portion of the kernel page tables is
+crippled by setting the NX bit in the top level. This ensures
+that any missed kernel->user CR3 switch will immediately crash
+userspace upon executing its first instruction.
+
+The userspace page tables map only the kernel data needed to enter
+and exit the kernel. This data is entirely contained in the 'struct
+cpu_entry_area' structure which is placed in the fixmap which gives
+each CPU's copy of the area a compile-time-fixed virtual address.
+
+For new userspace mappings, the kernel makes the entries in its
+page tables like normal. The only difference is when the kernel
+makes entries in the top (PGD) level. In addition to setting the
+entry in the main kernel PGD, a copy of the entry is made in the
+userspace page tables' PGD.
+
+This sharing at the PGD level also inherently shares all the lower
+layers of the page tables. This leaves a single, shared set of
+userspace page tables to manage. One PTE to lock, one set of
+accessed bits, dirty bits, etc...
+
+Overhead
+========
+
+Protection against side-channel attacks is important. But,
+this protection comes at a cost:
+
+1. Increased Memory Use
+
+ a. Each process now needs an order-1 PGD instead of order-0.
+ (Consumes an additional 4k per process).
+ b. The 'cpu_entry_area' structure must be 2MB in size and 2MB
+ aligned so that it can be mapped by setting a single PMD
+ entry. This consumes nearly 2MB of RAM once the kernel
+ is decompressed, but no space in the kernel image itself.
+
+2. Runtime Cost
+
+ a. CR3 manipulation to switch between the page table copies
+ must be done at interrupt, syscall, and exception entry
+ and exit (it can be skipped when the kernel is interrupted,
+ though.) Moves to CR3 are on the order of a hundred
+ cycles, and are required at every entry and exit.
+ b. A "trampoline" must be used for SYSCALL entry. This
+ trampoline depends on a smaller set of resources than the
+ non-PTI SYSCALL entry code, so requires mapping fewer
+ things into the userspace page tables. The downside is
+ that stacks must be switched at entry time.
+ c. Global pages are disabled for all kernel structures not
+ mapped into both kernel and userspace page tables. This
+ feature of the MMU allows different processes to share TLB
+ entries mapping the kernel. Losing the feature means more
+ TLB misses after a context switch. The actual loss of
+ performance is very small, however, never exceeding 1%.
+ d. Process Context IDentifiers (PCID) is a CPU feature that
+ allows us to skip flushing the entire TLB when switching page
+ tables by setting a special bit in CR3 when the page tables
+ are changed. This makes switching the page tables (at context
+ switch, or kernel entry/exit) cheaper. But, on systems with
+ PCID support, the context switch code must flush both the user
+ and kernel entries out of the TLB. The user PCID TLB flush is
+ deferred until the exit to userspace, minimizing the cost.
+ See intel.com/sdm for the gory PCID/INVPCID details.
+ e. The userspace page tables must be populated for each new
+ process. Even without PTI, the shared kernel mappings
+ are created by copying top-level (PGD) entries into each
+ new process. But, with PTI, there are now *two* kernel
+ mappings: one in the kernel page tables that maps everything
+ and one for the entry/exit structures. At fork(), we need to
+ copy both.
+ f. In addition to the fork()-time copying, there must also
+ be an update to the userspace PGD any time a set_pgd() is done
+ on a PGD used to map userspace. This ensures that the kernel
+ and userspace copies always map the same userspace
+ memory.
+ g. On systems without PCID support, each CR3 write flushes
+ the entire TLB. That means that each syscall, interrupt
+ or exception flushes the TLB.
+ h. INVPCID is a TLB-flushing instruction which allows flushing
+ of TLB entries for non-current PCIDs. Some systems support
+ PCIDs, but do not support INVPCID. On these systems, addresses
+ can only be flushed from the TLB for the current PCID. When
+ flushing a kernel address, we need to flush all PCIDs, so a
+ single kernel address flush will require a TLB-flushing CR3
+ write upon the next use of every PCID.
+
+Possible Future Work
+====================
+1. We can be more careful about not actually writing to CR3
+ unless its value is actually changed.
+2. Allow PTI to be enabled/disabled at runtime in addition to the
+ boot-time switching.
+
+Testing
+========
+
+To test stability of PTI, the following test procedure is recommended,
+ideally doing all of these in parallel:
+
+1. Set CONFIG_DEBUG_ENTRY=y
+2. Run several copies of all of the tools/testing/selftests/x86/ tests
+ (excluding MPX and protection_keys) in a loop on multiple CPUs for
+ several minutes. These tests frequently uncover corner cases in the
+ kernel entry code. In general, old kernels might cause these tests
+ themselves to crash, but they should never crash the kernel.
+3. Run the 'perf' tool in a mode (top or record) that generates many
+ frequent performance monitoring non-maskable interrupts (see "NMI"
+ in /proc/interrupts). This exercises the NMI entry/exit code which
+ is known to trigger bugs in code paths that did not expect to be
+ interrupted, including nested NMIs. Using "-c" boosts the rate of
+ NMIs, and using two -c with separate counters encourages nested NMIs
+ and less deterministic behavior.
+ ::
+
+ while true; do perf record -c 10000 -e instructions,cycles -a sleep 10; done
+
+4. Launch a KVM virtual machine.
+5. Run 32-bit binaries on systems supporting the SYSCALL instruction.
+ This has been a lightly-tested code path and needs extra scrutiny.
+
+Debugging
+=========
+
+Bugs in PTI cause a few different signatures of crashes
+that are worth noting here.
+
+ * Failures of the selftests/x86 code. Usually a bug in one of the
+ more obscure corners of entry_64.S
+ * Crashes in early boot, especially around CPU bringup. Bugs
+ in the trampoline code or mappings cause these.
+ * Crashes at the first interrupt. Caused by bugs in entry_64.S,
+ like screwing up a page table switch. Also caused by
+ incorrectly mapping the IRQ handler entry code.
+ * Crashes at the first NMI. The NMI code is separate from main
+ interrupt handlers and can have bugs that do not affect
+ normal interrupts. Also caused by incorrectly mapping NMI
+ code. NMIs that interrupt the entry code must be very
+ careful and can be the cause of crashes that show up when
+ running perf.
+ * Kernel crashes at the first exit to userspace. entry_64.S
+ bugs, or failing to map some of the exit code.
+ * Crashes at first interrupt that interrupts userspace. The paths
+ in entry_64.S that return to userspace are sometimes separate
+ from the ones that return to the kernel.
+ * Double faults: overflowing the kernel stack because of page
+ faults upon page faults. Caused by touching non-pti-mapped
+ data in the entry code, or forgetting to switch to kernel
+ CR3 before calling into C functions which are not pti-mapped.
+ * Userspace segfaults early in boot, sometimes manifesting
+ as mount(8) failing to mount the rootfs. These have
+ tended to be TLB invalidation issues. Usually invalidating
+ the wrong PCID, or otherwise missing an invalidation.
+
+.. [1] https://gruss.cc/files/kaiser.pdf
+.. [2] https://meltdownattack.com/meltdown.pdf
diff --git a/Documentation/x86/resctrl.rst b/Documentation/x86/resctrl.rst
new file mode 100644
index 000000000..71a531061
--- /dev/null
+++ b/Documentation/x86/resctrl.rst
@@ -0,0 +1,1304 @@
+.. SPDX-License-Identifier: GPL-2.0
+.. include:: <isonum.txt>
+
+===========================================
+User Interface for Resource Control feature
+===========================================
+
+:Copyright: |copy| 2016 Intel Corporation
+:Authors: - Fenghua Yu <fenghua.yu@intel.com>
+ - Tony Luck <tony.luck@intel.com>
+ - Vikas Shivappa <vikas.shivappa@intel.com>
+
+
+Intel refers to this feature as Intel Resource Director Technology(Intel(R) RDT).
+AMD refers to this feature as AMD Platform Quality of Service(AMD QoS).
+
+This feature is enabled by the CONFIG_X86_CPU_RESCTRL and the x86 /proc/cpuinfo
+flag bits:
+
+============================================= ================================
+RDT (Resource Director Technology) Allocation "rdt_a"
+CAT (Cache Allocation Technology) "cat_l3", "cat_l2"
+CDP (Code and Data Prioritization) "cdp_l3", "cdp_l2"
+CQM (Cache QoS Monitoring) "cqm_llc", "cqm_occup_llc"
+MBM (Memory Bandwidth Monitoring) "cqm_mbm_total", "cqm_mbm_local"
+MBA (Memory Bandwidth Allocation) "mba"
+============================================= ================================
+
+To use the feature mount the file system::
+
+ # mount -t resctrl resctrl [-o cdp[,cdpl2][,mba_MBps]] /sys/fs/resctrl
+
+mount options are:
+
+"cdp":
+ Enable code/data prioritization in L3 cache allocations.
+"cdpl2":
+ Enable code/data prioritization in L2 cache allocations.
+"mba_MBps":
+ Enable the MBA Software Controller(mba_sc) to specify MBA
+ bandwidth in MBps
+
+L2 and L3 CDP are controlled separately.
+
+RDT features are orthogonal. A particular system may support only
+monitoring, only control, or both monitoring and control. Cache
+pseudo-locking is a unique way of using cache control to "pin" or
+"lock" data in the cache. Details can be found in
+"Cache Pseudo-Locking".
+
+
+The mount succeeds if either of allocation or monitoring is present, but
+only those files and directories supported by the system will be created.
+For more details on the behavior of the interface during monitoring
+and allocation, see the "Resource alloc and monitor groups" section.
+
+Info directory
+==============
+
+The 'info' directory contains information about the enabled
+resources. Each resource has its own subdirectory. The subdirectory
+names reflect the resource names.
+
+Each subdirectory contains the following files with respect to
+allocation:
+
+Cache resource(L3/L2) subdirectory contains the following files
+related to allocation:
+
+"num_closids":
+ The number of CLOSIDs which are valid for this
+ resource. The kernel uses the smallest number of
+ CLOSIDs of all enabled resources as limit.
+"cbm_mask":
+ The bitmask which is valid for this resource.
+ This mask is equivalent to 100%.
+"min_cbm_bits":
+ The minimum number of consecutive bits which
+ must be set when writing a mask.
+
+"shareable_bits":
+ Bitmask of shareable resource with other executing
+ entities (e.g. I/O). User can use this when
+ setting up exclusive cache partitions. Note that
+ some platforms support devices that have their
+ own settings for cache use which can over-ride
+ these bits.
+"bit_usage":
+ Annotated capacity bitmasks showing how all
+ instances of the resource are used. The legend is:
+
+ "0":
+ Corresponding region is unused. When the system's
+ resources have been allocated and a "0" is found
+ in "bit_usage" it is a sign that resources are
+ wasted.
+
+ "H":
+ Corresponding region is used by hardware only
+ but available for software use. If a resource
+ has bits set in "shareable_bits" but not all
+ of these bits appear in the resource groups'
+ schematas then the bits appearing in
+ "shareable_bits" but no resource group will
+ be marked as "H".
+ "X":
+ Corresponding region is available for sharing and
+ used by hardware and software. These are the
+ bits that appear in "shareable_bits" as
+ well as a resource group's allocation.
+ "S":
+ Corresponding region is used by software
+ and available for sharing.
+ "E":
+ Corresponding region is used exclusively by
+ one resource group. No sharing allowed.
+ "P":
+ Corresponding region is pseudo-locked. No
+ sharing allowed.
+
+Memory bandwidth(MB) subdirectory contains the following files
+with respect to allocation:
+
+"min_bandwidth":
+ The minimum memory bandwidth percentage which
+ user can request.
+
+"bandwidth_gran":
+ The granularity in which the memory bandwidth
+ percentage is allocated. The allocated
+ b/w percentage is rounded off to the next
+ control step available on the hardware. The
+ available bandwidth control steps are:
+ min_bandwidth + N * bandwidth_gran.
+
+"delay_linear":
+ Indicates if the delay scale is linear or
+ non-linear. This field is purely informational
+ only.
+
+"thread_throttle_mode":
+ Indicator on Intel systems of how tasks running on threads
+ of a physical core are throttled in cases where they
+ request different memory bandwidth percentages:
+
+ "max":
+ the smallest percentage is applied
+ to all threads
+ "per-thread":
+ bandwidth percentages are directly applied to
+ the threads running on the core
+
+If RDT monitoring is available there will be an "L3_MON" directory
+with the following files:
+
+"num_rmids":
+ The number of RMIDs available. This is the
+ upper bound for how many "CTRL_MON" + "MON"
+ groups can be created.
+
+"mon_features":
+ Lists the monitoring events if
+ monitoring is enabled for the resource.
+
+"max_threshold_occupancy":
+ Read/write file provides the largest value (in
+ bytes) at which a previously used LLC_occupancy
+ counter can be considered for re-use.
+
+Finally, in the top level of the "info" directory there is a file
+named "last_cmd_status". This is reset with every "command" issued
+via the file system (making new directories or writing to any of the
+control files). If the command was successful, it will read as "ok".
+If the command failed, it will provide more information that can be
+conveyed in the error returns from file operations. E.g.
+::
+
+ # echo L3:0=f7 > schemata
+ bash: echo: write error: Invalid argument
+ # cat info/last_cmd_status
+ mask f7 has non-consecutive 1-bits
+
+Resource alloc and monitor groups
+=================================
+
+Resource groups are represented as directories in the resctrl file
+system. The default group is the root directory which, immediately
+after mounting, owns all the tasks and cpus in the system and can make
+full use of all resources.
+
+On a system with RDT control features additional directories can be
+created in the root directory that specify different amounts of each
+resource (see "schemata" below). The root and these additional top level
+directories are referred to as "CTRL_MON" groups below.
+
+On a system with RDT monitoring the root directory and other top level
+directories contain a directory named "mon_groups" in which additional
+directories can be created to monitor subsets of tasks in the CTRL_MON
+group that is their ancestor. These are called "MON" groups in the rest
+of this document.
+
+Removing a directory will move all tasks and cpus owned by the group it
+represents to the parent. Removing one of the created CTRL_MON groups
+will automatically remove all MON groups below it.
+
+All groups contain the following files:
+
+"tasks":
+ Reading this file shows the list of all tasks that belong to
+ this group. Writing a task id to the file will add a task to the
+ group. If the group is a CTRL_MON group the task is removed from
+ whichever previous CTRL_MON group owned the task and also from
+ any MON group that owned the task. If the group is a MON group,
+ then the task must already belong to the CTRL_MON parent of this
+ group. The task is removed from any previous MON group.
+
+
+"cpus":
+ Reading this file shows a bitmask of the logical CPUs owned by
+ this group. Writing a mask to this file will add and remove
+ CPUs to/from this group. As with the tasks file a hierarchy is
+ maintained where MON groups may only include CPUs owned by the
+ parent CTRL_MON group.
+ When the resource group is in pseudo-locked mode this file will
+ only be readable, reflecting the CPUs associated with the
+ pseudo-locked region.
+
+
+"cpus_list":
+ Just like "cpus", only using ranges of CPUs instead of bitmasks.
+
+
+When control is enabled all CTRL_MON groups will also contain:
+
+"schemata":
+ A list of all the resources available to this group.
+ Each resource has its own line and format - see below for details.
+
+"size":
+ Mirrors the display of the "schemata" file to display the size in
+ bytes of each allocation instead of the bits representing the
+ allocation.
+
+"mode":
+ The "mode" of the resource group dictates the sharing of its
+ allocations. A "shareable" resource group allows sharing of its
+ allocations while an "exclusive" resource group does not. A
+ cache pseudo-locked region is created by first writing
+ "pseudo-locksetup" to the "mode" file before writing the cache
+ pseudo-locked region's schemata to the resource group's "schemata"
+ file. On successful pseudo-locked region creation the mode will
+ automatically change to "pseudo-locked".
+
+When monitoring is enabled all MON groups will also contain:
+
+"mon_data":
+ This contains a set of files organized by L3 domain and by
+ RDT event. E.g. on a system with two L3 domains there will
+ be subdirectories "mon_L3_00" and "mon_L3_01". Each of these
+ directories have one file per event (e.g. "llc_occupancy",
+ "mbm_total_bytes", and "mbm_local_bytes"). In a MON group these
+ files provide a read out of the current value of the event for
+ all tasks in the group. In CTRL_MON groups these files provide
+ the sum for all tasks in the CTRL_MON group and all tasks in
+ MON groups. Please see example section for more details on usage.
+
+Resource allocation rules
+-------------------------
+
+When a task is running the following rules define which resources are
+available to it:
+
+1) If the task is a member of a non-default group, then the schemata
+ for that group is used.
+
+2) Else if the task belongs to the default group, but is running on a
+ CPU that is assigned to some specific group, then the schemata for the
+ CPU's group is used.
+
+3) Otherwise the schemata for the default group is used.
+
+Resource monitoring rules
+-------------------------
+1) If a task is a member of a MON group, or non-default CTRL_MON group
+ then RDT events for the task will be reported in that group.
+
+2) If a task is a member of the default CTRL_MON group, but is running
+ on a CPU that is assigned to some specific group, then the RDT events
+ for the task will be reported in that group.
+
+3) Otherwise RDT events for the task will be reported in the root level
+ "mon_data" group.
+
+
+Notes on cache occupancy monitoring and control
+===============================================
+When moving a task from one group to another you should remember that
+this only affects *new* cache allocations by the task. E.g. you may have
+a task in a monitor group showing 3 MB of cache occupancy. If you move
+to a new group and immediately check the occupancy of the old and new
+groups you will likely see that the old group is still showing 3 MB and
+the new group zero. When the task accesses locations still in cache from
+before the move, the h/w does not update any counters. On a busy system
+you will likely see the occupancy in the old group go down as cache lines
+are evicted and re-used while the occupancy in the new group rises as
+the task accesses memory and loads into the cache are counted based on
+membership in the new group.
+
+The same applies to cache allocation control. Moving a task to a group
+with a smaller cache partition will not evict any cache lines. The
+process may continue to use them from the old partition.
+
+Hardware uses CLOSid(Class of service ID) and an RMID(Resource monitoring ID)
+to identify a control group and a monitoring group respectively. Each of
+the resource groups are mapped to these IDs based on the kind of group. The
+number of CLOSid and RMID are limited by the hardware and hence the creation of
+a "CTRL_MON" directory may fail if we run out of either CLOSID or RMID
+and creation of "MON" group may fail if we run out of RMIDs.
+
+max_threshold_occupancy - generic concepts
+------------------------------------------
+
+Note that an RMID once freed may not be immediately available for use as
+the RMID is still tagged the cache lines of the previous user of RMID.
+Hence such RMIDs are placed on limbo list and checked back if the cache
+occupancy has gone down. If there is a time when system has a lot of
+limbo RMIDs but which are not ready to be used, user may see an -EBUSY
+during mkdir.
+
+max_threshold_occupancy is a user configurable value to determine the
+occupancy at which an RMID can be freed.
+
+Schemata files - general concepts
+---------------------------------
+Each line in the file describes one resource. The line starts with
+the name of the resource, followed by specific values to be applied
+in each of the instances of that resource on the system.
