OpenCloudOS-Kernel/arch/powerpc/kvm/book3s_64_mmu_radix.c

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// SPDX-License-Identifier: GPL-2.0-only
/*
*
* Copyright 2016 Paul Mackerras, IBM Corp. <paulus@au1.ibm.com>
*/
#include <linux/types.h>
#include <linux/string.h>
#include <linux/kvm.h>
#include <linux/kvm_host.h>
#include <linux/anon_inodes.h>
#include <linux/file.h>
#include <linux/debugfs.h>
#include <asm/kvm_ppc.h>
#include <asm/kvm_book3s.h>
#include <asm/page.h>
#include <asm/mmu.h>
#include <asm/pgtable.h>
#include <asm/pgalloc.h>
#include <asm/pte-walk.h>
/*
* Supported radix tree geometry.
* Like p9, we support either 5 or 9 bits at the first (lowest) level,
* for a page size of 64k or 4k.
*/
static int p9_supported_radix_bits[4] = { 5, 9, 9, 13 };
unsigned long __kvmhv_copy_tofrom_guest_radix(int lpid, int pid,
gva_t eaddr, void *to, void *from,
unsigned long n)
KVM: PPC: Book3S HV: Implement functions to access quadrants 1 & 2 The POWER9 radix mmu has the concept of quadrants. The quadrant number is the two high bits of the effective address and determines the fully qualified address to be used for the translation. The fully qualified address consists of the effective lpid, the effective pid and the effective address. This gives then 4 possible quadrants 0, 1, 2, and 3. When accessing these quadrants the fully qualified address is obtained as follows: Quadrant | Hypervisor | Guest -------------------------------------------------------------------------- | EA[0:1] = 0b00 | EA[0:1] = 0b00 0 | effLPID = 0 | effLPID = LPIDR | effPID = PIDR | effPID = PIDR -------------------------------------------------------------------------- | EA[0:1] = 0b01 | 1 | effLPID = LPIDR | Invalid Access | effPID = PIDR | -------------------------------------------------------------------------- | EA[0:1] = 0b10 | 2 | effLPID = LPIDR | Invalid Access | effPID = 0 | -------------------------------------------------------------------------- | EA[0:1] = 0b11 | EA[0:1] = 0b11 3 | effLPID = 0 | effLPID = LPIDR | effPID = 0 | effPID = 0 -------------------------------------------------------------------------- In the Guest; Quadrant 3 is normally used to address the operating system since this uses effPID=0 and effLPID=LPIDR, meaning the PID register doesn't need to be switched. Quadrant 0 is normally used to address user space since the effLPID and effPID are taken from the corresponding registers. In the Host; Quadrant 0 and 3 are used as above, however the effLPID is always 0 to address the host. Quadrants 1 and 2 can be used by the host to address guest memory using a guest effective address. Since the effLPID comes from the LPID register, the host loads the LPID of the guest it would like to access (and the PID of the process) and can perform accesses to a guest effective address. This means quadrant 1 can be used to address the guest user space and quadrant 2 can be used to address the guest operating system from the hypervisor, using a guest effective address. Access to the quadrants can cause a Hypervisor Data Storage Interrupt (HDSI) due to being unable to perform partition scoped translation. Previously this could only be generated from a guest and so the code path expects us to take the KVM trampoline in the interrupt handler. This is no longer the case so we modify the handler to call bad_page_fault() to check if we were expecting this fault so we can handle it gracefully and just return with an error code. In the hash mmu case we still raise an unknown exception since quadrants aren't defined for the hash mmu. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-14 13:29:05 +08:00
{
int uninitialized_var(old_pid), old_lpid;
KVM: PPC: Book3S HV: Implement functions to access quadrants 1 & 2 The POWER9 radix mmu has the concept of quadrants. The quadrant number is the two high bits of the effective address and determines the fully qualified address to be used for the translation. The fully qualified address consists of the effective lpid, the effective pid and the effective address. This gives then 4 possible quadrants 0, 1, 2, and 3. When accessing these quadrants the fully qualified address is obtained as follows: Quadrant | Hypervisor | Guest -------------------------------------------------------------------------- | EA[0:1] = 0b00 | EA[0:1] = 0b00 0 | effLPID = 0 | effLPID = LPIDR | effPID = PIDR | effPID = PIDR -------------------------------------------------------------------------- | EA[0:1] = 0b01 | 1 | effLPID = LPIDR | Invalid Access | effPID = PIDR | -------------------------------------------------------------------------- | EA[0:1] = 0b10 | 2 | effLPID = LPIDR | Invalid Access | effPID = 0 | -------------------------------------------------------------------------- | EA[0:1] = 0b11 | EA[0:1] = 0b11 3 | effLPID = 0 | effLPID = LPIDR | effPID = 0 | effPID = 0 -------------------------------------------------------------------------- In the Guest; Quadrant 3 is normally used to address the operating system since this uses effPID=0 and effLPID=LPIDR, meaning the PID register doesn't need to be switched. Quadrant 0 is normally used to address user space since the effLPID and effPID are taken from the corresponding registers. In the Host; Quadrant 0 and 3 are used as above, however the effLPID is always 0 to address the host. Quadrants 1 and 2 can be used by the host to address guest memory using a guest effective address. Since the effLPID comes from the LPID register, the host loads the LPID of the guest it would like to access (and the PID of the process) and can perform accesses to a guest effective address. This means quadrant 1 can be used to address the guest user space and quadrant 2 can be used to address the guest operating system from the hypervisor, using a guest effective address. Access to the quadrants can cause a Hypervisor Data Storage Interrupt (HDSI) due to being unable to perform partition scoped translation. Previously this could only be generated from a guest and so the code path expects us to take the KVM trampoline in the interrupt handler. This is no longer the case so we modify the handler to call bad_page_fault() to check if we were expecting this fault so we can handle it gracefully and just return with an error code. In the hash mmu case we still raise an unknown exception since quadrants aren't defined for the hash mmu. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-14 13:29:05 +08:00
unsigned long quadrant, ret = n;
bool is_load = !!to;
/* Can't access quadrants 1 or 2 in non-HV mode, call the HV to do it */
if (kvmhv_on_pseries())
return plpar_hcall_norets(H_COPY_TOFROM_GUEST, lpid, pid, eaddr,
(to != NULL) ? __pa(to): 0,
(from != NULL) ? __pa(from): 0, n);
