Paged virtual memory was invented along with virtual memory as a concept in 1962 on the Ferranti Atlas Computer which was the first computer with paged virtual memory. The feature migrated to newer computers and became a de facto feature of all Unix-like systems as time went by. In 1985 the feature was included in the Intel 80386, which was the CPU Linux 1.0 was developed on.
Page tables map virtual addresses as seen by the CPU into physical addresses as seen on the external memory bus.
Linux defines page tables as a hierarchy which is currently five levels in height. The architecture code for each supported architecture will then map this to the restrictions of the hardware.
The physical address corresponding to the virtual address is often referenced by the underlying physical page frame. The page frame number or pfn is the physical address of the page (as seen on the external memory bus) divided by PAGE_SIZE.
Physical memory address 0 will be pfn 0 and the highest pfn will be the last page of physical memory the external address bus of the CPU can address.
With a page granularity of 4KB and a address range of 32 bits, pfn 0 is at address 0x00000000, pfn 1 is at address 0x00001000, pfn 2 is at 0x00002000 and so on until we reach pfn 0xfffff at 0xfffff000. With 16KB pages pfs are at 0x00004000, 0x00008000 ... 0xffffc000 and pfn goes from 0 to 0x3fffff.
As you can see, with 4KB pages the page base address uses bits 12-31 of the address, and this is why PAGE_SHIFT in this case is defined as 12 and PAGE_SIZE is usually defined in terms of the page shift as (1 << PAGE_SHIFT)
Over time a deeper hierarchy has been developed in response to increasing memory sizes. When Linux was created, 4KB pages and a single page table called swapper_pg_dir with 1024 entries was used, covering 4MB which coincided with the fact that Torvald's first computer had 4MB of physical memory. Entries in this single table were referred to as PTE:s - page table entries.
The software page table hierarchy reflects the fact that page table hardware has become hierarchical and that in turn is done to save page table memory and speed up mapping.
One could of course imagine a single, linear page table with enormous amounts of entries, breaking down the whole memory into single pages. Such a page table would be very sparse, because large portions of the virtual memory usually remains unused. By using hierarchical page tables large holes in the virtual address space does not waste valuable page table memory, because it will suffice to mark large areas as unmapped at a higher level in the page table hierarchy.
Additionally, on modern CPUs, a higher level page table entry can point directly to a physical memory range, which allows mapping a contiguous range of several megabytes or even gigabytes in a single high-level page table entry, taking shortcuts in mapping virtual memory to physical memory: there is no need to traverse deeper in the hierarchy when you find a large mapped range like this.
The page table hierarchy has now developed into this:
+-----+ | PGD | +-----+ | | +-----+ +-->| P4D | +-----+ | | +-----+ +-->| PUD | +-----+ | | +-----+ +-->| PMD | +-----+ | | +-----+ +-->| PTE | +-----+
Symbols on the different levels of the page table hierarchy have the following meaning beginning from the bottom:
pte, pte_t, pteval_t = Page Table Entry - mentioned earlier. The pte is an array of PTRS_PER_PTE elements of the pteval_t type, each mapping a single page of virtual memory to a single page of physical memory. The architecture defines the size and contents of pteval_t.
A typical example is that the pteval_t is a 32- or 64-bit value with the upper bits being a pfn (page frame number), and the lower bits being some architecture-specific bits such as memory protection.
The entry part of the name is a bit confusing because while in Linux 1.0 this did refer to a single page table entry in the single top level page table, it was retrofitted to be an array of mapping elements when two-level page tables were first introduced, so the pte is the lowermost page table, not a page table entry.
pmd, pmd_t, pmdval_t = Page Middle Directory, the hierarchy right above the pte, with PTRS_PER_PMD references to the pte:s.
pud, pud_t, pudval_t = Page Upper Directory was introduced after the other levels to handle 4-level page tables. It is potentially unused, or folded as we will discuss later.
p4d, p4d_t, p4dval_t = Page Level 4 Directory was introduced to handle 5-level page tables after the pud was introduced. Now it was clear that we needed to replace pgd, pmd, pud etc with a figure indicating the directory level and that we cannot go on with ad hoc names any more. This is only used on systems which actually have 5 levels of page tables, otherwise it is folded.
pgd, pgd_t, pgdval_t = Page Global Directory - the Linux kernel main page table handling the PGD for the kernel memory is still found in swapper_pg_dir, but each userspace process in the system also has its own memory context and thus its own pgd, found in struct mm_struct which in turn is referenced to in each struct task_struct. So tasks have memory context in the form of a struct mm_struct and this in turn has a struct pgt_t *pgd pointer to the corresponding page global directory.
To repeat: each level in the page table hierarchy is a array of pointers, so the pgd contains PTRS_PER_PGD pointers to the next level below, p4d contains PTRS_PER_P4D pointers to pud items and so on. The number of pointers on each level is architecture-defined.:
PMD --> +-----+ PTE | ptr |-------> +-----+ | ptr |- | ptr |-------> PAGE | ptr | \ | ptr | | ptr | \ ... | ... | \ | ptr | \ PTE +-----+ +----> +-----+ | ptr |-------> PAGE | ptr | ...
Page Table Folding¶
If the architecture does not use all the page table levels, they can be folded which means skipped, and all operations performed on page tables will be compile-time augmented to just skip a level when accessing the next lower level.