+
+Cache IDs
+---------
+On current generation systems there is one L3 cache per socket and L2
+caches are generally just shared by the hyperthreads on a core, but this
+isn't an architectural requirement. We could have multiple separate L3
+caches on a socket, multiple cores could share an L2 cache. So instead
+of using "socket" or "core" to define the set of logical cpus sharing
+a resource we use a "Cache ID". At a given cache level this will be a
+unique number across the whole system (but it isn't guaranteed to be a
+contiguous sequence, there may be gaps). To find the ID for each logical
+CPU look in /sys/devices/system/cpu/cpu*/cache/index*/id
+
+Cache Bit Masks (CBM)
+---------------------
+For cache resources we describe the portion of the cache that is available
+for allocation using a bitmask. The maximum value of the mask is defined
+by each cpu model (and may be different for different cache levels). It
+is found using CPUID, but is also provided in the "info" directory of
+the resctrl file system in "info/{resource}/cbm_mask". Intel hardware
+requires that these masks have all the '1' bits in a contiguous block. So
+0x3, 0x6 and 0xC are legal 4-bit masks with two bits set, but 0x5, 0x9
+and 0xA are not. On a system with a 20-bit mask each bit represents 5%
+of the capacity of the cache. You could partition the cache into four
+equal parts with masks: 0x1f, 0x3e0, 0x7c00, 0xf8000.
+
+Memory bandwidth Allocation and monitoring
+==========================================
+
+For Memory bandwidth resource, by default the user controls the resource
+by indicating the percentage of total memory bandwidth.
+
+The minimum bandwidth percentage value for each cpu model is predefined
+and can be looked up through "info/MB/min_bandwidth". The bandwidth
+granularity that is allocated is also dependent on the cpu model and can
+be looked up at "info/MB/bandwidth_gran". The available bandwidth
+control steps are: min_bw + N * bw_gran. Intermediate values are rounded
+to the next control step available on the hardware.
+
+The bandwidth throttling is a core specific mechanism on some of Intel
+SKUs. Using a high bandwidth and a low bandwidth setting on two threads
+sharing a core may result in both threads being throttled to use the
+low bandwidth (see "thread_throttle_mode").
+
+The fact that Memory bandwidth allocation(MBA) may be a core
+specific mechanism where as memory bandwidth monitoring(MBM) is done at
+the package level may lead to confusion when users try to apply control
+via the MBA and then monitor the bandwidth to see if the controls are
+effective. Below are such scenarios:
+
+1. User may *not* see increase in actual bandwidth when percentage
+ values are increased:
+
+This can occur when aggregate L2 external bandwidth is more than L3
+external bandwidth. Consider an SKL SKU with 24 cores on a package and
+where L2 external is 10GBps (hence aggregate L2 external bandwidth is
+240GBps) and L3 external bandwidth is 100GBps. Now a workload with '20
+threads, having 50% bandwidth, each consuming 5GBps' consumes the max L3
+bandwidth of 100GBps although the percentage value specified is only 50%
+<< 100%. Hence increasing the bandwidth percentage will not yield any
+more bandwidth. This is because although the L2 external bandwidth still
+has capacity, the L3 external bandwidth is fully used. Also note that
+this would be dependent on number of cores the benchmark is run on.
+
+2. Same bandwidth percentage may mean different actual bandwidth
+ depending on # of threads:
+
+For the same SKU in #1, a 'single thread, with 10% bandwidth' and '4
+thread, with 10% bandwidth' can consume upto 10GBps and 40GBps although
+they have same percentage bandwidth of 10%. This is simply because as
+threads start using more cores in an rdtgroup, the actual bandwidth may
+increase or vary although user specified bandwidth percentage is same.
+
+In order to mitigate this and make the interface more user friendly,
+resctrl added support for specifying the bandwidth in MBps as well. The
+kernel underneath would use a software feedback mechanism or a "Software
+Controller(mba_sc)" which reads the actual bandwidth using MBM counters
+and adjust the memory bandwidth percentages to ensure::
+
+ "actual bandwidth < user specified bandwidth".
+
+By default, the schemata would take the bandwidth percentage values
+where as user can switch to the "MBA software controller" mode using
+a mount option 'mba_MBps'. The schemata format is specified in the below
+sections.
+
+L3 schemata file details (code and data prioritization disabled)
+----------------------------------------------------------------
+With CDP disabled the L3 schemata format is::
+
+ L3:<cache_id0>=<cbm>;<cache_id1>=<cbm>;...
+
+L3 schemata file details (CDP enabled via mount option to resctrl)
+------------------------------------------------------------------
+When CDP is enabled L3 control is split into two separate resources
+so you can specify independent masks for code and data like this::
+
+ L3DATA:<cache_id0>=<cbm>;<cache_id1>=<cbm>;...
+ L3CODE:<cache_id0>=<cbm>;<cache_id1>=<cbm>;...
+
+L2 schemata file details
+------------------------
+CDP is supported at L2 using the 'cdpl2' mount option. The schemata
+format is either::
+
+ L2:<cache_id0>=<cbm>;<cache_id1>=<cbm>;...
+
+or
+
+ L2DATA:<cache_id0>=<cbm>;<cache_id1>=<cbm>;...
+ L2CODE:<cache_id0>=<cbm>;<cache_id1>=<cbm>;...
+
+
+Memory bandwidth Allocation (default mode)
+------------------------------------------
+
+Memory b/w domain is L3 cache.
+::
+
+ MB:<cache_id0>=bandwidth0;<cache_id1>=bandwidth1;...
+
+Memory bandwidth Allocation specified in MBps
+---------------------------------------------
+
+Memory bandwidth domain is L3 cache.
+::
+
+ MB:<cache_id0>=bw_MBps0;<cache_id1>=bw_MBps1;...
+
+Reading/writing the schemata file
+---------------------------------
+Reading the schemata file will show the state of all resources
+on all domains. When writing you only need to specify those values
+which you wish to change. E.g.
+::
+
+ # cat schemata
+ L3DATA:0=fffff;1=fffff;2=fffff;3=fffff
+ L3CODE:0=fffff;1=fffff;2=fffff;3=fffff
+ # echo "L3DATA:2=3c0;" > schemata
+ # cat schemata
+ L3DATA:0=fffff;1=fffff;2=3c0;3=fffff
+ L3CODE:0=fffff;1=fffff;2=fffff;3=fffff
+
+Cache Pseudo-Locking
+====================
+CAT enables a user to specify the amount of cache space that an
+application can fill. Cache pseudo-locking builds on the fact that a
+CPU can still read and write data pre-allocated outside its current
+allocated area on a cache hit. With cache pseudo-locking, data can be
+preloaded into a reserved portion of cache that no application can
+fill, and from that point on will only serve cache hits. The cache
+pseudo-locked memory is made accessible to user space where an
+application can map it into its virtual address space and thus have
+a region of memory with reduced average read latency.
+
+The creation of a cache pseudo-locked region is triggered by a request
+from the user to do so that is accompanied by a schemata of the region
+to be pseudo-locked. The cache pseudo-locked region is created as follows:
+
+- Create a CAT allocation CLOSNEW with a CBM matching the schemata
+ from the user of the cache region that will contain the pseudo-locked
+ memory. This region must not overlap with any current CAT allocation/CLOS
+ on the system and no future overlap with this cache region is allowed
+ while the pseudo-locked region exists.
+- Create a contiguous region of memory of the same size as the cache
+ region.
+- Flush the cache, disable hardware prefetchers, disable preemption.
+- Make CLOSNEW the active CLOS and touch the allocated memory to load
+ it into the cache.
+- Set the previous CLOS as active.
+- At this point the closid CLOSNEW can be released - the cache
+ pseudo-locked region is protected as long as its CBM does not appear in
+ any CAT allocation. Even though the cache pseudo-locked region will from
+ this point on not appear in any CBM of any CLOS an application running with
+ any CLOS will be able to access the memory in the pseudo-locked region since
+ the region continues to serve cache hits.
+- The contiguous region of memory loaded into the cache is exposed to
+ user-space as a character device.
+
+Cache pseudo-locking increases the probability that data will remain
+in the cache via carefully configuring the CAT feature and controlling
+application behavior. There is no guarantee that data is placed in
+cache. Instructions like INVD, WBINVD, CLFLUSH, etc. can still evict
+“locked” data from cache. Power management C-states may shrink or
+power off cache. Deeper C-states will automatically be restricted on
+pseudo-locked region creation.
+
+It is required that an application using a pseudo-locked region runs
+with affinity to the cores (or a subset of the cores) associated
+with the cache on which the pseudo-locked region resides. A sanity check
+within the code will not allow an application to map pseudo-locked memory
+unless it runs with affinity to cores associated with the cache on which the
+pseudo-locked region resides. The sanity check is only done during the
+initial mmap() handling, there is no enforcement afterwards and the
+application self needs to ensure it remains affine to the correct cores.
+
+Pseudo-locking is accomplished in two stages:
+
+1) During the first stage the system administrator allocates a portion
+ of cache that should be dedicated to pseudo-locking. At this time an
+ equivalent portion of memory is allocated, loaded into allocated
+ cache portion, and exposed as a character device.
+2) During the second stage a user-space application maps (mmap()) the
+ pseudo-locked memory into its address space.
+
+Cache Pseudo-Locking Interface
+------------------------------
+A pseudo-locked region is created using the resctrl interface as follows:
+
+1) Create a new resource group by creating a new directory in /sys/fs/resctrl.
+2) Change the new resource group's mode to "pseudo-locksetup" by writing
+ "pseudo-locksetup" to the "mode" file.
+3) Write the schemata of the pseudo-locked region to the "schemata" file. All
+ bits within the schemata should be "unused" according to the "bit_usage"
+ file.
+
+On successful pseudo-locked region creation the "mode" file will contain
+"pseudo-locked" and a new character device with the same name as the resource
+group will exist in /dev/pseudo_lock. This character device can be mmap()'ed
+by user space in order to obtain access to the pseudo-locked memory region.
+
+An example of cache pseudo-locked region creation and usage can be found below.
+
+Cache Pseudo-Locking Debugging Interface
+----------------------------------------
+The pseudo-locking debugging interface is enabled by default (if
+CONFIG_DEBUG_FS is enabled) and can be found in /sys/kernel/debug/resctrl.
+
+There is no explicit way for the kernel to test if a provided memory
+location is present in the cache. The pseudo-locking debugging interface uses
+the tracing infrastructure to provide two ways to measure cache residency of
+the pseudo-locked region:
+
+1) Memory access latency using the pseudo_lock_mem_latency tracepoint. Data
+ from these measurements are best visualized using a hist trigger (see
+ example below). In this test the pseudo-locked region is traversed at
+ a stride of 32 bytes while hardware prefetchers and preemption
+ are disabled. This also provides a substitute visualization of cache
+ hits and misses.
+2) Cache hit and miss measurements using model specific precision counters if
+ available. Depending on the levels of cache on the system the pseudo_lock_l2
+ and pseudo_lock_l3 tracepoints are available.
+
+When a pseudo-locked region is created a new debugfs directory is created for
+it in debugfs as /sys/kernel/debug/resctrl/<newdir>. A single
+write-only file, pseudo_lock_measure, is present in this directory. The
+measurement of the pseudo-locked region depends on the number written to this
+debugfs file:
+
+1:
+ writing "1" to the pseudo_lock_measure file will trigger the latency
+ measurement captured in the pseudo_lock_mem_latency tracepoint. See
+ example below.
+2:
+ writing "2" to the pseudo_lock_measure file will trigger the L2 cache
+ residency (cache hits and misses) measurement captured in the
+ pseudo_lock_l2 tracepoint. See example below.
+3:
+ writing "3" to the pseudo_lock_measure file will trigger the L3 cache
+ residency (cache hits and misses) measurement captured in the
+ pseudo_lock_l3 tracepoint.
+
+All measurements are recorded with the tracing infrastructure. This requires
+the relevant tracepoints to be enabled before the measurement is triggered.
+
+Example of latency debugging interface
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+In this example a pseudo-locked region named "newlock" was created. Here is
+how we can measure the latency in cycles of reading from this region and
+visualize this data with a histogram that is available if CONFIG_HIST_TRIGGERS
+is set::
+
+ # :> /sys/kernel/debug/tracing/trace
+ # echo 'hist:keys=latency' > /sys/kernel/debug/tracing/events/resctrl/pseudo_lock_mem_latency/trigger
+ # echo 1 > /sys/kernel/debug/tracing/events/resctrl/pseudo_lock_mem_latency/enable
+ # echo 1 > /sys/kernel/debug/resctrl/newlock/pseudo_lock_measure
+ # echo 0 > /sys/kernel/debug/tracing/events/resctrl/pseudo_lock_mem_latency/enable
+ # cat /sys/kernel/debug/tracing/events/resctrl/pseudo_lock_mem_latency/hist
+
+ # event histogram
+ #
+ # trigger info: hist:keys=latency:vals=hitcount:sort=hitcount:size=2048 [active]
+ #
+
+ { latency: 456 } hitcount: 1
+ { latency: 50 } hitcount: 83
+ { latency: 36 } hitcount: 96
+ { latency: 44 } hitcount: 174
+ { latency: 48 } hitcount: 195
+ { latency: 46 } hitcount: 262
+ { latency: 42 } hitcount: 693
+ { latency: 40 } hitcount: 3204
+ { latency: 38 } hitcount: 3484
+
+ Totals:
+ Hits: 8192
+ Entries: 9
+ Dropped: 0
+
+Example of cache hits/misses debugging
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+In this example a pseudo-locked region named "newlock" was created on the L2
+cache of a platform. Here is how we can obtain details of the cache hits
+and misses using the platform's precision counters.
+::
+
+ # :> /sys/kernel/debug/tracing/trace
+ # echo 1 > /sys/kernel/debug/tracing/events/resctrl/pseudo_lock_l2/enable
+ # echo 2 > /sys/kernel/debug/resctrl/newlock/pseudo_lock_measure
+ # echo 0 > /sys/kernel/debug/tracing/events/resctrl/pseudo_lock_l2/enable
+ # cat /sys/kernel/debug/tracing/trace
+
+ # tracer: nop
+ #
+ # _-----=> irqs-off
+ # / _----=> need-resched
+ # | / _---=> hardirq/softirq
+ # || / _--=> preempt-depth
+ # ||| / delay
+ # TASK-PID CPU# |||| TIMESTAMP FUNCTION
+ # | | | |||| | |
+ pseudo_lock_mea-1672 [002] .... 3132.860500: pseudo_lock_l2: hits=4097 miss=0
+
+
+Examples for RDT allocation usage
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+1) Example 1
+
+On a two socket machine (one L3 cache per socket) with just four bits
+for cache bit masks, minimum b/w of 10% with a memory bandwidth
+granularity of 10%.
+::
+
+ # mount -t resctrl resctrl /sys/fs/resctrl
+ # cd /sys/fs/resctrl
+ # mkdir p0 p1
+ # echo "L3:0=3;1=c\nMB:0=50;1=50" > /sys/fs/resctrl/p0/schemata
+ # echo "L3:0=3;1=3\nMB:0=50;1=50" > /sys/fs/resctrl/p1/schemata
+
+The default resource group is unmodified, so we have access to all parts
+of all caches (its schemata file reads "L3:0=f;1=f").
+
+Tasks that are under the control of group "p0" may only allocate from the
+"lower" 50% on cache ID 0, and the "upper" 50% of cache ID 1.
+Tasks in group "p1" use the "lower" 50% of cache on both sockets.
+
+Similarly, tasks that are under the control of group "p0" may use a
+maximum memory b/w of 50% on socket0 and 50% on socket 1.
+Tasks in group "p1" may also use 50% memory b/w on both sockets.
+Note that unlike cache masks, memory b/w cannot specify whether these
+allocations can overlap or not. The allocations specifies the maximum
+b/w that the group may be able to use and the system admin can configure
+the b/w accordingly.
+
+If resctrl is using the software controller (mba_sc) then user can enter the
+max b/w in MB rather than the percentage values.
+::
+
+ # echo "L3:0=3;1=c\nMB:0=1024;1=500" > /sys/fs/resctrl/p0/schemata
+ # echo "L3:0=3;1=3\nMB:0=1024;1=500" > /sys/fs/resctrl/p1/schemata
+
+In the above example the tasks in "p1" and "p0" on socket 0 would use a max b/w
+of 1024MB where as on socket 1 they would use 500MB.
+
+2) Example 2
+
+Again two sockets, but this time with a more realistic 20-bit mask.
+
+Two real time tasks pid=1234 running on processor 0 and pid=5678 running on
+processor 1 on socket 0 on a 2-socket and dual core machine. To avoid noisy
+neighbors, each of the two real-time tasks exclusively occupies one quarter
+of L3 cache on socket 0.
+::
+
+ # mount -t resctrl resctrl /sys/fs/resctrl
+ # cd /sys/fs/resctrl
+
+First we reset the schemata for the default group so that the "upper"
+50% of the L3 cache on socket 0 and 50% of memory b/w cannot be used by
+ordinary tasks::
+
+ # echo "L3:0=3ff;1=fffff\nMB:0=50;1=100" > schemata
+
+Next we make a resource group for our first real time task and give
+it access to the "top" 25% of the cache on socket 0.
+::
+
+ # mkdir p0
+ # echo "L3:0=f8000;1=fffff" > p0/schemata
+
+Finally we move our first real time task into this resource group. We
+also use taskset(1) to ensure the task always runs on a dedicated CPU
+on socket 0. Most uses of resource groups will also constrain which
+processors tasks run on.
+::
+
+ # echo 1234 > p0/tasks
+ # taskset -cp 1 1234
+
+Ditto for the second real time task (with the remaining 25% of cache)::
+
+ # mkdir p1
+ # echo "L3:0=7c00;1=fffff" > p1/schemata
+ # echo 5678 > p1/tasks
+ # taskset -cp 2 5678
+
+For the same 2 socket system with memory b/w resource and CAT L3 the
+schemata would look like(Assume min_bandwidth 10 and bandwidth_gran is
+10):
+
+For our first real time task this would request 20% memory b/w on socket 0.
+::
+
+ # echo -e "L3:0=f8000;1=fffff\nMB:0=20;1=100" > p0/schemata
+
+For our second real time task this would request an other 20% memory b/w
+on socket 0.
+::
+
+ # echo -e "L3:0=f8000;1=fffff\nMB:0=20;1=100" > p0/schemata
+
+3) Example 3
+
+A single socket system which has real-time tasks running on core 4-7 and
+non real-time workload assigned to core 0-3. The real-time tasks share text
+and data, so a per task association is not required and due to interaction
+with the kernel it's desired that the kernel on these cores shares L3 with
+the tasks.
+::
+
+ # mount -t resctrl resctrl /sys/fs/resctrl
+ # cd /sys/fs/resctrl
+
+First we reset the schemata for the default group so that the "upper"
+50% of the L3 cache on socket 0, and 50% of memory bandwidth on socket 0
+cannot be used by ordinary tasks::
+
+ # echo "L3:0=3ff\nMB:0=50" > schemata
+
+Next we make a resource group for our real time cores and give it access
+to the "top" 50% of the cache on socket 0 and 50% of memory bandwidth on
+socket 0.
+::
+
+ # mkdir p0
+ # echo "L3:0=ffc00\nMB:0=50" > p0/schemata
+
+Finally we move core 4-7 over to the new group and make sure that the
+kernel and the tasks running there get 50% of the cache. They should
+also get 50% of memory bandwidth assuming that the cores 4-7 are SMT
+siblings and only the real time threads are scheduled on the cores 4-7.