KVM: PPC: Book3S HV: Implement functions to access quadrants 1 & 2 The POWER9 radix mmu has the concept of quadrants. The quadrant number is the two high bits of the effective address and determines the fully qualified address to be used for the translation. The fully qualified address consists of the effective lpid, the effective pid and the effective address. This gives then 4 possible quadrants 0, 1, 2, and 3. When accessing these quadrants the fully qualified address is obtained as follows: Quadrant | Hypervisor | Guest -------------------------------------------------------------------------- | EA[0:1] = 0b00 | EA[0:1] = 0b00 0 | effLPID = 0 | effLPID = LPIDR | effPID = PIDR | effPID = PIDR -------------------------------------------------------------------------- | EA[0:1] = 0b01 | 1 | effLPID = LPIDR | Invalid Access | effPID = PIDR | -------------------------------------------------------------------------- | EA[0:1] = 0b10 | 2 | effLPID = LPIDR | Invalid Access | effPID = 0 | -------------------------------------------------------------------------- | EA[0:1] = 0b11 | EA[0:1] = 0b11 3 | effLPID = 0 | effLPID = LPIDR | effPID = 0 | effPID = 0 -------------------------------------------------------------------------- In the Guest; Quadrant 3 is normally used to address the operating system since this uses effPID=0 and effLPID=LPIDR, meaning the PID register doesn't need to be switched. Quadrant 0 is normally used to address user space since the effLPID and effPID are taken from the corresponding registers. In the Host; Quadrant 0 and 3 are used as above, however the effLPID is always 0 to address the host. Quadrants 1 and 2 can be used by the host to address guest memory using a guest effective address. Since the effLPID comes from the LPID register, the host loads the LPID of the guest it would like to access (and the PID of the process) and can perform accesses to a guest effective address. This means quadrant 1 can be used to address the guest user space and quadrant 2 can be used to address the guest operating system from the hypervisor, using a guest effective address. Access to the quadrants can cause a Hypervisor Data Storage Interrupt (HDSI) due to being unable to perform partition scoped translation. Previously this could only be generated from a guest and so the code path expects us to take the KVM trampoline in the interrupt handler. This is no longer the case so we modify the handler to call bad_page_fault() to check if we were expecting this fault so we can handle it gracefully and just return with an error code. In the hash mmu case we still raise an unknown exception since quadrants aren't defined for the hash mmu. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-14 13:29:05 +08:00
quadrant = 1;
if (!pid)
quadrant = 2;
if (is_load)
from = (void *) (eaddr | (quadrant << 62));
else
to = (void *) (eaddr | (quadrant << 62));
preempt_disable();
/* switch the lpid first to avoid running host with unallocated pid */
old_lpid = mfspr(SPRN_LPID);
if (old_lpid != lpid)
mtspr(SPRN_LPID, lpid);
if (quadrant == 1) {
old_pid = mfspr(SPRN_PID);
if (old_pid != pid)
mtspr(SPRN_PID, pid);
}
isync();
pagefault_disable();
if (is_load)
ret = raw_copy_from_user(to, from, n);
else
ret = raw_copy_to_user(to, from, n);
pagefault_enable();
/* switch the pid first to avoid running host with unallocated pid */
if (quadrant == 1 && pid != old_pid)
mtspr(SPRN_PID, old_pid);
if (lpid != old_lpid)
mtspr(SPRN_LPID, old_lpid);
isync();
preempt_enable();
return ret;
}
EXPORT_SYMBOL_GPL(__kvmhv_copy_tofrom_guest_radix);
KVM: PPC: Book3S HV: Implement functions to access quadrants 1 & 2 The POWER9 radix mmu has the concept of quadrants. The quadrant number is the two high bits of the effective address and determines the fully qualified address to be used for the translation. The fully qualified address consists of the effective lpid, the effective pid and the effective address. This gives then 4 possible quadrants 0, 1, 2, and 3. When accessing these quadrants the fully qualified address is obtained as follows: Quadrant | Hypervisor | Guest -------------------------------------------------------------------------- | EA[0:1] = 0b00 | EA[0:1] = 0b00 0 | effLPID = 0 | effLPID = LPIDR | effPID = PIDR | effPID = PIDR -------------------------------------------------------------------------- | EA[0:1] = 0b01 | 1 | effLPID = LPIDR | Invalid Access | effPID = PIDR | -------------------------------------------------------------------------- | EA[0:1] = 0b10 | 2 | effLPID = LPIDR | Invalid Access | effPID = 0 | -------------------------------------------------------------------------- | EA[0:1] = 0b11 | EA[0:1] = 0b11 3 | effLPID = 0 | effLPID = LPIDR | effPID = 0 | effPID = 0 -------------------------------------------------------------------------- In the Guest; Quadrant 3 is normally used to address the operating system since this uses effPID=0 and effLPID=LPIDR, meaning the PID register doesn't need to be switched. Quadrant 0 is normally used to address user space since the effLPID and effPID are taken from the corresponding registers. In the Host; Quadrant 0 and 3 are used as above, however the effLPID is always 0 to address the host. Quadrants 1 and 2 can be used by the host to address guest memory using a guest effective address. Since the effLPID comes from the LPID register, the host loads the LPID of the guest it would like to access (and the PID of the process) and can perform accesses to a guest effective address. This means quadrant 1 can be used to address the guest user space and quadrant 2 can be used to address the guest operating system from the hypervisor, using a guest effective address. Access to the quadrants can cause a Hypervisor Data Storage Interrupt (HDSI) due to being unable to perform partition scoped translation. Previously this could only be generated from a guest and so the code path expects us to take the KVM trampoline in the interrupt handler. This is no longer the case so we modify the handler to call bad_page_fault() to check if we were expecting this fault so we can handle it gracefully and just return with an error code. In the hash mmu case we still raise an unknown exception since quadrants aren't defined for the hash mmu. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-14 13:29:05 +08:00
static long kvmhv_copy_tofrom_guest_radix(struct kvm_vcpu *vcpu, gva_t eaddr,
void *to, void *from, unsigned long n)
{
int lpid = vcpu->kvm->arch.lpid;
int pid = vcpu->arch.pid;
/* This would cause a data segment intr so don't allow the access */
if (eaddr & (0x3FFUL << 52))
return -EINVAL;
/* Should we be using the nested lpid */
if (vcpu->arch.nested)
lpid = vcpu->arch.nested->shadow_lpid;
/* If accessing quadrant 3 then pid is expected to be 0 */
if (((eaddr >> 62) & 0x3) == 0x3)
pid = 0;
eaddr &= ~(0xFFFUL << 52);
return __kvmhv_copy_tofrom_guest_radix(lpid, pid, eaddr, to, from, n);
}
long kvmhv_copy_from_guest_radix(struct kvm_vcpu *vcpu, gva_t eaddr, void *to,
unsigned long n)
{
long ret;
ret = kvmhv_copy_tofrom_guest_radix(vcpu, eaddr, to, NULL, n);
if (ret > 0)
memset(to + (n - ret), 0, ret);
return ret;
}
EXPORT_SYMBOL_GPL(kvmhv_copy_from_guest_radix);
long kvmhv_copy_to_guest_radix(struct kvm_vcpu *vcpu, gva_t eaddr, void *from,
unsigned long n)
{
return kvmhv_copy_tofrom_guest_radix(vcpu, eaddr, NULL, from, n);
}
EXPORT_SYMBOL_GPL(kvmhv_copy_to_guest_radix);
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
int kvmppc_mmu_walk_radix_tree(struct kvm_vcpu *vcpu, gva_t eaddr,
struct kvmppc_pte *gpte, u64 root,
u64 *pte_ret_p)
{
struct kvm *kvm = vcpu->kvm;
int ret, level, ps;
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
unsigned long rts, bits, offset, index;
u64 pte, base, gpa;
__be64 rpte;
rts = ((root & RTS1_MASK) >> (RTS1_SHIFT - 3)) |
((root & RTS2_MASK) >> RTS2_SHIFT);
bits = root & RPDS_MASK;
base = root & RPDB_MASK;
offset = rts + 31;
/* Current implementations only support 52-bit space */
if (offset != 52)
return -EINVAL;
/* Walk each level of the radix tree */