Page table handling code that wishes to be architecture-neutral, such as the virtual memory manager, will need to be written so that it traverses all of the currently five levels. This style should also be preferred for architecture-specific code, so as to be robust to future changes.
MMU, TLB, and Page Faults¶
The Memory Management Unit (MMU) is a hardware component that handles virtual to physical address translations. It may use relatively small caches in hardware called Translation Lookaside Buffers (TLBs) and Page Walk Caches to speed up these translations.
When CPU accesses a memory location, it provides a virtual address to the MMU, which checks if there is the existing translation in the TLB or in the Page Walk Caches (on architectures that support them). If no translation is found, MMU uses the page walks to determine the physical address and create the map.
The dirty bit for a page is set (i.e., turned on) when the page is written to. Each page of memory has associated permission and dirty bits. The latter indicate that the page has been modified since it was loaded into memory.
If nothing prevents it, eventually the physical memory can be accessed and the requested operation on the physical frame is performed.
There are several reasons why the MMU can't find certain translations. It could happen because the CPU is trying to access memory that the current task is not permitted to, or because the data is not present into physical memory.
When these conditions happen, the MMU triggers page faults, which are types of exceptions that signal the CPU to pause the current execution and run a special function to handle the mentioned exceptions.
There are common and expected causes of page faults. These are triggered by process management optimization techniques called "Lazy Allocation" and "Copy-on-Write". Page faults may also happen when frames have been swapped out to persistent storage (swap partition or file) and evicted from their physical locations.
These techniques improve memory efficiency, reduce latency, and minimize space occupation. This document won't go deeper into the details of "Lazy Allocation" and "Copy-on-Write" because these subjects are out of scope as they belong to Process Address Management.
Swapping differentiates itself from the other mentioned techniques because it's undesirable since it's performed as a means to reduce memory under heavy pressure.
Swapping can't work for memory mapped by kernel logical addresses. These are a subset of the kernel virtual space that directly maps a contiguous range of physical memory. Given any logical address, its physical address is determined with simple arithmetic on an offset. Accesses to logical addresses are fast because they avoid the need for complex page table lookups at the expenses of frames not being evictable and pageable out.
If the kernel fails to make room for the data that must be present in the physical frames, the kernel invokes the out-of-memory (OOM) killer to make room by terminating lower priority processes until pressure reduces under a safe threshold.
Additionally, page faults may be also caused by code bugs or by maliciously crafted addresses that the CPU is instructed to access. A thread of a process could use instructions to address (non-shared) memory which does not belong to its own address space, or could try to execute an instruction that want to write to a read-only location.
If the above-mentioned conditions happen in user-space, the kernel sends a Segmentation Fault (SIGSEGV) signal to the current thread. That signal usually causes the termination of the thread and of the process it belongs to.
This document is going to simplify and show an high altitude view of how the Linux kernel handles these page faults, creates tables and tables' entries, check if memory is present and, if not, requests to load data from persistent storage or from other devices, and updates the MMU and its caches.
The first steps are architecture dependent. Most architectures jump to do_page_fault(), whereas the x86 interrupt handler is defined by the DEFINE_IDTENTRY_RAW_ERRORCODE() macro which calls handle_page_fault().
Whatever the routes, all architectures end up to the invocation of handle_mm_fault() which, in turn, (likely) ends up calling __handle_mm_fault() to carry out the actual work of allocating the page tables.
The unfortunate case of not being able to call __handle_mm_fault() means that the virtual address is pointing to areas of physical memory which are not permitted to be accessed (at least from the current context). This condition resolves to the kernel sending the above-mentioned SIGSEGV signal to the process and leads to the consequences already explained.
__handle_mm_fault() carries out its work by calling several functions to find the entry's offsets of the upper layers of the page tables and allocate the tables that it may need.
The functions that look for the offset have names like *_offset(), where the "*" is for pgd, p4d, pud, pmd, pte; instead the functions to allocate the corresponding tables, layer by layer, are called *_alloc, using the above-mentioned convention to name them after the corresponding types of tables in the hierarchy.
The page table walk may end at one of the middle or upper layers (PMD, PUD).
Linux supports larger page sizes than the usual 4KB (i.e., the so called huge pages). When using these kinds of larger pages, higher level pages can directly map them, with no need to use lower level page entries (PTE). Huge pages contain large contiguous physical regions that usually span from 2MB to 1GB. They are respectively mapped by the PMD and PUD page entries.
The huge pages bring with them several benefits like reduced TLB pressure, reduced page table overhead, memory allocation efficiency, and performance improvement for certain workloads. However, these benefits come with trade-offs, like wasted memory and allocation challenges.
At the very end of the walk with allocations, if it didn't return errors, __handle_mm_fault() finally calls handle_pte_fault(), which via do_fault() performs one of do_read_fault(), do_cow_fault(), do_shared_fault(). "read", "cow", "shared" give hints about the reasons and the kind of fault it's handling.
The actual implementation of the workflow is very complex. Its design allows Linux to handle page faults in a way that is tailored to the specific characteristics of each architecture, while still sharing a common overall structure.
To conclude this high altitude view of how Linux handles page faults, let's add that the page faults handler can be disabled and enabled respectively with pagefault_disable() and pagefault_enable().
Several code path make use of the latter two functions because they need to disable traps into the page faults handler, mostly to prevent deadlocks.