+::
+
+ # echo F0 > p0/cpus
+
+4) Example 4
+
+The resource groups in previous examples were all in the default "shareable"
+mode allowing sharing of their cache allocations. If one resource group
+configures a cache allocation then nothing prevents another resource group
+to overlap with that allocation.
+
+In this example a new exclusive resource group will be created on a L2 CAT
+system with two L2 cache instances that can be configured with an 8-bit
+capacity bitmask. The new exclusive resource group will be configured to use
+25% of each cache instance.
+::
+
+ # mount -t resctrl resctrl /sys/fs/resctrl/
+ # cd /sys/fs/resctrl
+
+First, we observe that the default group is configured to allocate to all L2
+cache::
+
+ # cat schemata
+ L2:0=ff;1=ff
+
+We could attempt to create the new resource group at this point, but it will
+fail because of the overlap with the schemata of the default group::
+
+ # mkdir p0
+ # echo 'L2:0=0x3;1=0x3' > p0/schemata
+ # cat p0/mode
+ shareable
+ # echo exclusive > p0/mode
+ -sh: echo: write error: Invalid argument
+ # cat info/last_cmd_status
+ schemata overlaps
+
+To ensure that there is no overlap with another resource group the default
+resource group's schemata has to change, making it possible for the new
+resource group to become exclusive.
+::
+
+ # echo 'L2:0=0xfc;1=0xfc' > schemata
+ # echo exclusive > p0/mode
+ # grep . p0/*
+ p0/cpus:0
+ p0/mode:exclusive
+ p0/schemata:L2:0=03;1=03
+ p0/size:L2:0=262144;1=262144
+
+A new resource group will on creation not overlap with an exclusive resource
+group::
+
+ # mkdir p1
+ # grep . p1/*
+ p1/cpus:0
+ p1/mode:shareable
+ p1/schemata:L2:0=fc;1=fc
+ p1/size:L2:0=786432;1=786432
+
+The bit_usage will reflect how the cache is used::
+
+ # cat info/L2/bit_usage
+ 0=SSSSSSEE;1=SSSSSSEE
+
+A resource group cannot be forced to overlap with an exclusive resource group::
+
+ # echo 'L2:0=0x1;1=0x1' > p1/schemata
+ -sh: echo: write error: Invalid argument
+ # cat info/last_cmd_status
+ overlaps with exclusive group
+
+Example of Cache Pseudo-Locking
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+Lock portion of L2 cache from cache id 1 using CBM 0x3. Pseudo-locked
+region is exposed at /dev/pseudo_lock/newlock that can be provided to
+application for argument to mmap().
+::
+
+ # mount -t resctrl resctrl /sys/fs/resctrl/
+ # cd /sys/fs/resctrl
+
+Ensure that there are bits available that can be pseudo-locked, since only
+unused bits can be pseudo-locked the bits to be pseudo-locked needs to be
+removed from the default resource group's schemata::
+
+ # cat info/L2/bit_usage
+ 0=SSSSSSSS;1=SSSSSSSS
+ # echo 'L2:1=0xfc' > schemata
+ # cat info/L2/bit_usage
+ 0=SSSSSSSS;1=SSSSSS00
+
+Create a new resource group that will be associated with the pseudo-locked
+region, indicate that it will be used for a pseudo-locked region, and
+configure the requested pseudo-locked region capacity bitmask::
+
+ # mkdir newlock
+ # echo pseudo-locksetup > newlock/mode
+ # echo 'L2:1=0x3' > newlock/schemata
+
+On success the resource group's mode will change to pseudo-locked, the
+bit_usage will reflect the pseudo-locked region, and the character device
+exposing the pseudo-locked region will exist::
+
+ # cat newlock/mode
+ pseudo-locked
+ # cat info/L2/bit_usage
+ 0=SSSSSSSS;1=SSSSSSPP
+ # ls -l /dev/pseudo_lock/newlock
+ crw------- 1 root root 243, 0 Apr 3 05:01 /dev/pseudo_lock/newlock
+
+::
+
+ /*
+ * Example code to access one page of pseudo-locked cache region
+ * from user space.
+ */
+ #define _GNU_SOURCE
+ #include <fcntl.h>
+ #include <sched.h>
+ #include <stdio.h>
+ #include <stdlib.h>
+ #include <unistd.h>
+ #include <sys/mman.h>
+
+ /*
+ * It is required that the application runs with affinity to only
+ * cores associated with the pseudo-locked region. Here the cpu
+ * is hardcoded for convenience of example.
+ */
+ static int cpuid = 2;
+
+ int main(int argc, char *argv[])
+ {
+ cpu_set_t cpuset;
+ long page_size;
+ void *mapping;
+ int dev_fd;
+ int ret;
+
+ page_size = sysconf(_SC_PAGESIZE);
+
+ CPU_ZERO(&cpuset);
+ CPU_SET(cpuid, &cpuset);
+ ret = sched_setaffinity(0, sizeof(cpuset), &cpuset);
+ if (ret < 0) {
+ perror("sched_setaffinity");
+ exit(EXIT_FAILURE);
+ }
+
+ dev_fd = open("/dev/pseudo_lock/newlock", O_RDWR);
+ if (dev_fd < 0) {
+ perror("open");
+ exit(EXIT_FAILURE);
+ }
+
+ mapping = mmap(0, page_size, PROT_READ | PROT_WRITE, MAP_SHARED,
+ dev_fd, 0);
+ if (mapping == MAP_FAILED) {
+ perror("mmap");
+ close(dev_fd);
+ exit(EXIT_FAILURE);
+ }
+
+ /* Application interacts with pseudo-locked memory @mapping */
+
+ ret = munmap(mapping, page_size);
+ if (ret < 0) {
+ perror("munmap");
+ close(dev_fd);
+ exit(EXIT_FAILURE);
+ }
+
+ close(dev_fd);
+ exit(EXIT_SUCCESS);
+ }
+
+Locking between applications
+----------------------------
+
+Certain operations on the resctrl filesystem, composed of read/writes
+to/from multiple files, must be atomic.
+
+As an example, the allocation of an exclusive reservation of L3 cache
+involves:
+
+ 1. Read the cbmmasks from each directory or the per-resource "bit_usage"
+ 2. Find a contiguous set of bits in the global CBM bitmask that is clear
+ in any of the directory cbmmasks
+ 3. Create a new directory
+ 4. Set the bits found in step 2 to the new directory "schemata" file
+
+If two applications attempt to allocate space concurrently then they can
+end up allocating the same bits so the reservations are shared instead of
+exclusive.
+
+To coordinate atomic operations on the resctrlfs and to avoid the problem
+above, the following locking procedure is recommended:
+
+Locking is based on flock, which is available in libc and also as a shell
+script command
+
+Write lock:
+
+ A) Take flock(LOCK_EX) on /sys/fs/resctrl
+ B) Read/write the directory structure.
+ C) funlock
+
+Read lock:
+
+ A) Take flock(LOCK_SH) on /sys/fs/resctrl
+ B) If success read the directory structure.
+ C) funlock
+
+Example with bash::
+
+ # Atomically read directory structure
+ $ flock -s /sys/fs/resctrl/ find /sys/fs/resctrl
+
+ # Read directory contents and create new subdirectory
+
+ $ cat create-dir.sh
+ find /sys/fs/resctrl/ > output.txt
+ mask = function-of(output.txt)
+ mkdir /sys/fs/resctrl/newres/
+ echo mask > /sys/fs/resctrl/newres/schemata
+
+ $ flock /sys/fs/resctrl/ ./create-dir.sh
+
+Example with C::
+
+ /*
+ * Example code do take advisory locks
+ * before accessing resctrl filesystem
+ */
+ #include <sys/file.h>
+ #include <stdlib.h>
+
+ void resctrl_take_shared_lock(int fd)
+ {
+ int ret;
+
+ /* take shared lock on resctrl filesystem */
+ ret = flock(fd, LOCK_SH);
+ if (ret) {
+ perror("flock");
+ exit(-1);
+ }
+ }
+
+ void resctrl_take_exclusive_lock(int fd)
+ {
+ int ret;
+
+ /* release lock on resctrl filesystem */
+ ret = flock(fd, LOCK_EX);
+ if (ret) {
+ perror("flock");
+ exit(-1);
+ }
+ }
+
+ void resctrl_release_lock(int fd)
+ {
+ int ret;
+
+ /* take shared lock on resctrl filesystem */
+ ret = flock(fd, LOCK_UN);
+ if (ret) {
+ perror("flock");
+ exit(-1);
+ }
+ }
+
+ void main(void)
+ {
+ int fd, ret;
+
+ fd = open("/sys/fs/resctrl", O_DIRECTORY);
+ if (fd == -1) {
+ perror("open");
+ exit(-1);
+ }
+ resctrl_take_shared_lock(fd);
+ /* code to read directory contents */
+ resctrl_release_lock(fd);
+
+ resctrl_take_exclusive_lock(fd);
+ /* code to read and write directory contents */
+ resctrl_release_lock(fd);
+ }
+
+Examples for RDT Monitoring along with allocation usage
+=======================================================
+Reading monitored data
+----------------------
+Reading an event file (for ex: mon_data/mon_L3_00/llc_occupancy) would
+show the current snapshot of LLC occupancy of the corresponding MON
+group or CTRL_MON group.
+
+
+Example 1 (Monitor CTRL_MON group and subset of tasks in CTRL_MON group)
+------------------------------------------------------------------------
+On a two socket machine (one L3 cache per socket) with just four bits
+for cache bit masks::
+
+ # mount -t resctrl resctrl /sys/fs/resctrl
+ # cd /sys/fs/resctrl
+ # mkdir p0 p1
+ # echo "L3:0=3;1=c" > /sys/fs/resctrl/p0/schemata
+ # echo "L3:0=3;1=3" > /sys/fs/resctrl/p1/schemata
+ # echo 5678 > p1/tasks
+ # echo 5679 > p1/tasks
+
+The default resource group is unmodified, so we have access to all parts
+of all caches (its schemata file reads "L3:0=f;1=f").
+
+Tasks that are under the control of group "p0" may only allocate from the
+"lower" 50% on cache ID 0, and the "upper" 50% of cache ID 1.
+Tasks in group "p1" use the "lower" 50% of cache on both sockets.
+
+Create monitor groups and assign a subset of tasks to each monitor group.
+::
+
+ # cd /sys/fs/resctrl/p1/mon_groups
+ # mkdir m11 m12
+ # echo 5678 > m11/tasks
+ # echo 5679 > m12/tasks
+
+fetch data (data shown in bytes)
+::
+
+ # cat m11/mon_data/mon_L3_00/llc_occupancy
+ 16234000
+ # cat m11/mon_data/mon_L3_01/llc_occupancy
+ 14789000
+ # cat m12/mon_data/mon_L3_00/llc_occupancy
+ 16789000
+
+The parent ctrl_mon group shows the aggregated data.
+::
+
+ # cat /sys/fs/resctrl/p1/mon_data/mon_l3_00/llc_occupancy
+ 31234000
+
+Example 2 (Monitor a task from its creation)
+--------------------------------------------
+On a two socket machine (one L3 cache per socket)::
+
+ # mount -t resctrl resctrl /sys/fs/resctrl
+ # cd /sys/fs/resctrl
+ # mkdir p0 p1
+
+An RMID is allocated to the group once its created and hence the <cmd>
+below is monitored from its creation.
+::
+
+ # echo $$ > /sys/fs/resctrl/p1/tasks
+ # <cmd>
+
+Fetch the data::
+
+ # cat /sys/fs/resctrl/p1/mon_data/mon_l3_00/llc_occupancy
+ 31789000
+
+Example 3 (Monitor without CAT support or before creating CAT groups)
+---------------------------------------------------------------------
+
+Assume a system like HSW has only CQM and no CAT support. In this case
+the resctrl will still mount but cannot create CTRL_MON directories.
+But user can create different MON groups within the root group thereby
+able to monitor all tasks including kernel threads.
+
+This can also be used to profile jobs cache size footprint before being
+able to allocate them to different allocation groups.
+::
+
+ # mount -t resctrl resctrl /sys/fs/resctrl
+ # cd /sys/fs/resctrl
+ # mkdir mon_groups/m01
+ # mkdir mon_groups/m02
+
+ # echo 3478 > /sys/fs/resctrl/mon_groups/m01/tasks
+ # echo 2467 > /sys/fs/resctrl/mon_groups/m02/tasks
+
+Monitor the groups separately and also get per domain data. From the
+below its apparent that the tasks are mostly doing work on
+domain(socket) 0.
+::
+
+ # cat /sys/fs/resctrl/mon_groups/m01/mon_L3_00/llc_occupancy
+ 31234000
+ # cat /sys/fs/resctrl/mon_groups/m01/mon_L3_01/llc_occupancy
+ 34555
+ # cat /sys/fs/resctrl/mon_groups/m02/mon_L3_00/llc_occupancy
+ 31234000
+ # cat /sys/fs/resctrl/mon_groups/m02/mon_L3_01/llc_occupancy
+ 32789
+
+
+Example 4 (Monitor real time tasks)
+-----------------------------------
+
+A single socket system which has real time tasks running on cores 4-7
+and non real time tasks on other cpus. We want to monitor the cache
+occupancy of the real time threads on these cores.
+::
+
+ # mount -t resctrl resctrl /sys/fs/resctrl
+ # cd /sys/fs/resctrl
+ # mkdir p1
+
+Move the cpus 4-7 over to p1::
+
+ # echo f0 > p1/cpus
+
+View the llc occupancy snapshot::
+
+ # cat /sys/fs/resctrl/p1/mon_data/mon_L3_00/llc_occupancy
+ 11234000
+
+Intel RDT Errata
+================
+
+Intel MBM Counters May Report System Memory Bandwidth Incorrectly
+-----------------------------------------------------------------
+
+Errata SKX99 for Skylake server and BDF102 for Broadwell server.
+
+Problem: Intel Memory Bandwidth Monitoring (MBM) counters track metrics
+according to the assigned Resource Monitor ID (RMID) for that logical
+core. The IA32_QM_CTR register (MSR 0xC8E), used to report these
+metrics, may report incorrect system bandwidth for certain RMID values.
+
+Implication: Due to the errata, system memory bandwidth may not match
+what is reported.
+
+Workaround: MBM total and local readings are corrected according to the
+following correction factor table:
+
++---------------+---------------+---------------+-----------------+
+|core count |rmid count |rmid threshold |correction factor|
++---------------+---------------+---------------+-----------------+
+|1 |8 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|2 |16 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|3 |24 |15 |0.969650 |
++---------------+---------------+---------------+-----------------+
+|4 |32 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|6 |48 |31 |0.969650 |
++---------------+---------------+---------------+-----------------+
+|7 |56 |47 |1.142857 |
++---------------+---------------+---------------+-----------------+
+|8 |64 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|9 |72 |63 |1.185115 |
++---------------+---------------+---------------+-----------------+
+|10 |80 |63 |1.066553 |
++---------------+---------------+---------------+-----------------+
+|11 |88 |79 |1.454545 |
++---------------+---------------+---------------+-----------------+
+|12 |96 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|13 |104 |95 |1.230769 |
++---------------+---------------+---------------+-----------------+
+|14 |112 |95 |1.142857 |
++---------------+---------------+---------------+-----------------+
+|15 |120 |95 |1.066667 |
++---------------+---------------+---------------+-----------------+
+|16 |128 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|17 |136 |127 |1.254863 |
++---------------+---------------+---------------+-----------------+
+|18 |144 |127 |1.185255 |
++---------------+---------------+---------------+-----------------+
+|19 |152 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|20 |160 |127 |1.066667 |
++---------------+---------------+---------------+-----------------+
+|21 |168 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|22 |176 |159 |1.454334 |
++---------------+---------------+---------------+-----------------+
+|23 |184 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|24 |192 |127 |0.969744 |
++---------------+---------------+---------------+-----------------+
+|25 |200 |191 |1.280246 |
++---------------+---------------+---------------+-----------------+
+|26 |208 |191 |1.230921 |
++---------------+---------------+---------------+-----------------+
+|27 |216 |0 |1.000000 |
++---------------+---------------+---------------+-----------------+
+|28 |224 |191 |1.143118 |
++---------------+---------------+---------------+-----------------+
+
+If rmid > rmid threshold, MBM total and local values should be multiplied
+by the correction factor.
+
+See:
+
+1. Erratum SKX99 in Intel Xeon Processor Scalable Family Specification Update:
+http://web.archive.org/web/20200716124958/https://www.intel.com/content/www/us/en/processors/xeon/scalable/xeon-scalable-spec-update.html
+
+2. Erratum BDF102 in Intel Xeon E5-2600 v4 Processor Product Family Specification Update:
+http://web.archive.org/web/20191125200531/https://www.intel.com/content/dam/www/public/us/en/documents/specification-updates/xeon-e5-v4-spec-update.pdf
+
+3. The errata in Intel Resource Director Technology (Intel RDT) on 2nd Generation Intel Xeon Scalable Processors Reference Manual:
+https://software.intel.com/content/www/us/en/develop/articles/intel-resource-director-technology-rdt-reference-manual.html
+
+for further information.
diff --git a/Documentation/x86/sgx.rst b/Documentation/x86/sgx.rst
new file mode 100644
index 000000000..2bcbffacb
--- /dev/null
+++ b/Documentation/x86/sgx.rst
@@ -0,0 +1,302 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===============================
+Software Guard eXtensions (SGX)
+===============================
+
+Overview
+========
+
+Software Guard eXtensions (SGX) hardware enables for user space applications
+to set aside private memory regions of code and data:
+
+* Privileged (ring-0) ENCLS functions orchestrate the construction of the
+ regions.
+* Unprivileged (ring-3) ENCLU functions allow an application to enter and
+ execute inside the regions.
+
+These memory regions are called enclaves. An enclave can be only entered at a
+fixed set of entry points. Each entry point can hold a single hardware thread
+at a time. While the enclave is loaded from a regular binary file by using
+ENCLS functions, only the threads inside the enclave can access its memory. The
+region is denied from outside access by the CPU, and encrypted before it leaves
+from LLC.
+
+The support can be determined by
+
+ ``grep sgx /proc/cpuinfo``
+
+SGX must both be supported in the processor and enabled by the BIOS. If SGX
+appears to be unsupported on a system which has hardware support, ensure
+support is enabled in the BIOS. If a BIOS presents a choice between "Enabled"
+and "Software Enabled" modes for SGX, choose "Enabled".
+
+Enclave Page Cache
+==================
+
+SGX utilizes an *Enclave Page Cache (EPC)* to store pages that are associated
+with an enclave. It is contained in a BIOS-reserved region of physical memory.
+Unlike pages used for regular memory, pages can only be accessed from outside of
+the enclave during enclave construction with special, limited SGX instructions.
+
+Only a CPU executing inside an enclave can directly access enclave memory.
+However, a CPU executing inside an enclave may access normal memory outside the
+enclave.
+
+The kernel manages enclave memory similar to how it treats device memory.
+
+Enclave Page Types
+------------------
+
+**SGX Enclave Control Structure (SECS)**
+ Enclave's address range, attributes and other global data are defined
+ by this structure.
+
+**Regular (REG)**
+ Regular EPC pages contain the code and data of an enclave.