for (level = 3; level >= 0; --level) {
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
u64 addr;
/* Check a valid size */
if (level && bits != p9_supported_radix_bits[level])
return -EINVAL;
if (level == 0 && !(bits == 5 || bits == 9))
return -EINVAL;
offset -= bits;
index = (eaddr >> offset) & ((1UL << bits) - 1);
/* Check that low bits of page table base are zero */
if (base & ((1UL << (bits + 3)) - 1))
return -EINVAL;
/* Read the entry from guest memory */
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
addr = base + (index * sizeof(rpte));
ret = kvm_read_guest(kvm, addr, &rpte, sizeof(rpte));
if (ret) {
if (pte_ret_p)
*pte_ret_p = addr;
return ret;
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
}
pte = __be64_to_cpu(rpte);
if (!(pte & _PAGE_PRESENT))
return -ENOENT;
/* Check if a leaf entry */
if (pte & _PAGE_PTE)
break;
/* Get ready to walk the next level */
base = pte & RPDB_MASK;
bits = pte & RPDS_MASK;
}
/* Need a leaf at lowest level; 512GB pages not supported */
if (level < 0 || level == 3)
return -EINVAL;
/* We found a valid leaf PTE */
/* Offset is now log base 2 of the page size */
gpa = pte & 0x01fffffffffff000ul;
if (gpa & ((1ul << offset) - 1))
return -EINVAL;
gpa |= eaddr & ((1ul << offset) - 1);
for (ps = MMU_PAGE_4K; ps < MMU_PAGE_COUNT; ++ps)
if (offset == mmu_psize_defs[ps].shift)
break;
gpte->page_size = ps;
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
gpte->page_shift = offset;
gpte->eaddr = eaddr;
gpte->raddr = gpa;
/* Work out permissions */
gpte->may_read = !!(pte & _PAGE_READ);
gpte->may_write = !!(pte & _PAGE_WRITE);
gpte->may_execute = !!(pte & _PAGE_EXEC);
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
gpte->rc = pte & (_PAGE_ACCESSED | _PAGE_DIRTY);
if (pte_ret_p)
*pte_ret_p = pte;
return 0;
}
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
/*
* Used to walk a partition or process table radix tree in guest memory
* Note: We exploit the fact that a partition table and a process
* table have the same layout, a partition-scoped page table and a
* process-scoped page table have the same layout, and the 2nd
* doubleword of a partition table entry has the same layout as
* the PTCR register.
*/
int kvmppc_mmu_radix_translate_table(struct kvm_vcpu *vcpu, gva_t eaddr,
struct kvmppc_pte *gpte, u64 table,
int table_index, u64 *pte_ret_p)
{
struct kvm *kvm = vcpu->kvm;
int ret;
unsigned long size, ptbl, root;
struct prtb_entry entry;
if ((table & PRTS_MASK) > 24)
return -EINVAL;
size = 1ul << ((table & PRTS_MASK) + 12);
/* Is the table big enough to contain this entry? */
if ((table_index * sizeof(entry)) >= size)
return -EINVAL;
/* Read the table to find the root of the radix tree */
ptbl = (table & PRTB_MASK) + (table_index * sizeof(entry));
ret = kvm_read_guest(kvm, ptbl, &entry, sizeof(entry));
if (ret)
return ret;
/* Root is stored in the first double word */
root = be64_to_cpu(entry.prtb0);
return kvmppc_mmu_walk_radix_tree(vcpu, eaddr, gpte, root, pte_ret_p);
}
int kvmppc_mmu_radix_xlate(struct kvm_vcpu *vcpu, gva_t eaddr,
struct kvmppc_pte *gpte, bool data, bool iswrite)
{
u32 pid;
u64 pte;
int ret;
/* Work out effective PID */
switch (eaddr >> 62) {
case 0:
pid = vcpu->arch.pid;
break;
case 3:
pid = 0;
break;
default:
return -EINVAL;
}
ret = kvmppc_mmu_radix_translate_table(vcpu, eaddr, gpte,
vcpu->kvm->arch.process_table, pid, &pte);
if (ret)
return ret;
/* Check privilege (applies only to process scoped translations) */
if (kvmppc_get_msr(vcpu) & MSR_PR) {
if (pte & _PAGE_PRIVILEGED) {
gpte->may_read = 0;
gpte->may_write = 0;
gpte->may_execute = 0;
}
} else {
if (!(pte & _PAGE_PRIVILEGED)) {
/* Check AMR/IAMR to see if strict mode is in force */
if (vcpu->arch.amr & (1ul << 62))
gpte->may_read = 0;
if (vcpu->arch.amr & (1ul << 63))
gpte->may_write = 0;
if (vcpu->arch.iamr & (1ul << 62))
gpte->may_execute = 0;
}
}
return 0;
}
void kvmppc_radix_tlbie_page(struct kvm *kvm, unsigned long addr,
unsigned int pshift, unsigned int lpid)
{
unsigned long psize = PAGE_SIZE;
int psi;
long rc;
unsigned long rb;
if (pshift)
psize = 1UL << pshift;
else
pshift = PAGE_SHIFT;
addr &= ~(psize - 1);
if (!kvmhv_on_pseries()) {
radix__flush_tlb_lpid_page(lpid, addr, psize);
return;
}
psi = shift_to_mmu_psize(pshift);
rb = addr | (mmu_get_ap(psi) << PPC_BITLSHIFT(58));
rc = plpar_hcall_norets(H_TLB_INVALIDATE, H_TLBIE_P1_ENC(0, 0, 1),
lpid, rb);
if (rc)
pr_err("KVM: TLB page invalidation hcall failed, rc=%ld\n", rc);
}
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
static void kvmppc_radix_flush_pwc(struct kvm *kvm, unsigned int lpid)
{
long rc;
if (!kvmhv_on_pseries()) {
radix__flush_pwc_lpid(lpid);
return;
}
rc = plpar_hcall_norets(H_TLB_INVALIDATE, H_TLBIE_P1_ENC(1, 0, 1),
lpid, TLBIEL_INVAL_SET_LPID);
if (rc)
pr_err("KVM: TLB PWC invalidation hcall failed, rc=%ld\n", rc);
}
static unsigned long kvmppc_radix_update_pte(struct kvm *kvm, pte_t *ptep,
unsigned long clr, unsigned long set,
unsigned long addr, unsigned int shift)
{
return __radix_pte_update(ptep, clr, set);
}
void kvmppc_radix_set_pte_at(struct kvm *kvm, unsigned long addr,
pte_t *ptep, pte_t pte)
{
radix__set_pte_at(kvm->mm, addr, ptep, pte, 0);
}
static struct kmem_cache *kvm_pte_cache;
static struct kmem_cache *kvm_pmd_cache;
static pte_t *kvmppc_pte_alloc(void)
{
pte_t *pte;
pte = kmem_cache_alloc(kvm_pte_cache, GFP_KERNEL);
/* pmd_populate() will only reference _pa(pte). */
kmemleak_ignore(pte);
return pte;
}
static void kvmppc_pte_free(pte_t *ptep)
{
kmem_cache_free(kvm_pte_cache, ptep);
}
static pmd_t *kvmppc_pmd_alloc(void)
{
pmd_t *pmd;
pmd = kmem_cache_alloc(kvm_pmd_cache, GFP_KERNEL);
/* pud_populate() will only reference _pa(pmd). */
kmemleak_ignore(pmd);
return pmd;
}
static void kvmppc_pmd_free(pmd_t *pmdp)
{
kmem_cache_free(kvm_pmd_cache, pmdp);
}
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
/* Called with kvm->mmu_lock held */
void kvmppc_unmap_pte(struct kvm *kvm, pte_t *pte, unsigned long gpa,
unsigned int shift,
const struct kvm_memory_slot *memslot,
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
unsigned int lpid)
{
unsigned long old;
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
unsigned long gfn = gpa >> PAGE_SHIFT;
unsigned long page_size = PAGE_SIZE;
unsigned long hpa;
old = kvmppc_radix_update_pte(kvm, pte, ~0UL, 0, gpa, shift);
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_radix_tlbie_page(kvm, gpa, shift, lpid);
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
/* The following only applies to L1 entries */
if (lpid != kvm->arch.lpid)
return;
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
if (!memslot) {
memslot = gfn_to_memslot(kvm, gfn);
if (!memslot)
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
return;
}
if (shift) { /* 1GB or 2MB page */
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
page_size = 1ul << shift;
if (shift == PMD_SHIFT)
kvm->stat.num_2M_pages--;
else if (shift == PUD_SHIFT)
kvm->stat.num_1G_pages--;
}
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
gpa &= ~(page_size - 1);
hpa = old & PTE_RPN_MASK;
kvmhv_remove_nest_rmap_range(kvm, memslot, gpa, hpa, page_size);
if ((old & _PAGE_DIRTY) && memslot->dirty_bitmap)
kvmppc_update_dirty_map(memslot, gfn, page_size);
}
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
/*
* kvmppc_free_p?d are used to free existing page tables, and recursively
* descend and clear and free children.
* Callers are responsible for flushing the PWC.
*
* When page tables are being unmapped/freed as part of page fault path
* (full == false), ptes are not expected. There is code to unmap them
* and emit a warning if encountered, but there may already be data
* corruption due to the unexpected mappings.
*/
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
static void kvmppc_unmap_free_pte(struct kvm *kvm, pte_t *pte, bool full,
unsigned int lpid)
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
{
if (full) {
memset(pte, 0, sizeof(long) << PTE_INDEX_SIZE);
} else {
pte_t *p = pte;
unsigned long it;
for (it = 0; it < PTRS_PER_PTE; ++it, ++p) {
if (pte_val(*p) == 0)
continue;
WARN_ON_ONCE(1);
kvmppc_unmap_pte(kvm, p,
pte_pfn(*p) << PAGE_SHIFT,
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
PAGE_SHIFT, NULL, lpid);
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
}
}
kvmppc_pte_free(pte);
}
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
static void kvmppc_unmap_free_pmd(struct kvm *kvm, pmd_t *pmd, bool full,
unsigned int lpid)
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
{
unsigned long im;
pmd_t *p = pmd;
for (im = 0; im < PTRS_PER_PMD; ++im, ++p) {
if (!pmd_present(*p))
continue;
if (pmd_is_leaf(*p)) {
if (full) {
pmd_clear(p);
} else {
WARN_ON_ONCE(1);
kvmppc_unmap_pte(kvm, (pte_t *)p,
pte_pfn(*(pte_t *)p) << PAGE_SHIFT,
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
PMD_SHIFT, NULL, lpid);
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
}
} else {
pte_t *pte;
pte = pte_offset_map(p, 0);
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_free_pte(kvm, pte, full, lpid);
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
pmd_clear(p);
}
}
kvmppc_pmd_free(pmd);
}
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
static void kvmppc_unmap_free_pud(struct kvm *kvm, pud_t *pud,
unsigned int lpid)
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
{
unsigned long iu;
pud_t *p = pud;
for (iu = 0; iu < PTRS_PER_PUD; ++iu, ++p) {
if (!pud_present(*p))
continue;
if (pud_is_leaf(*p)) {
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
pud_clear(p);
} else {
pmd_t *pmd;
pmd = pmd_offset(p, 0);
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_free_pmd(kvm, pmd, true, lpid);
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
pud_clear(p);
}
}
pud_free(kvm->mm, pud);
}
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
void kvmppc_free_pgtable_radix(struct kvm *kvm, pgd_t *pgd, unsigned int lpid)
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
{
unsigned long ig;
for (ig = 0; ig < PTRS_PER_PGD; ++ig, ++pgd) {
pud_t *pud;
if (!pgd_present(*pgd))
continue;
pud = pud_offset(pgd, 0);
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_free_pud(kvm, pud, lpid);
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
pgd_clear(pgd);
}
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
}
void kvmppc_free_radix(struct kvm *kvm)
{
if (kvm->arch.pgtable) {
kvmppc_free_pgtable_radix(kvm, kvm->arch.pgtable,
kvm->arch.lpid);
pgd_free(kvm->mm, kvm->arch.pgtable);
kvm->arch.pgtable = NULL;
}
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
}
static void kvmppc_unmap_free_pmd_entry_table(struct kvm *kvm, pmd_t *pmd,
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
unsigned long gpa, unsigned int lpid)
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
{
pte_t *pte = pte_offset_kernel(pmd, 0);
/*
* Clearing the pmd entry then flushing the PWC ensures that the pte
* page no longer be cached by the MMU, so can be freed without
* flushing the PWC again.
*/
pmd_clear(pmd);
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_radix_flush_pwc(kvm, lpid);
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_free_pte(kvm, pte, false, lpid);
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
}
static void kvmppc_unmap_free_pud_entry_table(struct kvm *kvm, pud_t *pud,
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
unsigned long gpa, unsigned int lpid)
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
{
pmd_t *pmd = pmd_offset(pud, 0);
/*
* Clearing the pud entry then flushing the PWC ensures that the pmd
* page and any children pte pages will no longer be cached by the MMU,
* so can be freed without flushing the PWC again.
*/
pud_clear(pud);
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_radix_flush_pwc(kvm, lpid);
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_free_pmd(kvm, pmd, false, lpid);
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
}
/*
* There are a number of bits which may differ between different faults to
* the same partition scope entry. RC bits, in the course of cleaning and
* aging. And the write bit can change, either the access could have been
* upgraded, or a read fault could happen concurrently with a write fault
* that sets those bits first.