+
+**Thread Control Structure (TCS)**
+ Thread Control Structure pages define the entry points to an enclave and
+ track the execution state of an enclave thread.
+
+**Version Array (VA)**
+ Version Array pages contain 512 slots, each of which can contain a version
+ number for a page evicted from the EPC.
+
+Enclave Page Cache Map
+----------------------
+
+The processor tracks EPC pages in a hardware metadata structure called the
+*Enclave Page Cache Map (EPCM)*. The EPCM contains an entry for each EPC page
+which describes the owning enclave, access rights and page type among the other
+things.
+
+EPCM permissions are separate from the normal page tables. This prevents the
+kernel from, for instance, allowing writes to data which an enclave wishes to
+remain read-only. EPCM permissions may only impose additional restrictions on
+top of normal x86 page permissions.
+
+For all intents and purposes, the SGX architecture allows the processor to
+invalidate all EPCM entries at will. This requires that software be prepared to
+handle an EPCM fault at any time. In practice, this can happen on events like
+power transitions when the ephemeral key that encrypts enclave memory is lost.
+
+Application interface
+=====================
+
+Enclave build functions
+-----------------------
+
+In addition to the traditional compiler and linker build process, SGX has a
+separate enclave “build” process. Enclaves must be built before they can be
+executed (entered). The first step in building an enclave is opening the
+**/dev/sgx_enclave** device. Since enclave memory is protected from direct
+access, special privileged instructions are then used to copy data into enclave
+pages and establish enclave page permissions.
+
+.. kernel-doc:: arch/x86/kernel/cpu/sgx/ioctl.c
+ :functions: sgx_ioc_enclave_create
+ sgx_ioc_enclave_add_pages
+ sgx_ioc_enclave_init
+ sgx_ioc_enclave_provision
+
+Enclave runtime management
+--------------------------
+
+Systems supporting SGX2 additionally support changes to initialized
+enclaves: modifying enclave page permissions and type, and dynamically
+adding and removing of enclave pages. When an enclave accesses an address
+within its address range that does not have a backing page then a new
+regular page will be dynamically added to the enclave. The enclave is
+still required to run EACCEPT on the new page before it can be used.
+
+.. kernel-doc:: arch/x86/kernel/cpu/sgx/ioctl.c
+ :functions: sgx_ioc_enclave_restrict_permissions
+ sgx_ioc_enclave_modify_types
+ sgx_ioc_enclave_remove_pages
+
+Enclave vDSO
+------------
+
+Entering an enclave can only be done through SGX-specific EENTER and ERESUME
+functions, and is a non-trivial process. Because of the complexity of
+transitioning to and from an enclave, enclaves typically utilize a library to
+handle the actual transitions. This is roughly analogous to how glibc
+implementations are used by most applications to wrap system calls.
+
+Another crucial characteristic of enclaves is that they can generate exceptions
+as part of their normal operation that need to be handled in the enclave or are
+unique to SGX.
+
+Instead of the traditional signal mechanism to handle these exceptions, SGX
+can leverage special exception fixup provided by the vDSO. The kernel-provided
+vDSO function wraps low-level transitions to/from the enclave like EENTER and
+ERESUME. The vDSO function intercepts exceptions that would otherwise generate
+a signal and return the fault information directly to its caller. This avoids
+the need to juggle signal handlers.
+
+.. kernel-doc:: arch/x86/include/uapi/asm/sgx.h
+ :functions: vdso_sgx_enter_enclave_t
+
+ksgxd
+=====
+
+SGX support includes a kernel thread called *ksgxd*.
+
+EPC sanitization
+----------------
+
+ksgxd is started when SGX initializes. Enclave memory is typically ready
+for use when the processor powers on or resets. However, if SGX has been in
+use since the reset, enclave pages may be in an inconsistent state. This might
+occur after a crash and kexec() cycle, for instance. At boot, ksgxd
+reinitializes all enclave pages so that they can be allocated and re-used.
+
+The sanitization is done by going through EPC address space and applying the
+EREMOVE function to each physical page. Some enclave pages like SECS pages have
+hardware dependencies on other pages which prevents EREMOVE from functioning.
+Executing two EREMOVE passes removes the dependencies.
+
+Page reclaimer
+--------------
+
+Similar to the core kswapd, ksgxd, is responsible for managing the
+overcommitment of enclave memory. If the system runs out of enclave memory,
+*ksgxd* “swaps” enclave memory to normal memory.
+
+Launch Control
+==============
+
+SGX provides a launch control mechanism. After all enclave pages have been
+copied, kernel executes EINIT function, which initializes the enclave. Only after
+this the CPU can execute inside the enclave.
+
+EINIT function takes an RSA-3072 signature of the enclave measurement. The function
+checks that the measurement is correct and signature is signed with the key
+hashed to the four **IA32_SGXLEPUBKEYHASH{0, 1, 2, 3}** MSRs representing the
+SHA256 of a public key.
+
+Those MSRs can be configured by the BIOS to be either readable or writable.
+Linux supports only writable configuration in order to give full control to the
+kernel on launch control policy. Before calling EINIT function, the driver sets
+the MSRs to match the enclave's signing key.
+
+Encryption engines
+==================
+
+In order to conceal the enclave data while it is out of the CPU package, the
+memory controller has an encryption engine to transparently encrypt and decrypt
+enclave memory.
+
+In CPUs prior to Ice Lake, the Memory Encryption Engine (MEE) is used to
+encrypt pages leaving the CPU caches. MEE uses a n-ary Merkle tree with root in
+SRAM to maintain integrity of the encrypted data. This provides integrity and
+anti-replay protection but does not scale to large memory sizes because the time
+required to update the Merkle tree grows logarithmically in relation to the
+memory size.
+
+CPUs starting from Icelake use Total Memory Encryption (TME) in the place of
+MEE. TME-based SGX implementations do not have an integrity Merkle tree, which
+means integrity and replay-attacks are not mitigated. B, it includes
+additional changes to prevent cipher text from being returned and SW memory
+aliases from being created.
+
+DMA to enclave memory is blocked by range registers on both MEE and TME systems
+(SDM section 41.10).
+
+Usage Models
+============
+
+Shared Library
+--------------
+
+Sensitive data and the code that acts on it is partitioned from the application
+into a separate library. The library is then linked as a DSO which can be loaded
+into an enclave. The application can then make individual function calls into
+the enclave through special SGX instructions. A run-time within the enclave is
+configured to marshal function parameters into and out of the enclave and to
+call the correct library function.
+
+Application Container
+---------------------
+
+An application may be loaded into a container enclave which is specially
+configured with a library OS and run-time which permits the application to run.
+The enclave run-time and library OS work together to execute the application
+when a thread enters the enclave.
+
+Impact of Potential Kernel SGX Bugs
+===================================
+
+EPC leaks
+---------
+
+When EPC page leaks happen, a WARNING like this is shown in dmesg:
+
+"EREMOVE returned ... and an EPC page was leaked. SGX may become unusable..."
+
+This is effectively a kernel use-after-free of an EPC page, and due
+to the way SGX works, the bug is detected at freeing. Rather than
+adding the page back to the pool of available EPC pages, the kernel
+intentionally leaks the page to avoid additional errors in the future.
+
+When this happens, the kernel will likely soon leak more EPC pages, and
+SGX will likely become unusable because the memory available to SGX is
+limited. However, while this may be fatal to SGX, the rest of the kernel
+is unlikely to be impacted and should continue to work.
+
+As a result, when this happpens, user should stop running any new
+SGX workloads, (or just any new workloads), and migrate all valuable
+workloads. Although a machine reboot can recover all EPC memory, the bug
+should be reported to Linux developers.
+
+
+Virtual EPC
+===========
+
+The implementation has also a virtual EPC driver to support SGX enclaves
+in guests. Unlike the SGX driver, an EPC page allocated by the virtual
+EPC driver doesn't have a specific enclave associated with it. This is
+because KVM doesn't track how a guest uses EPC pages.
+
+As a result, the SGX core page reclaimer doesn't support reclaiming EPC
+pages allocated to KVM guests through the virtual EPC driver. If the
+user wants to deploy SGX applications both on the host and in guests
+on the same machine, the user should reserve enough EPC (by taking out
+total virtual EPC size of all SGX VMs from the physical EPC size) for
+host SGX applications so they can run with acceptable performance.
+
+Architectural behavior is to restore all EPC pages to an uninitialized
+state also after a guest reboot. Because this state can be reached only
+through the privileged ``ENCLS[EREMOVE]`` instruction, ``/dev/sgx_vepc``
+provides the ``SGX_IOC_VEPC_REMOVE_ALL`` ioctl to execute the instruction
+on all pages in the virtual EPC.
+
+``EREMOVE`` can fail for three reasons. Userspace must pay attention
+to expected failures and handle them as follows:
+
+1. Page removal will always fail when any thread is running in the
+ enclave to which the page belongs. In this case the ioctl will
+ return ``EBUSY`` independent of whether it has successfully removed
+ some pages; userspace can avoid these failures by preventing execution
+ of any vcpu which maps the virtual EPC.
+
+2. Page removal will cause a general protection fault if two calls to
+ ``EREMOVE`` happen concurrently for pages that refer to the same
+ "SECS" metadata pages. This can happen if there are concurrent
+ invocations to ``SGX_IOC_VEPC_REMOVE_ALL``, or if a ``/dev/sgx_vepc``
+ file descriptor in the guest is closed at the same time as
+ ``SGX_IOC_VEPC_REMOVE_ALL``; it will also be reported as ``EBUSY``.
+ This can be avoided in userspace by serializing calls to the ioctl()
+ and to close(), but in general it should not be a problem.
+
+3. Finally, page removal will fail for SECS metadata pages which still
+ have child pages. Child pages can be removed by executing
+ ``SGX_IOC_VEPC_REMOVE_ALL`` on all ``/dev/sgx_vepc`` file descriptors
+ mapped into the guest. This means that the ioctl() must be called
+ twice: an initial set of calls to remove child pages and a subsequent
+ set of calls to remove SECS pages. The second set of calls is only
+ required for those mappings that returned a nonzero value from the
+ first call. It indicates a bug in the kernel or the userspace client
+ if any of the second round of ``SGX_IOC_VEPC_REMOVE_ALL`` calls has
+ a return code other than 0.
diff --git a/Documentation/x86/sva.rst b/Documentation/x86/sva.rst
new file mode 100644
index 000000000..2e9b8b0f9
--- /dev/null
+++ b/Documentation/x86/sva.rst
@@ -0,0 +1,286 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===========================================
+Shared Virtual Addressing (SVA) with ENQCMD
+===========================================
+
+Background
+==========
+
+Shared Virtual Addressing (SVA) allows the processor and device to use the
+same virtual addresses avoiding the need for software to translate virtual
+addresses to physical addresses. SVA is what PCIe calls Shared Virtual
+Memory (SVM).
+
+In addition to the convenience of using application virtual addresses
+by the device, it also doesn't require pinning pages for DMA.
+PCIe Address Translation Services (ATS) along with Page Request Interface
+(PRI) allow devices to function much the same way as the CPU handling
+application page-faults. For more information please refer to the PCIe
+specification Chapter 10: ATS Specification.
+
+Use of SVA requires IOMMU support in the platform. IOMMU is also
+required to support the PCIe features ATS and PRI. ATS allows devices
+to cache translations for virtual addresses. The IOMMU driver uses the
+mmu_notifier() support to keep the device TLB cache and the CPU cache in
+sync. When an ATS lookup fails for a virtual address, the device should
+use the PRI in order to request the virtual address to be paged into the
+CPU page tables. The device must use ATS again in order the fetch the
+translation before use.
+
+Shared Hardware Workqueues
+==========================
+
+Unlike Single Root I/O Virtualization (SR-IOV), Scalable IOV (SIOV) permits
+the use of Shared Work Queues (SWQ) by both applications and Virtual
+Machines (VM's). This allows better hardware utilization vs. hard
+partitioning resources that could result in under utilization. In order to
+allow the hardware to distinguish the context for which work is being
+executed in the hardware by SWQ interface, SIOV uses Process Address Space
+ID (PASID), which is a 20-bit number defined by the PCIe SIG.
+
+PASID value is encoded in all transactions from the device. This allows the
+IOMMU to track I/O on a per-PASID granularity in addition to using the PCIe
+Resource Identifier (RID) which is the Bus/Device/Function.
+
+
+ENQCMD
+======
+
+ENQCMD is a new instruction on Intel platforms that atomically submits a
+work descriptor to a device. The descriptor includes the operation to be
+performed, virtual addresses of all parameters, virtual address of a completion
+record, and the PASID (process address space ID) of the current process.
+
+ENQCMD works with non-posted semantics and carries a status back if the
+command was accepted by hardware. This allows the submitter to know if the
+submission needs to be retried or other device specific mechanisms to
+implement fairness or ensure forward progress should be provided.
+
+ENQCMD is the glue that ensures applications can directly submit commands
+to the hardware and also permits hardware to be aware of application context
+to perform I/O operations via use of PASID.
+
+Process Address Space Tagging
+=============================
+
+A new thread-scoped MSR (IA32_PASID) provides the connection between
+user processes and the rest of the hardware. When an application first
+accesses an SVA-capable device, this MSR is initialized with a newly
+allocated PASID. The driver for the device calls an IOMMU-specific API
+that sets up the routing for DMA and page-requests.
+
+For example, the Intel Data Streaming Accelerator (DSA) uses
+iommu_sva_bind_device(), which will do the following:
+
+- Allocate the PASID, and program the process page-table (%cr3 register) in the
+ PASID context entries.
+- Register for mmu_notifier() to track any page-table invalidations to keep
+ the device TLB in sync. For example, when a page-table entry is invalidated,
+ the IOMMU propagates the invalidation to the device TLB. This will force any
+ future access by the device to this virtual address to participate in
+ ATS. If the IOMMU responds with proper response that a page is not
+ present, the device would request the page to be paged in via the PCIe PRI
+ protocol before performing I/O.
+
+This MSR is managed with the XSAVE feature set as "supervisor state" to
+ensure the MSR is updated during context switch.
+
+PASID Management
+================
+
+The kernel must allocate a PASID on behalf of each process which will use
+ENQCMD and program it into the new MSR to communicate the process identity to
+platform hardware. ENQCMD uses the PASID stored in this MSR to tag requests
+from this process. When a user submits a work descriptor to a device using the
+ENQCMD instruction, the PASID field in the descriptor is auto-filled with the
+value from MSR_IA32_PASID. Requests for DMA from the device are also tagged
+with the same PASID. The platform IOMMU uses the PASID in the transaction to
+perform address translation. The IOMMU APIs setup the corresponding PASID
+entry in IOMMU with the process address used by the CPU (e.g. %cr3 register in
+x86).
+
+The MSR must be configured on each logical CPU before any application
+thread can interact with a device. Threads that belong to the same
+process share the same page tables, thus the same MSR value.
+
+PASID Life Cycle Management
+===========================
+
+PASID is initialized as INVALID_IOASID (-1) when a process is created.
+
+Only processes that access SVA-capable devices need to have a PASID
+allocated. This allocation happens when a process opens/binds an SVA-capable
+device but finds no PASID for this process. Subsequent binds of the same, or
+other devices will share the same PASID.
+
+Although the PASID is allocated to the process by opening a device,
+it is not active in any of the threads of that process. It's loaded to the
+IA32_PASID MSR lazily when a thread tries to submit a work descriptor
+to a device using the ENQCMD.
+
+That first access will trigger a #GP fault because the IA32_PASID MSR
+has not been initialized with the PASID value assigned to the process
+when the device was opened. The Linux #GP handler notes that a PASID has
+been allocated for the process, and so initializes the IA32_PASID MSR
+and returns so that the ENQCMD instruction is re-executed.
+
+On fork(2) or exec(2) the PASID is removed from the process as it no
+longer has the same address space that it had when the device was opened.
+
+On clone(2) the new task shares the same address space, so will be
+able to use the PASID allocated to the process. The IA32_PASID is not
+preemptively initialized as the PASID value might not be allocated yet or
+the kernel does not know whether this thread is going to access the device
+and the cleared IA32_PASID MSR reduces context switch overhead by xstate
+init optimization. Since #GP faults have to be handled on any threads that
+were created before the PASID was assigned to the mm of the process, newly
+created threads might as well be treated in a consistent way.
+
+Due to complexity of freeing the PASID and clearing all IA32_PASID MSRs in
+all threads in unbind, free the PASID lazily only on mm exit.
+
+If a process does a close(2) of the device file descriptor and munmap(2)
+of the device MMIO portal, then the driver will unbind the device. The
+PASID is still marked VALID in the PASID_MSR for any threads in the
+process that accessed the device. But this is harmless as without the
+MMIO portal they cannot submit new work to the device.
+
+Relationships
+=============
+
+ * Each process has many threads, but only one PASID.
+ * Devices have a limited number (~10's to 1000's) of hardware workqueues.
+ The device driver manages allocating hardware workqueues.
+ * A single mmap() maps a single hardware workqueue as a "portal" and
+ each portal maps down to a single workqueue.
+ * For each device with which a process interacts, there must be
+ one or more mmap()'d portals.
+ * Many threads within a process can share a single portal to access
+ a single device.
+ * Multiple processes can separately mmap() the same portal, in
+ which case they still share one device hardware workqueue.
+ * The single process-wide PASID is used by all threads to interact
+ with all devices. There is not, for instance, a PASID for each
+ thread or each thread<->device pair.
+
+FAQ
+===
+
+* What is SVA/SVM?
+
+Shared Virtual Addressing (SVA) permits I/O hardware and the processor to
+work in the same address space, i.e., to share it. Some call it Shared
+Virtual Memory (SVM), but Linux community wanted to avoid confusing it with
+POSIX Shared Memory and Secure Virtual Machines which were terms already in
+circulation.
+
+* What is a PASID?
+
+A Process Address Space ID (PASID) is a PCIe-defined Transaction Layer Packet
+(TLP) prefix. A PASID is a 20-bit number allocated and managed by the OS.
+PASID is included in all transactions between the platform and the device.
+
+* How are shared workqueues different?
+
+Traditionally, in order for userspace applications to interact with hardware,
+there is a separate hardware instance required per process. For example,
+consider doorbells as a mechanism of informing hardware about work to process.
+Each doorbell is required to be spaced 4k (or page-size) apart for process
+isolation. This requires hardware to provision that space and reserve it in
+MMIO. This doesn't scale as the number of threads becomes quite large. The
+hardware also manages the queue depth for Shared Work Queues (SWQ), and
+consumers don't need to track queue depth. If there is no space to accept
+a command, the device will return an error indicating retry.
+
+A user should check Deferrable Memory Write (DMWr) capability on the device
+and only submits ENQCMD when the device supports it. In the new DMWr PCIe
+terminology, devices need to support DMWr completer capability. In addition,
+it requires all switch ports to support DMWr routing and must be enabled by
+the PCIe subsystem, much like how PCIe atomic operations are managed for
+instance.
+
+SWQ allows hardware to provision just a single address in the device. When
+used with ENQCMD to submit work, the device can distinguish the process
+submitting the work since it will include the PASID assigned to that
+process. This helps the device scale to a large number of processes.
+
+* Is this the same as a user space device driver?
+
+Communicating with the device via the shared workqueue is much simpler
+than a full blown user space driver. The kernel driver does all the
+initialization of the hardware. User space only needs to worry about
+submitting work and processing completions.
+
+* Is this the same as SR-IOV?
+
+Single Root I/O Virtualization (SR-IOV) focuses on providing independent
+hardware interfaces for virtualizing hardware. Hence, it's required to be
+almost fully functional interface to software supporting the traditional
+BARs, space for interrupts via MSI-X, its own register layout.
+Virtual Functions (VFs) are assisted by the Physical Function (PF)
+driver.
+
+Scalable I/O Virtualization builds on the PASID concept to create device
+instances for virtualization. SIOV requires host software to assist in
+creating virtual devices; each virtual device is represented by a PASID
+along with the bus/device/function of the device. This allows device
+hardware to optimize device resource creation and can grow dynamically on
+demand. SR-IOV creation and management is very static in nature. Consult
+references below for more details.
+
+* Why not just create a virtual function for each app?
+
+Creating PCIe SR-IOV type Virtual Functions (VF) is expensive. VFs require
+duplicated hardware for PCI config space and interrupts such as MSI-X.
+Resources such as interrupts have to be hard partitioned between VFs at
+creation time, and cannot scale dynamically on demand. The VFs are not
+completely independent from the Physical Function (PF). Most VFs require
+some communication and assistance from the PF driver. SIOV, in contrast,
+creates a software-defined device where all the configuration and control
+aspects are mediated via the slow path. The work submission and completion
+happen without any mediation.
+
+* Does this support virtualization?
+
+ENQCMD can be used from within a guest VM. In these cases, the VMM helps
+with setting up a translation table to translate from Guest PASID to Host
+PASID. Please consult the ENQCMD instruction set reference for more
+details.
+
+* Does memory need to be pinned?
+
+When devices support SVA along with platform hardware such as IOMMU
+supporting such devices, there is no need to pin memory for DMA purposes.
+Devices that support SVA also support other PCIe features that remove the
+pinning requirement for memory.
+
+Device TLB support - Device requests the IOMMU to lookup an address before
+use via Address Translation Service (ATS) requests. If the mapping exists
+but there is no page allocated by the OS, IOMMU hardware returns that no
+mapping exists.
+
+Device requests the virtual address to be mapped via Page Request
+Interface (PRI). Once the OS has successfully completed the mapping, it
+returns the response back to the device. The device requests again for
+a translation and continues.
+
+IOMMU works with the OS in managing consistency of page-tables with the
+device. When removing pages, it interacts with the device to remove any
+device TLB entry that might have been cached before removing the mappings from
+the OS.
+
+References
+==========
+
+VT-D:
+https://01.org/blogs/ashokraj/2018/recent-enhancements-intel-virtualization-technology-directed-i/o-intel-vt-d
+
+SIOV:
+https://01.org/blogs/2019/assignable-interfaces-intel-scalable-i/o-virtualization-linux
+
+ENQCMD in ISE:
+https://software.intel.com/sites/default/files/managed/c5/15/architecture-instruction-set-extensions-programming-reference.pdf
+
+DSA spec:
+https://software.intel.com/sites/default/files/341204-intel-data-streaming-accelerator-spec.pdf
diff --git a/Documentation/x86/tdx.rst b/Documentation/x86/tdx.rst
new file mode 100644
index 000000000..b8fa4329e
--- /dev/null
+++ b/Documentation/x86/tdx.rst
@@ -0,0 +1,218 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=====================================
+Intel Trust Domain Extensions (TDX)
+=====================================
+
+Intel's Trust Domain Extensions (TDX) protect confidential guest VMs from
+the host and physical attacks by isolating the guest register state and by
+encrypting the guest memory. In TDX, a special module running in a special
+mode sits between the host and the guest and manages the guest/host
+separation.
+
+Since the host cannot directly access guest registers or memory, much
+normal functionality of a hypervisor must be moved into the guest. This is
+implemented using a Virtualization Exception (#VE) that is handled by the
+guest kernel. A #VE is handled entirely inside the guest kernel, but some
+require the hypervisor to be consulted.
+
+TDX includes new hypercall-like mechanisms for communicating from the
+guest to the hypervisor or the TDX module.
+
+New TDX Exceptions
+==================
+
+TDX guests behave differently from bare-metal and traditional VMX guests.
+In TDX guests, otherwise normal instructions or memory accesses can cause
+#VE or #GP exceptions.
+
+Instructions marked with an '*' conditionally cause exceptions. The
+details for these instructions are discussed below.
+
+Instruction-based #VE
+---------------------
+
+- Port I/O (INS, OUTS, IN, OUT)
+- HLT
+- MONITOR, MWAIT
+- WBINVD, INVD
+- VMCALL
+- RDMSR*,WRMSR*
+- CPUID*
+
+Instruction-based #GP
+---------------------
+
+- All VMX instructions: INVEPT, INVVPID, VMCLEAR, VMFUNC, VMLAUNCH,
+ VMPTRLD, VMPTRST, VMREAD, VMRESUME, VMWRITE, VMXOFF, VMXON
+- ENCLS, ENCLU
+- GETSEC
+- RSM
+- ENQCMD
+- RDMSR*,WRMSR*
+
+RDMSR/WRMSR Behavior
+--------------------
+
+MSR access behavior falls into three categories:
+
+- #GP generated
+- #VE generated
+- "Just works"
+
+In general, the #GP MSRs should not be used in guests. Their use likely
+indicates a bug in the guest. The guest may try to handle the #GP with a
+hypercall but it is unlikely to succeed.
+
+The #VE MSRs are typically able to be handled by the hypervisor. Guests
+can make a hypercall to the hypervisor to handle the #VE.
+
+The "just works" MSRs do not need any special guest handling. They might
+be implemented by directly passing through the MSR to the hardware or by
+trapping and handling in the TDX module. Other than possibly being slow,
+these MSRs appear to function just as they would on bare metal.
+
+CPUID Behavior
+--------------
+
+For some CPUID leaves and sub-leaves, the virtualized bit fields of CPUID
+return values (in guest EAX/EBX/ECX/EDX) are configurable by the
+hypervisor. For such cases, the Intel TDX module architecture defines two
+virtualization types:
+
+- Bit fields for which the hypervisor controls the value seen by the guest
+ TD.
+
+- Bit fields for which the hypervisor configures the value such that the
+ guest TD either sees their native value or a value of 0. For these bit
+ fields, the hypervisor can mask off the native values, but it can not
+ turn *on* values.
+
+A #VE is generated for CPUID leaves and sub-leaves that the TDX module does
+not know how to handle. The guest kernel may ask the hypervisor for the
+value with a hypercall.
+
+#VE on Memory Accesses
+======================
+
+There are essentially two classes of TDX memory: private and shared.
+Private memory receives full TDX protections. Its content is protected
+against access from the hypervisor. Shared memory is expected to be
+shared between guest and hypervisor and does not receive full TDX
+protections.
+
+A TD guest is in control of whether its memory accesses are treated as
+private or shared. It selects the behavior with a bit in its page table
+entries. This helps ensure that a guest does not place sensitive
+information in shared memory, exposing it to the untrusted hypervisor.
+
+#VE on Shared Memory
+--------------------
+
+Access to shared mappings can cause a #VE. The hypervisor ultimately
+controls whether a shared memory access causes a #VE, so the guest must be
+careful to only reference shared pages it can safely handle a #VE. For
+instance, the guest should be careful not to access shared memory in the
+#VE handler before it reads the #VE info structure (TDG.VP.VEINFO.GET).
+
+Shared mapping content is entirely controlled by the hypervisor. The guest
+should only use shared mappings for communicating with the hypervisor.
+Shared mappings must never be used for sensitive memory content like kernel
+stacks. A good rule of thumb is that hypervisor-shared memory should be
+treated the same as memory mapped to userspace. Both the hypervisor and
+userspace are completely untrusted.
+
+MMIO for virtual devices is implemented as shared memory. The guest must
+be careful not to access device MMIO regions unless it is also prepared to
+handle a #VE.
+
+#VE on Private Pages
+--------------------
+
+An access to private mappings can also cause a #VE. Since all kernel
+memory is also private memory, the kernel might theoretically need to
+handle a #VE on arbitrary kernel memory accesses. This is not feasible, so
+TDX guests ensure that all guest memory has been "accepted" before memory
+is used by the kernel.
+
+A modest amount of memory (typically 512M) is pre-accepted by the firmware
+before the kernel runs to ensure that the kernel can start up without
+being subjected to a #VE.
+
+The hypervisor is permitted to unilaterally move accepted pages to a
+"blocked" state. However, if it does this, page access will not generate a
+#VE. It will, instead, cause a "TD Exit" where the hypervisor is required
+to handle the exception.
+
+Linux #VE handler
+=================
+
+Just like page faults or #GP's, #VE exceptions can be either handled or be
+fatal. Typically, an unhandled userspace #VE results in a SIGSEGV.
+An unhandled kernel #VE results in an oops.
+
+Handling nested exceptions on x86 is typically nasty business. A #VE
+could be interrupted by an NMI which triggers another #VE and hilarity
+ensues. The TDX #VE architecture anticipated this scenario and includes a
+feature to make it slightly less nasty.
+
+During #VE handling, the TDX module ensures that all interrupts (including
+NMIs) are blocked. The block remains in place until the guest makes a
+TDG.VP.VEINFO.GET TDCALL. This allows the guest to control when interrupts
+or a new #VE can be delivered.
+
+However, the guest kernel must still be careful to avoid potential
+#VE-triggering actions (discussed above) while this block is in place.
+While the block is in place, any #VE is elevated to a double fault (#DF)
+which is not recoverable.
+
+MMIO handling
+=============
+
+In non-TDX VMs, MMIO is usually implemented by giving a guest access to a
+mapping which will cause a VMEXIT on access, and then the hypervisor
+emulates the access. That is not possible in TDX guests because VMEXIT
+will expose the register state to the host. TDX guests don't trust the host
+and can't have their state exposed to the host.
+
+In TDX, MMIO regions typically trigger a #VE exception in the guest. The
+guest #VE handler then emulates the MMIO instruction inside the guest and
+converts it into a controlled TDCALL to the host, rather than exposing
+guest state to the host.
+
+MMIO addresses on x86 are just special physical addresses. They can
+theoretically be accessed with any instruction that accesses memory.
+However, the kernel instruction decoding method is limited. It is only
+designed to decode instructions like those generated by io.h macros.
+
+MMIO access via other means (like structure overlays) may result in an
+oops.
+
+Shared Memory Conversions
+=========================
+
+All TDX guest memory starts out as private at boot. This memory can not
+be accessed by the hypervisor. However, some kernel users like device
+drivers might have a need to share data with the hypervisor. To do this,
+memory must be converted between shared and private. This can be
+accomplished using some existing memory encryption helpers:
+
+ * set_memory_decrypted() converts a range of pages to shared.
+ * set_memory_encrypted() converts memory back to private.
+
+Device drivers are the primary user of shared memory, but there's no need
+to touch every driver. DMA buffers and ioremap() do the conversions
+automatically.
+
+TDX uses SWIOTLB for most DMA allocations. The SWIOTLB buffer is
+converted to shared on boot.
+
+For coherent DMA allocation, the DMA buffer gets converted on the
+allocation. Check force_dma_unencrypted() for details.
+
+References
+==========
+
+TDX reference material is collected here:
+
+https://www.intel.com/content/www/us/en/developer/articles/technical/intel-trust-domain-extensions.html
diff --git a/Documentation/x86/tlb.rst b/Documentation/x86/tlb.rst
new file mode 100644
index 000000000..82ec58ae6
--- /dev/null
+++ b/Documentation/x86/tlb.rst
@@ -0,0 +1,83 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=======
+The TLB
+=======
+
+When the kernel unmaps or modified the attributes of a range of
+memory, it has two choices:
+
+ 1. Flush the entire TLB with a two-instruction sequence. This is
+ a quick operation, but it causes collateral damage: TLB entries
+ from areas other than the one we are trying to flush will be
+ destroyed and must be refilled later, at some cost.
+ 2. Use the invlpg instruction to invalidate a single page at a
+ time. This could potentially cost many more instructions, but
+ it is a much more precise operation, causing no collateral
+ damage to other TLB entries.
+
+Which method to do depends on a few things:
+
+ 1. The size of the flush being performed. A flush of the entire
+ address space is obviously better performed by flushing the
+ entire TLB than doing 2^48/PAGE_SIZE individual flushes.
+ 2. The contents of the TLB. If the TLB is empty, then there will
+ be no collateral damage caused by doing the global flush, and
+ all of the individual flush will have ended up being wasted
+ work.
+ 3. The size of the TLB. The larger the TLB, the more collateral
+ damage we do with a full flush. So, the larger the TLB, the
+ more attractive an individual flush looks. Data and
+ instructions have separate TLBs, as do different page sizes.
+ 4. The microarchitecture. The TLB has become a multi-level
+ cache on modern CPUs, and the global flushes have become more
+ expensive relative to single-page flushes.
+
+There is obviously no way the kernel can know all these things,
+especially the contents of the TLB during a given flush. The
+sizes of the flush will vary greatly depending on the workload as
+well. There is essentially no "right" point to choose.
+
+You may be doing too many individual invalidations if you see the
+invlpg instruction (or instructions _near_ it) show up high in
+profiles. If you believe that individual invalidations being
+called too often, you can lower the tunable::
+
+ /sys/kernel/debug/x86/tlb_single_page_flush_ceiling
+
+This will cause us to do the global flush for more cases.
+Lowering it to 0 will disable the use of the individual flushes.
+Setting it to 1 is a very conservative setting and it should
+never need to be 0 under normal circumstances.
+
+Despite the fact that a single individual flush on x86 is
+guaranteed to flush a full 2MB [1]_, hugetlbfs always uses the full
+flushes. THP is treated exactly the same as normal memory.
+
+You might see invlpg inside of flush_tlb_mm_range() show up in
+profiles, or you can use the trace_tlb_flush() tracepoints. to
+determine how long the flush operations are taking.
+
+Essentially, you are balancing the cycles you spend doing invlpg
+with the cycles that you spend refilling the TLB later.
+
+You can measure how expensive TLB refills are by using
+performance counters and 'perf stat', like this::
+
+ perf stat -e
+ cpu/event=0x8,umask=0x84,name=dtlb_load_misses_walk_duration/,
+ cpu/event=0x8,umask=0x82,name=dtlb_load_misses_walk_completed/,
+ cpu/event=0x49,umask=0x4,name=dtlb_store_misses_walk_duration/,
+ cpu/event=0x49,umask=0x2,name=dtlb_store_misses_walk_completed/,
+ cpu/event=0x85,umask=0x4,name=itlb_misses_walk_duration/,
+ cpu/event=0x85,umask=0x2,name=itlb_misses_walk_completed/
+
+That works on an IvyBridge-era CPU (i5-3320M). Different CPUs
+may have differently-named counters, but they should at least
+be there in some form. You can use pmu-tools 'ocperf list'
+(https://github.com/andikleen/pmu-tools) to find the right
+counters for a given CPU.
+
+.. [1] A footnote in Intel's SDM "4.10.4.2 Recommended Invalidation"
+ says: "One execution of INVLPG is sufficient even for a page
+ with size greater than 4 KBytes."
diff --git a/Documentation/x86/topology.rst b/Documentation/x86/topology.rst
new file mode 100644
index 000000000..7f58010ea
--- /dev/null
+++ b/Documentation/x86/topology.rst
@@ -0,0 +1,234 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+============
+x86 Topology
+============
+
+This documents and clarifies the main aspects of x86 topology modelling and
+representation in the kernel. Update/change when doing changes to the
+respective code.
+
+The architecture-agnostic topology definitions are in
+Documentation/admin-guide/cputopology.rst. This file holds x86-specific
+differences/specialities which must not necessarily apply to the generic
+definitions. Thus, the way to read up on Linux topology on x86 is to start
+with the generic one and look at this one in parallel for the x86 specifics.
+
+Needless to say, code should use the generic functions - this file is *only*
+here to *document* the inner workings of x86 topology.
+
+Started by Thomas Gleixner <tglx@linutronix.de> and Borislav Petkov <bp@alien8.de>.
+
+The main aim of the topology facilities is to present adequate interfaces to
+code which needs to know/query/use the structure of the running system wrt
+threads, cores, packages, etc.
+
+The kernel does not care about the concept of physical sockets because a
+socket has no relevance to software. It's an electromechanical component. In
+the past a socket always contained a single package (see below), but with the
+advent of Multi Chip Modules (MCM) a socket can hold more than one package. So
+there might be still references to sockets in the code, but they are of
+historical nature and should be cleaned up.
+
+The topology of a system is described in the units of:
+
+ - packages
+ - cores
+ - threads
+
+Package
+=======
+Packages contain a number of cores plus shared resources, e.g. DRAM
+controller, shared caches etc.
+
+Modern systems may also use the term 'Die' for package.
+
+AMD nomenclature for package is 'Node'.
+
+Package-related topology information in the kernel:
+
+ - cpuinfo_x86.x86_max_cores:
+
+ The number of cores in a package. This information is retrieved via CPUID.
+
+ - cpuinfo_x86.x86_max_dies:
+
+ The number of dies in a package. This information is retrieved via CPUID.
+
+ - cpuinfo_x86.cpu_die_id:
+
+ The physical ID of the die. This information is retrieved via CPUID.
+
+ - cpuinfo_x86.phys_proc_id:
+
+ The physical ID of the package. This information is retrieved via CPUID
+ and deduced from the APIC IDs of the cores in the package.
+
+ Modern systems use this value for the socket. There may be multiple
+ packages within a socket. This value may differ from cpu_die_id.
+
+ - cpuinfo_x86.logical_proc_id:
+
+ The logical ID of the package. As we do not trust BIOSes to enumerate the
+ packages in a consistent way, we introduced the concept of logical package
+ ID so we can sanely calculate the number of maximum possible packages in
+ the system and have the packages enumerated linearly.
+
+ - topology_max_packages():
+
+ The maximum possible number of packages in the system. Helpful for per
+ package facilities to preallocate per package information.
+
+ - cpu_llc_id:
+
+ A per-CPU variable containing:
+
+ - On Intel, the first APIC ID of the list of CPUs sharing the Last Level
+ Cache
+
+ - On AMD, the Node ID or Core Complex ID containing the Last Level
+ Cache. In general, it is a number identifying an LLC uniquely on the
+ system.
+
+Cores
+=====
+A core consists of 1 or more threads. It does not matter whether the threads
+are SMT- or CMT-type threads.
+
+AMDs nomenclature for a CMT core is "Compute Unit". The kernel always uses
+"core".
+
+Core-related topology information in the kernel:
+
+ - smp_num_siblings:
+
+ The number of threads in a core. The number of threads in a package can be
+ calculated by::
+
+ threads_per_package = cpuinfo_x86.x86_max_cores * smp_num_siblings
+
+
+Threads
+=======
+A thread is a single scheduling unit. It's the equivalent to a logical Linux
+CPU.
+
+AMDs nomenclature for CMT threads is "Compute Unit Core". The kernel always
+uses "thread".
+
+Thread-related topology information in the kernel:
+
+ - topology_core_cpumask():
+
+ The cpumask contains all online threads in the package to which a thread
+ belongs.
+
+ The number of online threads is also printed in /proc/cpuinfo "siblings."