*/
#define PTE_BITS_MUST_MATCH (~(_PAGE_WRITE | _PAGE_DIRTY | _PAGE_ACCESSED))
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
int kvmppc_create_pte(struct kvm *kvm, pgd_t *pgtable, pte_t pte,
unsigned long gpa, unsigned int level,
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
unsigned long mmu_seq, unsigned int lpid,
unsigned long *rmapp, struct rmap_nested **n_rmap)
{
pgd_t *pgd;
pud_t *pud, *new_pud = NULL;
pmd_t *pmd, *new_pmd = NULL;
pte_t *ptep, *new_ptep = NULL;
int ret;
/* Traverse the guest's 2nd-level tree, allocate new levels needed */
pgd = pgtable + pgd_index(gpa);
pud = NULL;
if (pgd_present(*pgd))
pud = pud_offset(pgd, gpa);
else
new_pud = pud_alloc_one(kvm->mm, gpa);
pmd = NULL;
if (pud && pud_present(*pud) && !pud_is_leaf(*pud))
pmd = pmd_offset(pud, gpa);
else if (level <= 1)
new_pmd = kvmppc_pmd_alloc();
if (level == 0 && !(pmd && pmd_present(*pmd) && !pmd_is_leaf(*pmd)))
new_ptep = kvmppc_pte_alloc();
/* Check if we might have been invalidated; let the guest retry if so */
spin_lock(&kvm->mmu_lock);
ret = -EAGAIN;
if (mmu_notifier_retry(kvm, mmu_seq))
goto out_unlock;
/* Now traverse again under the lock and change the tree */
ret = -ENOMEM;
if (pgd_none(*pgd)) {
if (!new_pud)
goto out_unlock;
pgd_populate(kvm->mm, pgd, new_pud);
new_pud = NULL;
}
pud = pud_offset(pgd, gpa);
if (pud_is_leaf(*pud)) {
unsigned long hgpa = gpa & PUD_MASK;
/* Check if we raced and someone else has set the same thing */
if (level == 2) {
if (pud_raw(*pud) == pte_raw(pte)) {
ret = 0;
goto out_unlock;
}
/* Valid 1GB page here already, add our extra bits */
WARN_ON_ONCE((pud_val(*pud) ^ pte_val(pte)) &
PTE_BITS_MUST_MATCH);
kvmppc_radix_update_pte(kvm, (pte_t *)pud,
0, pte_val(pte), hgpa, PUD_SHIFT);
ret = 0;
goto out_unlock;
}
/*
* If we raced with another CPU which has just put
* a 1GB pte in after we saw a pmd page, try again.
*/
if (!new_pmd) {
ret = -EAGAIN;
goto out_unlock;
}
/* Valid 1GB page here already, remove it */
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_pte(kvm, (pte_t *)pud, hgpa, PUD_SHIFT, NULL,
lpid);
}
if (level == 2) {
if (!pud_none(*pud)) {
/*
* There's a page table page here, but we wanted to
* install a large page, so remove and free the page
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
* table page.
*/
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_free_pud_entry_table(kvm, pud, gpa, lpid);
}
kvmppc_radix_set_pte_at(kvm, gpa, (pte_t *)pud, pte);
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
if (rmapp && n_rmap)
kvmhv_insert_nest_rmap(kvm, rmapp, n_rmap);
ret = 0;
goto out_unlock;
}
if (pud_none(*pud)) {
if (!new_pmd)
goto out_unlock;
pud_populate(kvm->mm, pud, new_pmd);
new_pmd = NULL;
}
pmd = pmd_offset(pud, gpa);
if (pmd_is_leaf(*pmd)) {
unsigned long lgpa = gpa & PMD_MASK;
/* Check if we raced and someone else has set the same thing */
if (level == 1) {
if (pmd_raw(*pmd) == pte_raw(pte)) {
ret = 0;
goto out_unlock;
}
/* Valid 2MB page here already, add our extra bits */
WARN_ON_ONCE((pmd_val(*pmd) ^ pte_val(pte)) &
PTE_BITS_MUST_MATCH);
kvmppc_radix_update_pte(kvm, pmdp_ptep(pmd),
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
0, pte_val(pte), lgpa, PMD_SHIFT);
ret = 0;
goto out_unlock;
}
/*
* If we raced with another CPU which has just put
* a 2MB pte in after we saw a pte page, try again.
*/
if (!new_ptep) {
ret = -EAGAIN;
goto out_unlock;
}
/* Valid 2MB page here already, remove it */
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_pte(kvm, pmdp_ptep(pmd), lgpa, PMD_SHIFT, NULL,
lpid);
}
if (level == 1) {
if (!pmd_none(*pmd)) {
/*
* There's a page table page here, but we wanted to
* install a large page, so remove and free the page
KVM: PPC: Book3S HV: Recursively unmap all page table entries when unmapping When partition scope mappings are unmapped with kvm_unmap_radix, the pte is cleared, but the page table structure is left in place. If the next page fault requests a different page table geometry (e.g., due to THP promotion or split), kvmppc_create_pte is responsible for changing the page tables. When a page table entry is to be converted to a large pte, the page table entry is cleared, the PWC flushed, then the page table it points to freed. This will cause pte page tables to leak when a 1GB page is to replace a pud entry points to a pmd table with pte tables under it: The pmd table will be freed, but its pte tables will be missed. Fix this by replacing the simple clear and free code with one that walks down the page tables and frees children. Care must be taken to clear the root entry being unmapped then flushing the PWC before freeing any page tables, as explained in comments. This requires PWC flush to logically become a flush-all-PWC (which it already is in hardware, but the KVM API needs to be changed to avoid confusion). This code also checks that no unexpected pte entries exist in any page table being freed, and unmaps those and emits a WARN. This is an expensive operation for the pte page level, but partition scope changes are rare, so it's unconditional for now to iron out bugs. It can be put under a CONFIG option or removed after some time. Signed-off-by: Nicholas Piggin <npiggin@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-05-17 15:06:27 +08:00
* table page.
*/
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_free_pmd_entry_table(kvm, pmd, gpa, lpid);
}
kvmppc_radix_set_pte_at(kvm, gpa, pmdp_ptep(pmd), pte);
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
if (rmapp && n_rmap)
kvmhv_insert_nest_rmap(kvm, rmapp, n_rmap);
ret = 0;
goto out_unlock;
}
if (pmd_none(*pmd)) {
if (!new_ptep)
goto out_unlock;
pmd_populate(kvm->mm, pmd, new_ptep);
new_ptep = NULL;
}
ptep = pte_offset_kernel(pmd, gpa);
if (pte_present(*ptep)) {
/* Check if someone else set the same thing */
if (pte_raw(*ptep) == pte_raw(pte)) {
ret = 0;
goto out_unlock;
}
/* Valid page here already, add our extra bits */
WARN_ON_ONCE((pte_val(*ptep) ^ pte_val(pte)) &
PTE_BITS_MUST_MATCH);
kvmppc_radix_update_pte(kvm, ptep, 0, pte_val(pte), gpa, 0);
ret = 0;
goto out_unlock;
}
kvmppc_radix_set_pte_at(kvm, gpa, ptep, pte);
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
if (rmapp && n_rmap)
kvmhv_insert_nest_rmap(kvm, rmapp, n_rmap);
ret = 0;
out_unlock:
spin_unlock(&kvm->mmu_lock);
if (new_pud)
pud_free(kvm->mm, new_pud);
if (new_pmd)
kvmppc_pmd_free(new_pmd);
if (new_ptep)
kvmppc_pte_free(new_ptep);
return ret;
}
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
bool kvmppc_hv_handle_set_rc(struct kvm *kvm, pgd_t *pgtable, bool writing,
unsigned long gpa, unsigned int lpid)
{
unsigned long pgflags;
unsigned int shift;
pte_t *ptep;
/*
* Need to set an R or C bit in the 2nd-level tables;
* since we are just helping out the hardware here,
* it is sufficient to do what the hardware does.
*/
pgflags = _PAGE_ACCESSED;
if (writing)
pgflags |= _PAGE_DIRTY;
/*
* We are walking the secondary (partition-scoped) page table here.
* We can do this without disabling irq because the Linux MM
* subsystem doesn't do THP splits and collapses on this tree.