+
+ - topology_sibling_cpumask():
+
+ The cpumask contains all online threads in the core to which a thread
+ belongs.
+
+ - topology_logical_package_id():
+
+ The logical package ID to which a thread belongs.
+
+ - topology_physical_package_id():
+
+ The physical package ID to which a thread belongs.
+
+ - topology_core_id();
+
+ The ID of the core to which a thread belongs. It is also printed in /proc/cpuinfo
+ "core_id."
+
+
+
+System topology examples
+========================
+
+.. note::
+ The alternative Linux CPU enumeration depends on how the BIOS enumerates the
+ threads. Many BIOSes enumerate all threads 0 first and then all threads 1.
+ That has the "advantage" that the logical Linux CPU numbers of threads 0 stay
+ the same whether threads are enabled or not. That's merely an implementation
+ detail and has no practical impact.
+
+1) Single Package, Single Core::
+
+ [package 0] -> [core 0] -> [thread 0] -> Linux CPU 0
+
+2) Single Package, Dual Core
+
+ a) One thread per core::
+
+ [package 0] -> [core 0] -> [thread 0] -> Linux CPU 0
+ -> [core 1] -> [thread 0] -> Linux CPU 1
+
+ b) Two threads per core::
+
+ [package 0] -> [core 0] -> [thread 0] -> Linux CPU 0
+ -> [thread 1] -> Linux CPU 1
+ -> [core 1] -> [thread 0] -> Linux CPU 2
+ -> [thread 1] -> Linux CPU 3
+
+ Alternative enumeration::
+
+ [package 0] -> [core 0] -> [thread 0] -> Linux CPU 0
+ -> [thread 1] -> Linux CPU 2
+ -> [core 1] -> [thread 0] -> Linux CPU 1
+ -> [thread 1] -> Linux CPU 3
+
+ AMD nomenclature for CMT systems::
+
+ [node 0] -> [Compute Unit 0] -> [Compute Unit Core 0] -> Linux CPU 0
+ -> [Compute Unit Core 1] -> Linux CPU 1
+ -> [Compute Unit 1] -> [Compute Unit Core 0] -> Linux CPU 2
+ -> [Compute Unit Core 1] -> Linux CPU 3
+
+4) Dual Package, Dual Core
+
+ a) One thread per core::
+
+ [package 0] -> [core 0] -> [thread 0] -> Linux CPU 0
+ -> [core 1] -> [thread 0] -> Linux CPU 1
+
+ [package 1] -> [core 0] -> [thread 0] -> Linux CPU 2
+ -> [core 1] -> [thread 0] -> Linux CPU 3
+
+ b) Two threads per core::
+
+ [package 0] -> [core 0] -> [thread 0] -> Linux CPU 0
+ -> [thread 1] -> Linux CPU 1
+ -> [core 1] -> [thread 0] -> Linux CPU 2
+ -> [thread 1] -> Linux CPU 3
+
+ [package 1] -> [core 0] -> [thread 0] -> Linux CPU 4
+ -> [thread 1] -> Linux CPU 5
+ -> [core 1] -> [thread 0] -> Linux CPU 6
+ -> [thread 1] -> Linux CPU 7
+
+ Alternative enumeration::
+
+ [package 0] -> [core 0] -> [thread 0] -> Linux CPU 0
+ -> [thread 1] -> Linux CPU 4
+ -> [core 1] -> [thread 0] -> Linux CPU 1
+ -> [thread 1] -> Linux CPU 5
+
+ [package 1] -> [core 0] -> [thread 0] -> Linux CPU 2
+ -> [thread 1] -> Linux CPU 6
+ -> [core 1] -> [thread 0] -> Linux CPU 3
+ -> [thread 1] -> Linux CPU 7
+
+ AMD nomenclature for CMT systems::
+
+ [node 0] -> [Compute Unit 0] -> [Compute Unit Core 0] -> Linux CPU 0
+ -> [Compute Unit Core 1] -> Linux CPU 1
+ -> [Compute Unit 1] -> [Compute Unit Core 0] -> Linux CPU 2
+ -> [Compute Unit Core 1] -> Linux CPU 3
+
+ [node 1] -> [Compute Unit 0] -> [Compute Unit Core 0] -> Linux CPU 4
+ -> [Compute Unit Core 1] -> Linux CPU 5
+ -> [Compute Unit 1] -> [Compute Unit Core 0] -> Linux CPU 6
+ -> [Compute Unit Core 1] -> Linux CPU 7
diff --git a/Documentation/x86/tsx_async_abort.rst b/Documentation/x86/tsx_async_abort.rst
new file mode 100644
index 000000000..583ddc185
--- /dev/null
+++ b/Documentation/x86/tsx_async_abort.rst
@@ -0,0 +1,117 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+TSX Async Abort (TAA) mitigation
+================================
+
+.. _tsx_async_abort:
+
+Overview
+--------
+
+TSX Async Abort (TAA) is a side channel attack on internal buffers in some
+Intel processors similar to Microachitectural Data Sampling (MDS). In this
+case certain loads may speculatively pass invalid data to dependent operations
+when an asynchronous abort condition is pending in a Transactional
+Synchronization Extensions (TSX) transaction. This includes loads with no
+fault or assist condition. Such loads may speculatively expose stale data from
+the same uarch data structures as in MDS, with same scope of exposure i.e.
+same-thread and cross-thread. This issue affects all current processors that
+support TSX.
+
+Mitigation strategy
+-------------------
+
+a) TSX disable - one of the mitigations is to disable TSX. A new MSR
+IA32_TSX_CTRL will be available in future and current processors after
+microcode update which can be used to disable TSX. In addition, it
+controls the enumeration of the TSX feature bits (RTM and HLE) in CPUID.
+
+b) Clear CPU buffers - similar to MDS, clearing the CPU buffers mitigates this
+vulnerability. More details on this approach can be found in
+:ref:`Documentation/admin-guide/hw-vuln/mds.rst <mds>`.
+
+Kernel internal mitigation modes
+--------------------------------
+
+ ============= ============================================================
+ off Mitigation is disabled. Either the CPU is not affected or
+ tsx_async_abort=off is supplied on the kernel command line.
+
+ tsx disabled Mitigation is enabled. TSX feature is disabled by default at
+ bootup on processors that support TSX control.
+
+ verw Mitigation is enabled. CPU is affected and MD_CLEAR is
+ advertised in CPUID.
+
+ ucode needed Mitigation is enabled. CPU is affected and MD_CLEAR is not
+ advertised in CPUID. That is mainly for virtualization
+ scenarios where the host has the updated microcode but the
+ hypervisor does not expose MD_CLEAR in CPUID. It's a best
+ effort approach without guarantee.
+ ============= ============================================================
+
+If the CPU is affected and the "tsx_async_abort" kernel command line parameter is
+not provided then the kernel selects an appropriate mitigation depending on the
+status of RTM and MD_CLEAR CPUID bits.
+
+Below tables indicate the impact of tsx=on|off|auto cmdline options on state of
+TAA mitigation, VERW behavior and TSX feature for various combinations of
+MSR_IA32_ARCH_CAPABILITIES bits.
+
+1. "tsx=off"
+
+========= ========= ============ ============ ============== =================== ======================
+MSR_IA32_ARCH_CAPABILITIES bits Result with cmdline tsx=off
+---------------------------------- -------------------------------------------------------------------------
+TAA_NO MDS_NO TSX_CTRL_MSR TSX state VERW can clear TAA mitigation TAA mitigation
+ after bootup CPU buffers tsx_async_abort=off tsx_async_abort=full
+========= ========= ============ ============ ============== =================== ======================
+ 0 0 0 HW default Yes Same as MDS Same as MDS
+ 0 0 1 Invalid case Invalid case Invalid case Invalid case
+ 0 1 0 HW default No Need ucode update Need ucode update
+ 0 1 1 Disabled Yes TSX disabled TSX disabled
+ 1 X 1 Disabled X None needed None needed
+========= ========= ============ ============ ============== =================== ======================
+
+2. "tsx=on"
+
+========= ========= ============ ============ ============== =================== ======================
+MSR_IA32_ARCH_CAPABILITIES bits Result with cmdline tsx=on
+---------------------------------- -------------------------------------------------------------------------
+TAA_NO MDS_NO TSX_CTRL_MSR TSX state VERW can clear TAA mitigation TAA mitigation
+ after bootup CPU buffers tsx_async_abort=off tsx_async_abort=full
+========= ========= ============ ============ ============== =================== ======================
+ 0 0 0 HW default Yes Same as MDS Same as MDS
+ 0 0 1 Invalid case Invalid case Invalid case Invalid case
+ 0 1 0 HW default No Need ucode update Need ucode update
+ 0 1 1 Enabled Yes None Same as MDS
+ 1 X 1 Enabled X None needed None needed
+========= ========= ============ ============ ============== =================== ======================
+
+3. "tsx=auto"
+
+========= ========= ============ ============ ============== =================== ======================
+MSR_IA32_ARCH_CAPABILITIES bits Result with cmdline tsx=auto
+---------------------------------- -------------------------------------------------------------------------
+TAA_NO MDS_NO TSX_CTRL_MSR TSX state VERW can clear TAA mitigation TAA mitigation
+ after bootup CPU buffers tsx_async_abort=off tsx_async_abort=full
+========= ========= ============ ============ ============== =================== ======================
+ 0 0 0 HW default Yes Same as MDS Same as MDS
+ 0 0 1 Invalid case Invalid case Invalid case Invalid case
+ 0 1 0 HW default No Need ucode update Need ucode update
+ 0 1 1 Disabled Yes TSX disabled TSX disabled
+ 1 X 1 Enabled X None needed None needed
+========= ========= ============ ============ ============== =================== ======================
+
+In the tables, TSX_CTRL_MSR is a new bit in MSR_IA32_ARCH_CAPABILITIES that
+indicates whether MSR_IA32_TSX_CTRL is supported.
+
+There are two control bits in IA32_TSX_CTRL MSR:
+
+ Bit 0: When set it disables the Restricted Transactional Memory (RTM)
+ sub-feature of TSX (will force all transactions to abort on the
+ XBEGIN instruction).
+
+ Bit 1: When set it disables the enumeration of the RTM and HLE feature
+ (i.e. it will make CPUID(EAX=7).EBX{bit4} and
+ CPUID(EAX=7).EBX{bit11} read as 0).
diff --git a/Documentation/x86/usb-legacy-support.rst b/Documentation/x86/usb-legacy-support.rst
new file mode 100644
index 000000000..e01c08b7c
--- /dev/null
+++ b/Documentation/x86/usb-legacy-support.rst
@@ -0,0 +1,50 @@
+
+.. SPDX-License-Identifier: GPL-2.0
+
+==================
+USB Legacy support
+==================
+
+:Author: Vojtech Pavlik <vojtech@suse.cz>, January 2004
+
+
+Also known as "USB Keyboard" or "USB Mouse support" in the BIOS Setup is a
+feature that allows one to use the USB mouse and keyboard as if they were
+their classic PS/2 counterparts. This means one can use an USB keyboard to
+type in LILO for example.
+
+It has several drawbacks, though:
+
+1) On some machines, the emulated PS/2 mouse takes over even when no USB
+ mouse is present and a real PS/2 mouse is present. In that case the extra
+ features (wheel, extra buttons, touchpad mode) of the real PS/2 mouse may
+ not be available.
+
+2) If CONFIG_HIGHMEM64G is enabled, the PS/2 mouse emulation can cause
+ system crashes, because the SMM BIOS is not expecting to be in PAE mode.
+ The Intel E7505 is a typical machine where this happens.
+
+3) If AMD64 64-bit mode is enabled, again system crashes often happen,
+ because the SMM BIOS isn't expecting the CPU to be in 64-bit mode. The
+ BIOS manufacturers only test with Windows, and Windows doesn't do 64-bit
+ yet.
+
+Solutions:
+
+Problem 1)
+ can be solved by loading the USB drivers prior to loading the
+ PS/2 mouse driver. Since the PS/2 mouse driver is in 2.6 compiled into
+ the kernel unconditionally, this means the USB drivers need to be
+ compiled-in, too.
+
+Problem 2)
+ can currently only be solved by either disabling HIGHMEM64G
+ in the kernel config or USB Legacy support in the BIOS. A BIOS update
+ could help, but so far no such update exists.
+
+Problem 3)
+ is usually fixed by a BIOS update. Check the board
+ manufacturers web site. If an update is not available, disable USB
+ Legacy support in the BIOS. If this alone doesn't help, try also adding
+ idle=poll on the kernel command line. The BIOS may be entering the SMM
+ on the HLT instruction as well.
diff --git a/Documentation/x86/x86_64/5level-paging.rst b/Documentation/x86/x86_64/5level-paging.rst
new file mode 100644
index 000000000..b792bbdc0
--- /dev/null
+++ b/Documentation/x86/x86_64/5level-paging.rst
@@ -0,0 +1,67 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==============
+5-level paging
+==============
+
+Overview
+========
+Original x86-64 was limited by 4-level paging to 256 TiB of virtual address
+space and 64 TiB of physical address space. We are already bumping into
+this limit: some vendors offer servers with 64 TiB of memory today.
+
+To overcome the limitation upcoming hardware will introduce support for
+5-level paging. It is a straight-forward extension of the current page
+table structure adding one more layer of translation.
+
+It bumps the limits to 128 PiB of virtual address space and 4 PiB of
+physical address space. This "ought to be enough for anybody" ©.
+
+QEMU 2.9 and later support 5-level paging.
+
+Virtual memory layout for 5-level paging is described in
+Documentation/x86/x86_64/mm.rst
+
+
+Enabling 5-level paging
+=======================
+CONFIG_X86_5LEVEL=y enables the feature.
+
+Kernel with CONFIG_X86_5LEVEL=y still able to boot on 4-level hardware.
+In this case additional page table level -- p4d -- will be folded at
+runtime.
+
+User-space and large virtual address space
+==========================================
+On x86, 5-level paging enables 56-bit userspace virtual address space.
+Not all user space is ready to handle wide addresses. It's known that
+at least some JIT compilers use higher bits in pointers to encode their
+information. It collides with valid pointers with 5-level paging and
+leads to crashes.
+
+To mitigate this, we are not going to allocate virtual address space
+above 47-bit by default.
+
+But userspace can ask for allocation from full address space by
+specifying hint address (with or without MAP_FIXED) above 47-bits.
+
+If hint address set above 47-bit, but MAP_FIXED is not specified, we try
+to look for unmapped area by specified address. If it's already
+occupied, we look for unmapped area in *full* address space, rather than
+from 47-bit window.
+
+A high hint address would only affect the allocation in question, but not
+any future mmap()s.
+
+Specifying high hint address on older kernel or on machine without 5-level
+paging support is safe. The hint will be ignored and kernel will fall back
+to allocation from 47-bit address space.
+
+This approach helps to easily make application's memory allocator aware
+about large address space without manually tracking allocated virtual
+address space.
+
+One important case we need to handle here is interaction with MPX.
+MPX (without MAWA extension) cannot handle addresses above 47-bit, so we
+need to make sure that MPX cannot be enabled we already have VMA above
+the boundary and forbid creating such VMAs once MPX is enabled.
diff --git a/Documentation/x86/x86_64/boot-options.rst b/Documentation/x86/x86_64/boot-options.rst
new file mode 100644
index 000000000..cbd14124a
--- /dev/null
+++ b/Documentation/x86/x86_64/boot-options.rst
@@ -0,0 +1,319 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===========================
+AMD64 Specific Boot Options
+===========================
+
+There are many others (usually documented in driver documentation), but
+only the AMD64 specific ones are listed here.
+
+Machine check
+=============
+Please see Documentation/x86/x86_64/machinecheck.rst for sysfs runtime tunables.
+
+ mce=off
+ Disable machine check
+ mce=no_cmci
+ Disable CMCI(Corrected Machine Check Interrupt) that
+ Intel processor supports. Usually this disablement is
+ not recommended, but it might be handy if your hardware
+ is misbehaving.
+ Note that you'll get more problems without CMCI than with
+ due to the shared banks, i.e. you might get duplicated
+ error logs.
+ mce=dont_log_ce
+ Don't make logs for corrected errors. All events reported
+ as corrected are silently cleared by OS.
+ This option will be useful if you have no interest in any
+ of corrected errors.
+ mce=ignore_ce
+ Disable features for corrected errors, e.g. polling timer
+ and CMCI. All events reported as corrected are not cleared
+ by OS and remained in its error banks.
+ Usually this disablement is not recommended, however if
+ there is an agent checking/clearing corrected errors
+ (e.g. BIOS or hardware monitoring applications), conflicting
+ with OS's error handling, and you cannot deactivate the agent,
+ then this option will be a help.
+ mce=no_lmce
+ Do not opt-in to Local MCE delivery. Use legacy method
+ to broadcast MCEs.
+ mce=bootlog
+ Enable logging of machine checks left over from booting.
+ Disabled by default on AMD Fam10h and older because some BIOS
+ leave bogus ones.
+ If your BIOS doesn't do that it's a good idea to enable though
+ to make sure you log even machine check events that result
+ in a reboot. On Intel systems it is enabled by default.
+ mce=nobootlog
+ Disable boot machine check logging.
+ mce=monarchtimeout (number)
+ monarchtimeout:
+ Sets the time in us to wait for other CPUs on machine checks. 0
+ to disable.
+ mce=bios_cmci_threshold
+ Don't overwrite the bios-set CMCI threshold. This boot option
+ prevents Linux from overwriting the CMCI threshold set by the
+ bios. Without this option, Linux always sets the CMCI
+ threshold to 1. Enabling this may make memory predictive failure
+ analysis less effective if the bios sets thresholds for memory
+ errors since we will not see details for all errors.
+ mce=recovery
+ Force-enable recoverable machine check code paths
+
+ nomce (for compatibility with i386)
+ same as mce=off
+
+ Everything else is in sysfs now.
+
+APICs
+=====
+
+ apic
+ Use IO-APIC. Default
+
+ noapic
+ Don't use the IO-APIC.
+
+ disableapic
+ Don't use the local APIC
+
+ nolapic
+ Don't use the local APIC (alias for i386 compatibility)
+
+ pirq=...
+ See Documentation/x86/i386/IO-APIC.rst
+
+ noapictimer
+ Don't set up the APIC timer
+
+ no_timer_check
+ Don't check the IO-APIC timer. This can work around
+ problems with incorrect timer initialization on some boards.
+
+ apicpmtimer
+ Do APIC timer calibration using the pmtimer. Implies
+ apicmaintimer. Useful when your PIT timer is totally broken.
+
+Timing
+======
+
+ notsc
+ Deprecated, use tsc=unstable instead.
+
+ nohpet
+ Don't use the HPET timer.
+
+Idle loop
+=========
+
+ idle=poll
+ Don't do power saving in the idle loop using HLT, but poll for rescheduling
+ event. This will make the CPUs eat a lot more power, but may be useful
+ to get slightly better performance in multiprocessor benchmarks. It also
+ makes some profiling using performance counters more accurate.
+ Please note that on systems with MONITOR/MWAIT support (like Intel EM64T
+ CPUs) this option has no performance advantage over the normal idle loop.
+ It may also interact badly with hyperthreading.
+
+Rebooting
+=========
+
+ reboot=b[ios] | t[riple] | k[bd] | a[cpi] | e[fi] | p[ci] [, [w]arm | [c]old]
+ bios
+ Use the CPU reboot vector for warm reset
+ warm
+ Don't set the cold reboot flag
+ cold
+ Set the cold reboot flag
+ triple
+ Force a triple fault (init)
+ kbd
+ Use the keyboard controller. cold reset (default)
+ acpi
+ Use the ACPI RESET_REG in the FADT. If ACPI is not configured or
+ the ACPI reset does not work, the reboot path attempts the reset
+ using the keyboard controller.
+ efi
+ Use efi reset_system runtime service. If EFI is not configured or
+ the EFI reset does not work, the reboot path attempts the reset using
+ the keyboard controller.
+ pci
+ Use a write to the PCI config space register 0xcf9 to trigger reboot.
+
+ Using warm reset will be much faster especially on big memory
+ systems because the BIOS will not go through the memory check.
+ Disadvantage is that not all hardware will be completely reinitialized
+ on reboot so there may be boot problems on some systems.
+
+ reboot=force
+ Don't stop other CPUs on reboot. This can make reboot more reliable
+ in some cases.
+
+ reboot=default
+ There are some built-in platform specific "quirks" - you may see:
+ "reboot: <name> series board detected. Selecting <type> for reboots."
+ In the case where you think the quirk is in error (e.g. you have
+ newer BIOS, or newer board) using this option will ignore the built-in
+ quirk table, and use the generic default reboot actions.
+
+NUMA
+====
+
+ numa=off
+ Only set up a single NUMA node spanning all memory.
+
+ numa=noacpi
+ Don't parse the SRAT table for NUMA setup
+
+ numa=nohmat
+ Don't parse the HMAT table for NUMA setup, or soft-reserved memory
+ partitioning.
+
+ numa=fake=<size>[MG]
+ If given as a memory unit, fills all system RAM with nodes of
+ size interleaved over physical nodes.
+
+ numa=fake=<N>
+ If given as an integer, fills all system RAM with N fake nodes
+ interleaved over physical nodes.
+
+ numa=fake=<N>U
+ If given as an integer followed by 'U', it will divide each
+ physical node into N emulated nodes.
+
+ACPI
+====
+
+ acpi=off
+ Don't enable ACPI
+ acpi=ht
+ Use ACPI boot table parsing, but don't enable ACPI interpreter
+ acpi=force
+ Force ACPI on (currently not needed)
+ acpi=strict
+ Disable out of spec ACPI workarounds.
+ acpi_sci={edge,level,high,low}
+ Set up ACPI SCI interrupt.
+ acpi=noirq
+ Don't route interrupts
+ acpi=nocmcff
+ Disable firmware first mode for corrected errors. This
+ disables parsing the HEST CMC error source to check if
+ firmware has set the FF flag. This may result in
+ duplicate corrected error reports.
+
+PCI
+===
+
+ pci=off
+ Don't use PCI
+ pci=conf1
+ Use conf1 access.
+ pci=conf2
+ Use conf2 access.
+ pci=rom
+ Assign ROMs.
+ pci=assign-busses
+ Assign busses
+ pci=irqmask=MASK
+ Set PCI interrupt mask to MASK
+ pci=lastbus=NUMBER
+ Scan up to NUMBER busses, no matter what the mptable says.
+ pci=noacpi
+ Don't use ACPI to set up PCI interrupt routing.
+
+IOMMU (input/output memory management unit)
+===========================================
+Multiple x86-64 PCI-DMA mapping implementations exist, for example:
+
+ 1. <kernel/dma/direct.c>: use no hardware/software IOMMU at all
+ (e.g. because you have < 3 GB memory).
+ Kernel boot message: "PCI-DMA: Disabling IOMMU"
+
+ 2. <arch/x86/kernel/amd_gart_64.c>: AMD GART based hardware IOMMU.
+ Kernel boot message: "PCI-DMA: using GART IOMMU"
+
+ 3. <arch/x86_64/kernel/pci-swiotlb.c> : Software IOMMU implementation. Used
+ e.g. if there is no hardware IOMMU in the system and it is need because
+ you have >3GB memory or told the kernel to us it (iommu=soft))
+ Kernel boot message: "PCI-DMA: Using software bounce buffering
+ for IO (SWIOTLB)"
+
+::
+
+ iommu=[<size>][,noagp][,off][,force][,noforce]
+ [,memaper[=<order>]][,merge][,fullflush][,nomerge]
+ [,noaperture]
+
+General iommu options:
+
+ off
+ Don't initialize and use any kind of IOMMU.
+ noforce
+ Don't force hardware IOMMU usage when it is not needed. (default).
+ force
+ Force the use of the hardware IOMMU even when it is
+ not actually needed (e.g. because < 3 GB memory).
+ soft
+ Use software bounce buffering (SWIOTLB) (default for
+ Intel machines). This can be used to prevent the usage
+ of an available hardware IOMMU.
+
+iommu options only relevant to the AMD GART hardware IOMMU:
+
+ <size>
+ Set the size of the remapping area in bytes.
+ allowed
+ Overwrite iommu off workarounds for specific chipsets.
+ fullflush
+ Flush IOMMU on each allocation (default).
+ nofullflush
+ Don't use IOMMU fullflush.
+ memaper[=<order>]
+ Allocate an own aperture over RAM with size 32MB<<order.
+ (default: order=1, i.e. 64MB)
+ merge
+ Do scatter-gather (SG) merging. Implies "force" (experimental).
+ nomerge
+ Don't do scatter-gather (SG) merging.
+ noaperture
+ Ask the IOMMU not to touch the aperture for AGP.
+ noagp
+ Don't initialize the AGP driver and use full aperture.
+ panic
+ Always panic when IOMMU overflows.
+
+iommu options only relevant to the software bounce buffering (SWIOTLB) IOMMU
+implementation:
+
+ swiotlb=<slots>[,force,noforce]
+ <slots>
+ Prereserve that many 2K slots for the software IO bounce buffering.
+ force
+ Force all IO through the software TLB.
+ noforce
+ Do not initialize the software TLB.
+
+
+Miscellaneous
+=============
+
+ nogbpages
+ Do not use GB pages for kernel direct mappings.
+ gbpages
+ Use GB pages for kernel direct mappings.
+
+
+AMD SEV (Secure Encrypted Virtualization)
+=========================================
+Options relating to AMD SEV, specified via the following format:
+
+::
+
+ sev=option1[,option2]
+
+The available options are:
+
+ debug
+ Enable debug messages.
diff --git a/Documentation/x86/x86_64/cpu-hotplug-spec.rst b/Documentation/x86/x86_64/cpu-hotplug-spec.rst
new file mode 100644
index 000000000..8d1c91f0c
--- /dev/null
+++ b/Documentation/x86/x86_64/cpu-hotplug-spec.rst
@@ -0,0 +1,24 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===================================================
+Firmware support for CPU hotplug under Linux/x86-64
+===================================================
+
+Linux/x86-64 supports CPU hotplug now. For various reasons Linux wants to
+know in advance of boot time the maximum number of CPUs that could be plugged
+into the system. ACPI 3.0 currently has no official way to supply
+this information from the firmware to the operating system.
+
+In ACPI each CPU needs an LAPIC object in the MADT table (5.2.11.5 in the
+ACPI 3.0 specification). ACPI already has the concept of disabled LAPIC
+objects by setting the Enabled bit in the LAPIC object to zero.
+
+For CPU hotplug Linux/x86-64 expects now that any possible future hotpluggable
+CPU is already available in the MADT. If the CPU is not available yet
+it should have its LAPIC Enabled bit set to 0. Linux will use the number
+of disabled LAPICs to compute the maximum number of future CPUs.
+
+In the worst case the user can overwrite this choice using a command line
+option (additional_cpus=...), but it is recommended to supply the correct
+number (or a reasonable approximation of it, with erring towards more not less)
+in the MADT to avoid manual configuration.
diff --git a/Documentation/x86/x86_64/fake-numa-for-cpusets.rst b/Documentation/x86/x86_64/fake-numa-for-cpusets.rst
new file mode 100644
index 000000000..ff9bcfd2c
--- /dev/null
+++ b/Documentation/x86/x86_64/fake-numa-for-cpusets.rst
@@ -0,0 +1,78 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=====================
+Fake NUMA For CPUSets
+=====================
+
+:Author: David Rientjes <rientjes@cs.washington.edu>
+
+Using numa=fake and CPUSets for Resource Management
+
+This document describes how the numa=fake x86_64 command-line option can be used
+in conjunction with cpusets for coarse memory management. Using this feature,
+you can create fake NUMA nodes that represent contiguous chunks of memory and
+assign them to cpusets and their attached tasks. This is a way of limiting the
+amount of system memory that are available to a certain class of tasks.
+
+For more information on the features of cpusets, see
+Documentation/admin-guide/cgroup-v1/cpusets.rst.
+There are a number of different configurations you can use for your needs. For
+more information on the numa=fake command line option and its various ways of
+configuring fake nodes, see Documentation/x86/x86_64/boot-options.rst.
+
+For the purposes of this introduction, we'll assume a very primitive NUMA
+emulation setup of "numa=fake=4*512,". This will split our system memory into
+four equal chunks of 512M each that we can now use to assign to cpusets. As
+you become more familiar with using this combination for resource control,
+you'll determine a better setup to minimize the number of nodes you have to deal
+with.
+
+A machine may be split as follows with "numa=fake=4*512," as reported by dmesg::
+
+ Faking node 0 at 0000000000000000-0000000020000000 (512MB)
+ Faking node 1 at 0000000020000000-0000000040000000 (512MB)
+ Faking node 2 at 0000000040000000-0000000060000000 (512MB)
+ Faking node 3 at 0000000060000000-0000000080000000 (512MB)
+ ...
+ On node 0 totalpages: 130975
+ On node 1 totalpages: 131072
+ On node 2 totalpages: 131072
+ On node 3 totalpages: 131072
+
+Now following the instructions for mounting the cpusets filesystem from
+Documentation/admin-guide/cgroup-v1/cpusets.rst, you can assign fake nodes (i.e. contiguous memory
+address spaces) to individual cpusets::
+
+ [root@xroads /]# mkdir exampleset
+ [root@xroads /]# mount -t cpuset none exampleset
+ [root@xroads /]# mkdir exampleset/ddset
+ [root@xroads /]# cd exampleset/ddset
+ [root@xroads /exampleset/ddset]# echo 0-1 > cpus
+ [root@xroads /exampleset/ddset]# echo 0-1 > mems
+
+Now this cpuset, 'ddset', will only allowed access to fake nodes 0 and 1 for
+memory allocations (1G).
+
+You can now assign tasks to these cpusets to limit the memory resources
+available to them according to the fake nodes assigned as mems::
+
+ [root@xroads /exampleset/ddset]# echo $$ > tasks
+ [root@xroads /exampleset/ddset]# dd if=/dev/zero of=tmp bs=1024 count=1G
+ [1] 13425
+
+Notice the difference between the system memory usage as reported by
+/proc/meminfo between the restricted cpuset case above and the unrestricted
+case (i.e. running the same 'dd' command without assigning it to a fake NUMA
+cpuset):
+
+ ======== ============ ==========
+ Name Unrestricted Restricted
+ ======== ============ ==========
+ MemTotal 3091900 kB 3091900 kB
+ MemFree 42113 kB 1513236 kB
+ ======== ============ ==========
+
+This allows for coarse memory management for the tasks you assign to particular
+cpusets. Since cpusets can form a hierarchy, you can create some pretty
+interesting combinations of use-cases for various classes of tasks for your
+memory management needs.
diff --git a/Documentation/x86/x86_64/fsgs.rst b/Documentation/x86/x86_64/fsgs.rst
new file mode 100644
index 000000000..50960e09e
--- /dev/null
+++ b/Documentation/x86/x86_64/fsgs.rst
@@ -0,0 +1,199 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Using FS and GS segments in user space applications
+===================================================
+
+The x86 architecture supports segmentation. Instructions which access
+memory can use segment register based addressing mode. The following
+notation is used to address a byte within a segment:
+
+ Segment-register:Byte-address
+
+The segment base address is added to the Byte-address to compute the
+resulting virtual address which is accessed. This allows to access multiple
+instances of data with the identical Byte-address, i.e. the same code. The
+selection of a particular instance is purely based on the base-address in
+the segment register.
+
+In 32-bit mode the CPU provides 6 segments, which also support segment
+limits. The limits can be used to enforce address space protections.
+
+In 64-bit mode the CS/SS/DS/ES segments are ignored and the base address is
+always 0 to provide a full 64bit address space. The FS and GS segments are
+still functional in 64-bit mode.
+
+Common FS and GS usage
+------------------------------
+
+The FS segment is commonly used to address Thread Local Storage (TLS). FS
+is usually managed by runtime code or a threading library. Variables
+declared with the '__thread' storage class specifier are instantiated per
+thread and the compiler emits the FS: address prefix for accesses to these
+variables. Each thread has its own FS base address so common code can be
+used without complex address offset calculations to access the per thread
+instances. Applications should not use FS for other purposes when they use
+runtimes or threading libraries which manage the per thread FS.
+
+The GS segment has no common use and can be used freely by
+applications. GCC and Clang support GS based addressing via address space
+identifiers.
+
+Reading and writing the FS/GS base address
+------------------------------------------
+
+There exist two mechanisms to read and write the FS/GS base address:
+
+ - the arch_prctl() system call
+
+ - the FSGSBASE instruction family
+
+Accessing FS/GS base with arch_prctl()
+--------------------------------------
+
+ The arch_prctl(2) based mechanism is available on all 64-bit CPUs and all
+ kernel versions.
+
+ Reading the base:
+
+ arch_prctl(ARCH_GET_FS, &fsbase);
+ arch_prctl(ARCH_GET_GS, &gsbase);
+
+ Writing the base:
+
+ arch_prctl(ARCH_SET_FS, fsbase);
+ arch_prctl(ARCH_SET_GS, gsbase);
+
+ The ARCH_SET_GS prctl may be disabled depending on kernel configuration
+ and security settings.
+
+Accessing FS/GS base with the FSGSBASE instructions
+---------------------------------------------------
+
+ With the Ivy Bridge CPU generation Intel introduced a new set of
+ instructions to access the FS and GS base registers directly from user
+ space. These instructions are also supported on AMD Family 17H CPUs. The
+ following instructions are available:
+
+ =============== ===========================
+ RDFSBASE %reg Read the FS base register
+ RDGSBASE %reg Read the GS base register
+ WRFSBASE %reg Write the FS base register
+ WRGSBASE %reg Write the GS base register
+ =============== ===========================
+
+ The instructions avoid the overhead of the arch_prctl() syscall and allow
+ more flexible usage of the FS/GS addressing modes in user space
+ applications. This does not prevent conflicts between threading libraries
+ and runtimes which utilize FS and applications which want to use it for
+ their own purpose.
+
+FSGSBASE instructions enablement
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+ The instructions are enumerated in CPUID leaf 7, bit 0 of EBX. If
+ available /proc/cpuinfo shows 'fsgsbase' in the flag entry of the CPUs.
+
+ The availability of the instructions does not enable them
+ automatically. The kernel has to enable them explicitly in CR4. The
+ reason for this is that older kernels make assumptions about the values in
+ the GS register and enforce them when GS base is set via
+ arch_prctl(). Allowing user space to write arbitrary values to GS base
+ would violate these assumptions and cause malfunction.
+
+ On kernels which do not enable FSGSBASE the execution of the FSGSBASE
+ instructions will fault with a #UD exception.
+
+ The kernel provides reliable information about the enabled state in the
+ ELF AUX vector. If the HWCAP2_FSGSBASE bit is set in the AUX vector, the
+ kernel has FSGSBASE instructions enabled and applications can use them.
+ The following code example shows how this detection works::
+
+ #include <sys/auxv.h>
+ #include <elf.h>
+
+ /* Will be eventually in asm/hwcap.h */
+ #ifndef HWCAP2_FSGSBASE
+ #define HWCAP2_FSGSBASE (1 << 1)
+ #endif
+
+ ....
+
+ unsigned val = getauxval(AT_HWCAP2);
+
+ if (val & HWCAP2_FSGSBASE)
+ printf("FSGSBASE enabled\n");
+
+FSGSBASE instructions compiler support
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+GCC version 4.6.4 and newer provide instrinsics for the FSGSBASE
+instructions. Clang 5 supports them as well.
+
+ =================== ===========================
+ _readfsbase_u64() Read the FS base register
+ _readfsbase_u64() Read the GS base register
+ _writefsbase_u64() Write the FS base register
+ _writegsbase_u64() Write the GS base register
+ =================== ===========================
+
+To utilize these instrinsics <immintrin.h> must be included in the source
+code and the compiler option -mfsgsbase has to be added.
+
+Compiler support for FS/GS based addressing
+-------------------------------------------
+
+GCC version 6 and newer provide support for FS/GS based addressing via
+Named Address Spaces. GCC implements the following address space
+identifiers for x86:
+
+ ========= ====================================
+ __seg_fs Variable is addressed relative to FS
+ __seg_gs Variable is addressed relative to GS
+ ========= ====================================
+
+The preprocessor symbols __SEG_FS and __SEG_GS are defined when these
+address spaces are supported. Code which implements fallback modes should
+check whether these symbols are defined. Usage example::
+
+ #ifdef __SEG_GS
+
+ long data0 = 0;
+ long data1 = 1;
+
+ long __seg_gs *ptr;
+
+ /* Check whether FSGSBASE is enabled by the kernel (HWCAP2_FSGSBASE) */
+ ....
+
+ /* Set GS base to point to data0 */
+ _writegsbase_u64(&data0);
+
+ /* Access offset 0 of GS */
+ ptr = 0;
+ printf("data0 = %ld\n", *ptr);
+
+ /* Set GS base to point to data1 */
+ _writegsbase_u64(&data1);
+ /* ptr still addresses offset 0! */
+ printf("data1 = %ld\n", *ptr);
+
+
+Clang does not provide the GCC address space identifiers, but it provides
+address spaces via an attribute based mechanism in Clang 2.6 and newer
+versions:
+
+ ==================================== =====================================
+ __attribute__((address_space(256)) Variable is addressed relative to GS
+ __attribute__((address_space(257)) Variable is addressed relative to FS
+ ==================================== =====================================
+
+FS/GS based addressing with inline assembly
+-------------------------------------------
+
+In case the compiler does not support address spaces, inline assembly can
+be used for FS/GS based addressing mode::
+
+ mov %fs:offset, %reg
+ mov %gs:offset, %reg
+
+ mov %reg, %fs:offset
+ mov %reg, %gs:offset
diff --git a/Documentation/x86/x86_64/index.rst b/Documentation/x86/x86_64/index.rst
new file mode 100644
index 000000000..a56070fc8
--- /dev/null
+++ b/Documentation/x86/x86_64/index.rst
@@ -0,0 +1,17 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==============
+x86_64 Support
+==============
+
+.. toctree::
+ :maxdepth: 2
+
+ boot-options
+ uefi
+ mm
+ 5level-paging
+ fake-numa-for-cpusets
+ cpu-hotplug-spec
+ machinecheck
+ fsgs
diff --git a/Documentation/x86/x86_64/machinecheck.rst b/Documentation/x86/x86_64/machinecheck.rst
new file mode 100644
index 000000000..cea12ee97
--- /dev/null
+++ b/Documentation/x86/x86_64/machinecheck.rst
@@ -0,0 +1,33 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===============================================================
+Configurable sysfs parameters for the x86-64 machine check code
+===============================================================
+
+Machine checks report internal hardware error conditions detected
+by the CPU. Uncorrected errors typically cause a machine check
+(often with panic), corrected ones cause a machine check log entry.