*/
ptep = __find_linux_pte(pgtable, gpa, NULL, &shift);
if (ptep && pte_present(*ptep) && (!writing || pte_write(*ptep))) {
kvmppc_radix_update_pte(kvm, ptep, 0, pgflags, gpa, shift);
return true;
}
return false;
}
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
int kvmppc_book3s_instantiate_page(struct kvm_vcpu *vcpu,
unsigned long gpa,
struct kvm_memory_slot *memslot,
bool writing, bool kvm_ro,
pte_t *inserted_pte, unsigned int *levelp)
{
struct kvm *kvm = vcpu->kvm;
struct page *page = NULL;
unsigned long mmu_seq;
unsigned long hva, gfn = gpa >> PAGE_SHIFT;
bool upgrade_write = false;
bool *upgrade_p = &upgrade_write;
pte_t pte, *ptep;
unsigned int shift, level;
int ret;
bool large_enable;
/* used to check for invalidations in progress */
mmu_seq = kvm->mmu_notifier_seq;
smp_rmb();
/*
* Do a fast check first, since __gfn_to_pfn_memslot doesn't
* do it with !atomic && !async, which is how we call it.
* We always ask for write permission since the common case
* is that the page is writable.
*/
hva = gfn_to_hva_memslot(memslot, gfn);
if (!kvm_ro && __get_user_pages_fast(hva, 1, 1, &page) == 1) {
upgrade_write = true;
} else {
unsigned long pfn;
/* Call KVM generic code to do the slow-path check */
pfn = __gfn_to_pfn_memslot(memslot, gfn, false, NULL,
writing, upgrade_p);
if (is_error_noslot_pfn(pfn))
return -EFAULT;
page = NULL;
if (pfn_valid(pfn)) {
page = pfn_to_page(pfn);
if (PageReserved(page))
page = NULL;
}
}
/*
* Read the PTE from the process' radix tree and use that
* so we get the shift and attribute bits.
*/
local_irq_disable();
ptep = __find_linux_pte(vcpu->arch.pgdir, hva, NULL, &shift);
/*
* If the PTE disappeared temporarily due to a THP
* collapse, just return and let the guest try again.
*/
if (!ptep) {
local_irq_enable();
if (page)
put_page(page);
return RESUME_GUEST;
}
pte = *ptep;
local_irq_enable();
/* If we're logging dirty pages, always map single pages */
large_enable = !(memslot->flags & KVM_MEM_LOG_DIRTY_PAGES);
/* Get pte level from shift/size */
if (large_enable && shift == PUD_SHIFT &&
(gpa & (PUD_SIZE - PAGE_SIZE)) ==
(hva & (PUD_SIZE - PAGE_SIZE))) {
level = 2;
} else if (large_enable && shift == PMD_SHIFT &&
(gpa & (PMD_SIZE - PAGE_SIZE)) ==
(hva & (PMD_SIZE - PAGE_SIZE))) {
level = 1;
} else {
level = 0;
if (shift > PAGE_SHIFT) {
/*
* If the pte maps more than one page, bring over
* bits from the virtual address to get the real
* address of the specific single page we want.
*/
unsigned long rpnmask = (1ul << shift) - PAGE_SIZE;
pte = __pte(pte_val(pte) | (hva & rpnmask));
}
}
pte = __pte(pte_val(pte) | _PAGE_EXEC | _PAGE_ACCESSED);
if (writing || upgrade_write) {
if (pte_val(pte) & _PAGE_WRITE)
pte = __pte(pte_val(pte) | _PAGE_DIRTY);
} else {
pte = __pte(pte_val(pte) & ~(_PAGE_WRITE | _PAGE_DIRTY));
}
/* Allocate space in the tree and write the PTE */
ret = kvmppc_create_pte(kvm, kvm->arch.pgtable, pte, gpa, level,
KVM: PPC: Book3S HV: Introduce rmap to track nested guest mappings When a host (L0) page which is mapped into a (L1) guest is in turn mapped through to a nested (L2) guest we keep a reverse mapping (rmap) so that these mappings can be retrieved later. Whenever we create an entry in a shadow_pgtable for a nested guest we create a corresponding rmap entry and add it to the list for the L1 guest memslot at the index of the L1 guest page it maps. This means at the L1 guest memslot we end up with lists of rmaps. When we are notified of a host page being invalidated which has been mapped through to a (L1) guest, we can then walk the rmap list for that guest page, and find and invalidate all of the corresponding shadow_pgtable entries. In order to reduce memory consumption, we compress the information for each rmap entry down to 52 bits -- 12 bits for the LPID and 40 bits for the guest real page frame number -- which will fit in a single unsigned long. To avoid a scenario where a guest can trigger unbounded memory allocations, we scan the list when adding an entry to see if there is already an entry with the contents we need. This can occur, because we don't ever remove entries from the middle of a list. A struct nested guest rmap is a list pointer and an rmap entry; ---------------- | next pointer | ---------------- | rmap entry | ---------------- Thus the rmap pointer for each guest frame number in the memslot can be either NULL, a single entry, or a pointer to a list of nested rmap entries. gfn memslot rmap array ------------------------- 0 | NULL | (no rmap entry) ------------------------- 1 | single rmap entry | (rmap entry with low bit set) ------------------------- 2 | list head pointer | (list of rmap entries) ------------------------- The final entry always has the lowest bit set and is stored in the next pointer of the last list entry, or as a single rmap entry. With a list of rmap entries looking like; ----------------- ----------------- ------------------------- | list head ptr | ----> | next pointer | ----> | single rmap entry | ----------------- ----------------- ------------------------- | rmap entry | | rmap entry | ----------------- ------------------------- Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:08 +08:00
mmu_seq, kvm->arch.lpid, NULL, NULL);
if (inserted_pte)
*inserted_pte = pte;
if (levelp)
*levelp = level;
if (page) {
if (!ret && (pte_val(pte) & _PAGE_WRITE))
set_page_dirty_lock(page);
put_page(page);
}
/* Increment number of large pages if we (successfully) inserted one */
if (!ret) {
if (level == 1)
kvm->stat.num_2M_pages++;
else if (level == 2)
kvm->stat.num_1G_pages++;
}
return ret;
}
int kvmppc_book3s_radix_page_fault(struct kvm_run *run, struct kvm_vcpu *vcpu,
unsigned long ea, unsigned long dsisr)
{
struct kvm *kvm = vcpu->kvm;
unsigned long gpa, gfn;
struct kvm_memory_slot *memslot;
long ret;
bool writing = !!(dsisr & DSISR_ISSTORE);
bool kvm_ro = false;
/* Check for unusual errors */
if (dsisr & DSISR_UNSUPP_MMU) {
pr_err("KVM: Got unsupported MMU fault\n");
return -EFAULT;
}
if (dsisr & DSISR_BADACCESS) {
/* Reflect to the guest as DSI */
pr_err("KVM: Got radix HV page fault with DSISR=%lx\n", dsisr);
kvmppc_core_queue_data_storage(vcpu, ea, dsisr);
return RESUME_GUEST;
}
/* Translate the logical address */
gpa = vcpu->arch.fault_gpa & ~0xfffUL;
gpa &= ~0xF000000000000000ul;
gfn = gpa >> PAGE_SHIFT;
if (!(dsisr & DSISR_PRTABLE_FAULT))
gpa |= ea & 0xfff;
/* Get the corresponding memslot */
memslot = gfn_to_memslot(kvm, gfn);
/* No memslot means it's an emulated MMIO region */
if (!memslot || (memslot->flags & KVM_MEMSLOT_INVALID)) {
if (dsisr & (DSISR_PRTABLE_FAULT | DSISR_BADACCESS |
DSISR_SET_RC)) {
/*
* Bad address in guest page table tree, or other
* unusual error - reflect it to the guest as DSI.