+
+Machine checks are organized in banks (normally associated with
+a hardware subsystem) and subevents in a bank. The exact meaning
+of the banks and subevent is CPU specific.
+
+mcelog knows how to decode them.
+
+When you see the "Machine check errors logged" message in the system
+log then mcelog should run to collect and decode machine check entries
+from /dev/mcelog. Normally mcelog should be run regularly from a cronjob.
+
+Each CPU has a directory in /sys/devices/system/machinecheck/machinecheckN
+(N = CPU number).
+
+The directory contains some configurable entries. See
+Documentation/ABI/testing/sysfs-mce for more details.
+
+TBD document entries for AMD threshold interrupt configuration
+
+For more details about the x86 machine check architecture
+see the Intel and AMD architecture manuals from their developer websites.
+
+For more details about the architecture
+see http://one.firstfloor.org/~andi/mce.pdf
diff --git a/Documentation/x86/x86_64/mm.rst b/Documentation/x86/x86_64/mm.rst
new file mode 100644
index 000000000..9798676bb
--- /dev/null
+++ b/Documentation/x86/x86_64/mm.rst
@@ -0,0 +1,157 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=================
+Memory Management
+=================
+
+Complete virtual memory map with 4-level page tables
+====================================================
+
+.. note::
+
+ - Negative addresses such as "-23 TB" are absolute addresses in bytes, counted down
+ from the top of the 64-bit address space. It's easier to understand the layout
+ when seen both in absolute addresses and in distance-from-top notation.
+
+ For example 0xffffe90000000000 == -23 TB, it's 23 TB lower than the top of the
+ 64-bit address space (ffffffffffffffff).
+
+ Note that as we get closer to the top of the address space, the notation changes
+ from TB to GB and then MB/KB.
+
+ - "16M TB" might look weird at first sight, but it's an easier way to visualize size
+ notation than "16 EB", which few will recognize at first sight as 16 exabytes.
+ It also shows it nicely how incredibly large 64-bit address space is.
+
+::
+
+ ========================================================================================================================
+ Start addr | Offset | End addr | Size | VM area description
+ ========================================================================================================================
+ | | | |
+ 0000000000000000 | 0 | 00007fffffffffff | 128 TB | user-space virtual memory, different per mm
+ __________________|____________|__________________|_________|___________________________________________________________
+ | | | |
+ 0000800000000000 | +128 TB | ffff7fffffffffff | ~16M TB | ... huge, almost 64 bits wide hole of non-canonical
+ | | | | virtual memory addresses up to the -128 TB
+ | | | | starting offset of kernel mappings.
+ __________________|____________|__________________|_________|___________________________________________________________
+ |
+ | Kernel-space virtual memory, shared between all processes:
+ ____________________________________________________________|___________________________________________________________
+ | | | |
+ ffff800000000000 | -128 TB | ffff87ffffffffff | 8 TB | ... guard hole, also reserved for hypervisor
+ ffff880000000000 | -120 TB | ffff887fffffffff | 0.5 TB | LDT remap for PTI
+ ffff888000000000 | -119.5 TB | ffffc87fffffffff | 64 TB | direct mapping of all physical memory (page_offset_base)
+ ffffc88000000000 | -55.5 TB | ffffc8ffffffffff | 0.5 TB | ... unused hole
+ ffffc90000000000 | -55 TB | ffffe8ffffffffff | 32 TB | vmalloc/ioremap space (vmalloc_base)
+ ffffe90000000000 | -23 TB | ffffe9ffffffffff | 1 TB | ... unused hole
+ ffffea0000000000 | -22 TB | ffffeaffffffffff | 1 TB | virtual memory map (vmemmap_base)
+ ffffeb0000000000 | -21 TB | ffffebffffffffff | 1 TB | ... unused hole
+ ffffec0000000000 | -20 TB | fffffbffffffffff | 16 TB | KASAN shadow memory
+ __________________|____________|__________________|_________|____________________________________________________________
+ |
+ | Identical layout to the 56-bit one from here on:
+ ____________________________________________________________|____________________________________________________________
+ | | | |
+ fffffc0000000000 | -4 TB | fffffdffffffffff | 2 TB | ... unused hole
+ | | | | vaddr_end for KASLR
+ fffffe0000000000 | -2 TB | fffffe7fffffffff | 0.5 TB | cpu_entry_area mapping
+ fffffe8000000000 | -1.5 TB | fffffeffffffffff | 0.5 TB | ... unused hole
+ ffffff0000000000 | -1 TB | ffffff7fffffffff | 0.5 TB | %esp fixup stacks
+ ffffff8000000000 | -512 GB | ffffffeeffffffff | 444 GB | ... unused hole
+ ffffffef00000000 | -68 GB | fffffffeffffffff | 64 GB | EFI region mapping space
+ ffffffff00000000 | -4 GB | ffffffff7fffffff | 2 GB | ... unused hole
+ ffffffff80000000 | -2 GB | ffffffff9fffffff | 512 MB | kernel text mapping, mapped to physical address 0
+ ffffffff80000000 |-2048 MB | | |
+ ffffffffa0000000 |-1536 MB | fffffffffeffffff | 1520 MB | module mapping space
+ ffffffffff000000 | -16 MB | | |
+ FIXADDR_START | ~-11 MB | ffffffffff5fffff | ~0.5 MB | kernel-internal fixmap range, variable size and offset
+ ffffffffff600000 | -10 MB | ffffffffff600fff | 4 kB | legacy vsyscall ABI
+ ffffffffffe00000 | -2 MB | ffffffffffffffff | 2 MB | ... unused hole
+ __________________|____________|__________________|_________|___________________________________________________________
+
+
+Complete virtual memory map with 5-level page tables
+====================================================
+
+.. note::
+
+ - With 56-bit addresses, user-space memory gets expanded by a factor of 512x,
+ from 0.125 PB to 64 PB. All kernel mappings shift down to the -64 PB starting
+ offset and many of the regions expand to support the much larger physical
+ memory supported.
+
+::
+
+ ========================================================================================================================
+ Start addr | Offset | End addr | Size | VM area description
+ ========================================================================================================================
+ | | | |
+ 0000000000000000 | 0 | 00ffffffffffffff | 64 PB | user-space virtual memory, different per mm
+ __________________|____________|__________________|_________|___________________________________________________________
+ | | | |
+ 0100000000000000 | +64 PB | feffffffffffffff | ~16K PB | ... huge, still almost 64 bits wide hole of non-canonical
+ | | | | virtual memory addresses up to the -64 PB
+ | | | | starting offset of kernel mappings.
+ __________________|____________|__________________|_________|___________________________________________________________
+ |
+ | Kernel-space virtual memory, shared between all processes:
+ ____________________________________________________________|___________________________________________________________
+ | | | |
+ ff00000000000000 | -64 PB | ff0fffffffffffff | 4 PB | ... guard hole, also reserved for hypervisor
+ ff10000000000000 | -60 PB | ff10ffffffffffff | 0.25 PB | LDT remap for PTI
+ ff11000000000000 | -59.75 PB | ff90ffffffffffff | 32 PB | direct mapping of all physical memory (page_offset_base)
+ ff91000000000000 | -27.75 PB | ff9fffffffffffff | 3.75 PB | ... unused hole
+ ffa0000000000000 | -24 PB | ffd1ffffffffffff | 12.5 PB | vmalloc/ioremap space (vmalloc_base)
+ ffd2000000000000 | -11.5 PB | ffd3ffffffffffff | 0.5 PB | ... unused hole
+ ffd4000000000000 | -11 PB | ffd5ffffffffffff | 0.5 PB | virtual memory map (vmemmap_base)
+ ffd6000000000000 | -10.5 PB | ffdeffffffffffff | 2.25 PB | ... unused hole
+ ffdf000000000000 | -8.25 PB | fffffbffffffffff | ~8 PB | KASAN shadow memory
+ __________________|____________|__________________|_________|____________________________________________________________
+ |
+ | Identical layout to the 47-bit one from here on:
+ ____________________________________________________________|____________________________________________________________
+ | | | |
+ fffffc0000000000 | -4 TB | fffffdffffffffff | 2 TB | ... unused hole
+ | | | | vaddr_end for KASLR
+ fffffe0000000000 | -2 TB | fffffe7fffffffff | 0.5 TB | cpu_entry_area mapping
+ fffffe8000000000 | -1.5 TB | fffffeffffffffff | 0.5 TB | ... unused hole
+ ffffff0000000000 | -1 TB | ffffff7fffffffff | 0.5 TB | %esp fixup stacks
+ ffffff8000000000 | -512 GB | ffffffeeffffffff | 444 GB | ... unused hole
+ ffffffef00000000 | -68 GB | fffffffeffffffff | 64 GB | EFI region mapping space
+ ffffffff00000000 | -4 GB | ffffffff7fffffff | 2 GB | ... unused hole
+ ffffffff80000000 | -2 GB | ffffffff9fffffff | 512 MB | kernel text mapping, mapped to physical address 0
+ ffffffff80000000 |-2048 MB | | |
+ ffffffffa0000000 |-1536 MB | fffffffffeffffff | 1520 MB | module mapping space
+ ffffffffff000000 | -16 MB | | |
+ FIXADDR_START | ~-11 MB | ffffffffff5fffff | ~0.5 MB | kernel-internal fixmap range, variable size and offset
+ ffffffffff600000 | -10 MB | ffffffffff600fff | 4 kB | legacy vsyscall ABI
+ ffffffffffe00000 | -2 MB | ffffffffffffffff | 2 MB | ... unused hole
+ __________________|____________|__________________|_________|___________________________________________________________
+
+Architecture defines a 64-bit virtual address. Implementations can support
+less. Currently supported are 48- and 57-bit virtual addresses. Bits 63
+through to the most-significant implemented bit are sign extended.
+This causes hole between user space and kernel addresses if you interpret them
+as unsigned.
+
+The direct mapping covers all memory in the system up to the highest
+memory address (this means in some cases it can also include PCI memory
+holes).
+
+We map EFI runtime services in the 'efi_pgd' PGD in a 64Gb large virtual
+memory window (this size is arbitrary, it can be raised later if needed).
+The mappings are not part of any other kernel PGD and are only available
+during EFI runtime calls.
+
+Note that if CONFIG_RANDOMIZE_MEMORY is enabled, the direct mapping of all
+physical memory, vmalloc/ioremap space and virtual memory map are randomized.
+Their order is preserved but their base will be offset early at boot time.
+
+Be very careful vs. KASLR when changing anything here. The KASLR address
+range must not overlap with anything except the KASAN shadow area, which is
+correct as KASAN disables KASLR.
+
+For both 4- and 5-level layouts, the STACKLEAK_POISON value in the last 2MB
+hole: ffffffffffff4111
diff --git a/Documentation/x86/x86_64/uefi.rst b/Documentation/x86/x86_64/uefi.rst
new file mode 100644
index 000000000..fbc30c9a0
--- /dev/null
+++ b/Documentation/x86/x86_64/uefi.rst
@@ -0,0 +1,58 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=====================================
+General note on [U]EFI x86_64 support
+=====================================
+
+The nomenclature EFI and UEFI are used interchangeably in this document.
+
+Although the tools below are _not_ needed for building the kernel,
+the needed bootloader support and associated tools for x86_64 platforms
+with EFI firmware and specifications are listed below.
+
+1. UEFI specification: http://www.uefi.org
+
+2. Booting Linux kernel on UEFI x86_64 platform requires bootloader
+ support. Elilo with x86_64 support can be used.
+
+3. x86_64 platform with EFI/UEFI firmware.
+
+Mechanics
+---------
+
+- Build the kernel with the following configuration::
+
+ CONFIG_FB_EFI=y
+ CONFIG_FRAMEBUFFER_CONSOLE=y
+
+ If EFI runtime services are expected, the following configuration should
+ be selected::
+
+ CONFIG_EFI=y
+ CONFIG_EFIVAR_FS=y or m # optional
+
+- Create a VFAT partition on the disk
+- Copy the following to the VFAT partition:
+
+ elilo bootloader with x86_64 support, elilo configuration file,
+ kernel image built in first step and corresponding
+ initrd. Instructions on building elilo and its dependencies
+ can be found in the elilo sourceforge project.
+
+- Boot to EFI shell and invoke elilo choosing the kernel image built
+ in first step.
+- If some or all EFI runtime services don't work, you can try following
+ kernel command line parameters to turn off some or all EFI runtime
+ services.
+
+ noefi
+ turn off all EFI runtime services
+ reboot_type=k
+ turn off EFI reboot runtime service
+
+- If the EFI memory map has additional entries not in the E820 map,
+ you can include those entries in the kernels memory map of available
+ physical RAM by using the following kernel command line parameter.
+
+ add_efi_memmap
+ include EFI memory map of available physical RAM
diff --git a/Documentation/x86/xstate.rst b/Documentation/x86/xstate.rst
new file mode 100644
index 000000000..5cec7fb55
--- /dev/null
+++ b/Documentation/x86/xstate.rst
@@ -0,0 +1,74 @@
+Using XSTATE features in user space applications
+================================================
+
+The x86 architecture supports floating-point extensions which are
+enumerated via CPUID. Applications consult CPUID and use XGETBV to
+evaluate which features have been enabled by the kernel XCR0.
+
+Up to AVX-512 and PKRU states, these features are automatically enabled by
+the kernel if available. Features like AMX TILE_DATA (XSTATE component 18)
+are enabled by XCR0 as well, but the first use of related instruction is
+trapped by the kernel because by default the required large XSTATE buffers
+are not allocated automatically.
+
+Using dynamically enabled XSTATE features in user space applications
+--------------------------------------------------------------------
+
+The kernel provides an arch_prctl(2) based mechanism for applications to
+request the usage of such features. The arch_prctl(2) options related to
+this are:
+
+-ARCH_GET_XCOMP_SUPP
+
+ arch_prctl(ARCH_GET_XCOMP_SUPP, &features);
+
+ ARCH_GET_XCOMP_SUPP stores the supported features in userspace storage of
+ type uint64_t. The second argument is a pointer to that storage.
+
+-ARCH_GET_XCOMP_PERM
+
+ arch_prctl(ARCH_GET_XCOMP_PERM, &features);
+
+ ARCH_GET_XCOMP_PERM stores the features for which the userspace process
+ has permission in userspace storage of type uint64_t. The second argument
+ is a pointer to that storage.
+
+-ARCH_REQ_XCOMP_PERM
+
+ arch_prctl(ARCH_REQ_XCOMP_PERM, feature_nr);
+
+ ARCH_REQ_XCOMP_PERM allows to request permission for a dynamically enabled
+ feature or a feature set. A feature set can be mapped to a facility, e.g.
+ AMX, and can require one or more XSTATE components to be enabled.
+
+ The feature argument is the number of the highest XSTATE component which
+ is required for a facility to work.
+
+When requesting permission for a feature, the kernel checks the
+availability. The kernel ensures that sigaltstacks in the process's tasks
+are large enough to accommodate the resulting large signal frame. It
+enforces this both during ARCH_REQ_XCOMP_SUPP and during any subsequent
+sigaltstack(2) calls. If an installed sigaltstack is smaller than the
+resulting sigframe size, ARCH_REQ_XCOMP_SUPP results in -ENOSUPP. Also,
+sigaltstack(2) results in -ENOMEM if the requested altstack is too small
+for the permitted features.
+
+Permission, when granted, is valid per process. Permissions are inherited
+on fork(2) and cleared on exec(3).
+
+The first use of an instruction related to a dynamically enabled feature is
+trapped by the kernel. The trap handler checks whether the process has
+permission to use the feature. If the process has no permission then the
+kernel sends SIGILL to the application. If the process has permission then
+the handler allocates a larger xstate buffer for the task so the large
+state can be context switched. In the unlikely cases that the allocation
+fails, the kernel sends SIGSEGV.
+
+Dynamic features in signal frames
+---------------------------------
+
+Dynamcally enabled features are not written to the signal frame upon signal
+entry if the feature is in its initial configuration. This differs from
+non-dynamic features which are always written regardless of their
+configuration. Signal handlers can examine the XSAVE buffer's XSTATE_BV
+field to determine if a features was written.
diff --git a/Documentation/x86/zero-page.rst b/Documentation/x86/zero-page.rst
new file mode 100644
index 000000000..45aa9cceb
--- /dev/null
+++ b/Documentation/x86/zero-page.rst
@@ -0,0 +1,47 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=========
+Zero Page
+=========
+The additional fields in struct boot_params as a part of 32-bit boot
+protocol of kernel. These should be filled by bootloader or 16-bit
+real-mode setup code of the kernel. References/settings to it mainly
+are in::
+
+ arch/x86/include/uapi/asm/bootparam.h
+
+=========== ===== ======================= =================================================
+Offset/Size Proto Name Meaning
+
+000/040 ALL screen_info Text mode or frame buffer information
+ (struct screen_info)
+040/014 ALL apm_bios_info APM BIOS information (struct apm_bios_info)
+058/008 ALL tboot_addr Physical address of tboot shared page
+060/010 ALL ist_info Intel SpeedStep (IST) BIOS support information
+ (struct ist_info)
+070/008 ALL acpi_rsdp_addr Physical address of ACPI RSDP table
+080/010 ALL hd0_info hd0 disk parameter, OBSOLETE!!
+090/010 ALL hd1_info hd1 disk parameter, OBSOLETE!!
+0A0/010 ALL sys_desc_table System description table (struct sys_desc_table),
+ OBSOLETE!!
+0B0/010 ALL olpc_ofw_header OLPC's OpenFirmware CIF and friends
+0C0/004 ALL ext_ramdisk_image ramdisk_image high 32bits
+0C4/004 ALL ext_ramdisk_size ramdisk_size high 32bits
+0C8/004 ALL ext_cmd_line_ptr cmd_line_ptr high 32bits
+13C/004 ALL cc_blob_address Physical address of Confidential Computing blob
+140/080 ALL edid_info Video mode setup (struct edid_info)
+1C0/020 ALL efi_info EFI 32 information (struct efi_info)
+1E0/004 ALL alt_mem_k Alternative mem check, in KB
+1E4/004 ALL scratch Scratch field for the kernel setup code
+1E8/001 ALL e820_entries Number of entries in e820_table (below)
+1E9/001 ALL eddbuf_entries Number of entries in eddbuf (below)
+1EA/001 ALL edd_mbr_sig_buf_entries Number of entries in edd_mbr_sig_buffer
+ (below)
+1EB/001 ALL kbd_status Numlock is enabled
+1EC/001 ALL secure_boot Secure boot is enabled in the firmware
+1EF/001 ALL sentinel Used to detect broken bootloaders
+290/040 ALL edd_mbr_sig_buffer EDD MBR signatures
+2D0/A00 ALL e820_table E820 memory map table
+ (array of struct e820_entry)
+D00/1EC ALL eddbuf EDD data (array of struct edd_info)
+=========== ===== ======================= =================================================