*/
kvmppc_core_queue_data_storage(vcpu, ea, dsisr);
return RESUME_GUEST;
}
return kvmppc_hv_emulate_mmio(run, vcpu, gpa, ea, writing);
}
if (memslot->flags & KVM_MEM_READONLY) {
if (writing) {
/* give the guest a DSI */
kvmppc_core_queue_data_storage(vcpu, ea, DSISR_ISSTORE |
DSISR_PROTFAULT);
return RESUME_GUEST;
}
kvm_ro = true;
}
/* Failed to set the reference/change bits */
if (dsisr & DSISR_SET_RC) {
spin_lock(&kvm->mmu_lock);
if (kvmppc_hv_handle_set_rc(kvm, kvm->arch.pgtable,
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
writing, gpa, kvm->arch.lpid))
dsisr &= ~DSISR_SET_RC;
spin_unlock(&kvm->mmu_lock);
if (!(dsisr & (DSISR_BAD_FAULT_64S | DSISR_NOHPTE |
DSISR_PROTFAULT | DSISR_SET_RC)))
return RESUME_GUEST;
}
/* Try to insert a pte */
ret = kvmppc_book3s_instantiate_page(vcpu, gpa, memslot, writing,
kvm_ro, NULL, NULL);
if (ret == 0 || ret == -EAGAIN)
ret = RESUME_GUEST;
return ret;
}
/* Called with kvm->mmu_lock held */
int kvm_unmap_radix(struct kvm *kvm, struct kvm_memory_slot *memslot,
unsigned long gfn)
{
pte_t *ptep;
unsigned long gpa = gfn << PAGE_SHIFT;
unsigned int shift;
ptep = __find_linux_pte(kvm->arch.pgtable, gpa, NULL, &shift);
if (ptep && pte_present(*ptep))
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_unmap_pte(kvm, ptep, gpa, shift, memslot,
kvm->arch.lpid);
return 0;
}
/* Called with kvm->mmu_lock held */
int kvm_age_radix(struct kvm *kvm, struct kvm_memory_slot *memslot,
unsigned long gfn)
{
pte_t *ptep;
unsigned long gpa = gfn << PAGE_SHIFT;
unsigned int shift;
int ref = 0;
KVM: PPC: Book3S HV: Keep rc bits in shadow pgtable in sync with host The rc bits contained in ptes are used to track whether a page has been accessed and whether it is dirty. The accessed bit is used to age a page and the dirty bit to track whether a page is dirty or not. Now that we support nested guests there are three ptes which track the state of the same page: - The partition-scoped page table in the L1 guest, mapping L2->L1 address - The partition-scoped page table in the host for the L1 guest, mapping L1->L0 address - The shadow partition-scoped page table for the nested guest in the host, mapping L2->L0 address The idea is to attempt to keep the rc state of these three ptes in sync, both when setting and when clearing rc bits. When setting the bits we achieve consistency by: - Initially setting the bits in the shadow page table as the 'and' of the other two. - When updating in software the rc bits in the shadow page table we ensure the state is consistent with the other two locations first, and update these before reflecting the change into the shadow page table. i.e. only set the bits in the L2->L0 pte if also set in both the L2->L1 and the L1->L0 pte. When clearing the bits we achieve consistency by: - The rc bits in the shadow page table are only cleared when discarding a pte, and we don't need to record this as if either bit is set then it must also be set in the pte mapping L1->L0. - When L1 clears an rc bit in the L2->L1 mapping it __should__ issue a tlbie instruction - This means we will discard the pte from the shadow page table meaning the mapping will have to be setup again. - When setup the pte again in the shadow page table we will ensure consistency with the L2->L1 pte. - When the host clears an rc bit in the L1->L0 mapping we need to also clear the bit in any ptes in the shadow page table which map the same gfn so we will be notified if a nested guest accesses the page. This case is what this patch specifically concerns. - We can search the nest_rmap list for that given gfn and clear the same bit from all corresponding ptes in shadow page tables. - If a nested guest causes either of the rc bits to be set by software in future then we will update the L1->L0 pte and maintain consistency. With the process outlined above we aim to maintain consistency of the 3 pte locations where we track rc for a given guest page. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-21 11:28:43 +08:00
unsigned long old, *rmapp;
ptep = __find_linux_pte(kvm->arch.pgtable, gpa, NULL, &shift);
if (ptep && pte_present(*ptep) && pte_young(*ptep)) {
KVM: PPC: Book3S HV: Keep rc bits in shadow pgtable in sync with host The rc bits contained in ptes are used to track whether a page has been accessed and whether it is dirty. The accessed bit is used to age a page and the dirty bit to track whether a page is dirty or not. Now that we support nested guests there are three ptes which track the state of the same page: - The partition-scoped page table in the L1 guest, mapping L2->L1 address - The partition-scoped page table in the host for the L1 guest, mapping L1->L0 address - The shadow partition-scoped page table for the nested guest in the host, mapping L2->L0 address The idea is to attempt to keep the rc state of these three ptes in sync, both when setting and when clearing rc bits. When setting the bits we achieve consistency by: - Initially setting the bits in the shadow page table as the 'and' of the other two. - When updating in software the rc bits in the shadow page table we ensure the state is consistent with the other two locations first, and update these before reflecting the change into the shadow page table. i.e. only set the bits in the L2->L0 pte if also set in both the L2->L1 and the L1->L0 pte. When clearing the bits we achieve consistency by: - The rc bits in the shadow page table are only cleared when discarding a pte, and we don't need to record this as if either bit is set then it must also be set in the pte mapping L1->L0. - When L1 clears an rc bit in the L2->L1 mapping it __should__ issue a tlbie instruction - This means we will discard the pte from the shadow page table meaning the mapping will have to be setup again. - When setup the pte again in the shadow page table we will ensure consistency with the L2->L1 pte. - When the host clears an rc bit in the L1->L0 mapping we need to also clear the bit in any ptes in the shadow page table which map the same gfn so we will be notified if a nested guest accesses the page. This case is what this patch specifically concerns. - We can search the nest_rmap list for that given gfn and clear the same bit from all corresponding ptes in shadow page tables. - If a nested guest causes either of the rc bits to be set by software in future then we will update the L1->L0 pte and maintain consistency. With the process outlined above we aim to maintain consistency of the 3 pte locations where we track rc for a given guest page. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-21 11:28:43 +08:00
old = kvmppc_radix_update_pte(kvm, ptep, _PAGE_ACCESSED, 0,
gpa, shift);
/* XXX need to flush tlb here? */
KVM: PPC: Book3S HV: Keep rc bits in shadow pgtable in sync with host The rc bits contained in ptes are used to track whether a page has been accessed and whether it is dirty. The accessed bit is used to age a page and the dirty bit to track whether a page is dirty or not. Now that we support nested guests there are three ptes which track the state of the same page: - The partition-scoped page table in the L1 guest, mapping L2->L1 address - The partition-scoped page table in the host for the L1 guest, mapping L1->L0 address - The shadow partition-scoped page table for the nested guest in the host, mapping L2->L0 address The idea is to attempt to keep the rc state of these three ptes in sync, both when setting and when clearing rc bits. When setting the bits we achieve consistency by: - Initially setting the bits in the shadow page table as the 'and' of the other two. - When updating in software the rc bits in the shadow page table we ensure the state is consistent with the other two locations first, and update these before reflecting the change into the shadow page table. i.e. only set the bits in the L2->L0 pte if also set in both the L2->L1 and the L1->L0 pte. When clearing the bits we achieve consistency by: - The rc bits in the shadow page table are only cleared when discarding a pte, and we don't need to record this as if either bit is set then it must also be set in the pte mapping L1->L0. - When L1 clears an rc bit in the L2->L1 mapping it __should__ issue a tlbie instruction - This means we will discard the pte from the shadow page table meaning the mapping will have to be setup again. - When setup the pte again in the shadow page table we will ensure consistency with the L2->L1 pte. - When the host clears an rc bit in the L1->L0 mapping we need to also clear the bit in any ptes in the shadow page table which map the same gfn so we will be notified if a nested guest accesses the page. This case is what this patch specifically concerns. - We can search the nest_rmap list for that given gfn and clear the same bit from all corresponding ptes in shadow page tables. - If a nested guest causes either of the rc bits to be set by software in future then we will update the L1->L0 pte and maintain consistency. With the process outlined above we aim to maintain consistency of the 3 pte locations where we track rc for a given guest page. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-21 11:28:43 +08:00
/* Also clear bit in ptes in shadow pgtable for nested guests */
rmapp = &memslot->arch.rmap[gfn - memslot->base_gfn];
kvmhv_update_nest_rmap_rc_list(kvm, rmapp, _PAGE_ACCESSED, 0,
old & PTE_RPN_MASK,
1UL << shift);
ref = 1;
}
return ref;
}
/* Called with kvm->mmu_lock held */
int kvm_test_age_radix(struct kvm *kvm, struct kvm_memory_slot *memslot,
unsigned long gfn)
{
pte_t *ptep;
unsigned long gpa = gfn << PAGE_SHIFT;
unsigned int shift;
int ref = 0;
ptep = __find_linux_pte(kvm->arch.pgtable, gpa, NULL, &shift);
if (ptep && pte_present(*ptep) && pte_young(*ptep))
ref = 1;
return ref;
}
/* Returns the number of PAGE_SIZE pages that are dirty */
static int kvm_radix_test_clear_dirty(struct kvm *kvm,
struct kvm_memory_slot *memslot, int pagenum)
{
unsigned long gfn = memslot->base_gfn + pagenum;
unsigned long gpa = gfn << PAGE_SHIFT;
pte_t *ptep;
unsigned int shift;
int ret = 0;
KVM: PPC: Book3S HV: Keep rc bits in shadow pgtable in sync with host The rc bits contained in ptes are used to track whether a page has been accessed and whether it is dirty. The accessed bit is used to age a page and the dirty bit to track whether a page is dirty or not. Now that we support nested guests there are three ptes which track the state of the same page: - The partition-scoped page table in the L1 guest, mapping L2->L1 address - The partition-scoped page table in the host for the L1 guest, mapping L1->L0 address - The shadow partition-scoped page table for the nested guest in the host, mapping L2->L0 address The idea is to attempt to keep the rc state of these three ptes in sync, both when setting and when clearing rc bits. When setting the bits we achieve consistency by: - Initially setting the bits in the shadow page table as the 'and' of the other two. - When updating in software the rc bits in the shadow page table we ensure the state is consistent with the other two locations first, and update these before reflecting the change into the shadow page table. i.e. only set the bits in the L2->L0 pte if also set in both the L2->L1 and the L1->L0 pte. When clearing the bits we achieve consistency by: - The rc bits in the shadow page table are only cleared when discarding a pte, and we don't need to record this as if either bit is set then it must also be set in the pte mapping L1->L0. - When L1 clears an rc bit in the L2->L1 mapping it __should__ issue a tlbie instruction - This means we will discard the pte from the shadow page table meaning the mapping will have to be setup again. - When setup the pte again in the shadow page table we will ensure consistency with the L2->L1 pte. - When the host clears an rc bit in the L1->L0 mapping we need to also clear the bit in any ptes in the shadow page table which map the same gfn so we will be notified if a nested guest accesses the page. This case is what this patch specifically concerns. - We can search the nest_rmap list for that given gfn and clear the same bit from all corresponding ptes in shadow page tables. - If a nested guest causes either of the rc bits to be set by software in future then we will update the L1->L0 pte and maintain consistency. With the process outlined above we aim to maintain consistency of the 3 pte locations where we track rc for a given guest page. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-21 11:28:43 +08:00
unsigned long old, *rmapp;
ptep = __find_linux_pte(kvm->arch.pgtable, gpa, NULL, &shift);
if (ptep && pte_present(*ptep) && pte_dirty(*ptep)) {
ret = 1;
if (shift)
ret = 1 << (shift - PAGE_SHIFT);
KVM: PPC: Book3S HV: Keep rc bits in shadow pgtable in sync with host The rc bits contained in ptes are used to track whether a page has been accessed and whether it is dirty. The accessed bit is used to age a page and the dirty bit to track whether a page is dirty or not. Now that we support nested guests there are three ptes which track the state of the same page: - The partition-scoped page table in the L1 guest, mapping L2->L1 address - The partition-scoped page table in the host for the L1 guest, mapping L1->L0 address - The shadow partition-scoped page table for the nested guest in the host, mapping L2->L0 address The idea is to attempt to keep the rc state of these three ptes in sync, both when setting and when clearing rc bits. When setting the bits we achieve consistency by: - Initially setting the bits in the shadow page table as the 'and' of the other two. - When updating in software the rc bits in the shadow page table we ensure the state is consistent with the other two locations first, and update these before reflecting the change into the shadow page table. i.e. only set the bits in the L2->L0 pte if also set in both the L2->L1 and the L1->L0 pte. When clearing the bits we achieve consistency by: - The rc bits in the shadow page table are only cleared when discarding a pte, and we don't need to record this as if either bit is set then it must also be set in the pte mapping L1->L0. - When L1 clears an rc bit in the L2->L1 mapping it __should__ issue a tlbie instruction - This means we will discard the pte from the shadow page table meaning the mapping will have to be setup again. - When setup the pte again in the shadow page table we will ensure consistency with the L2->L1 pte. - When the host clears an rc bit in the L1->L0 mapping we need to also clear the bit in any ptes in the shadow page table which map the same gfn so we will be notified if a nested guest accesses the page. This case is what this patch specifically concerns. - We can search the nest_rmap list for that given gfn and clear the same bit from all corresponding ptes in shadow page tables. - If a nested guest causes either of the rc bits to be set by software in future then we will update the L1->L0 pte and maintain consistency. With the process outlined above we aim to maintain consistency of the 3 pte locations where we track rc for a given guest page. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-21 11:28:43 +08:00
spin_lock(&kvm->mmu_lock);
old = kvmppc_radix_update_pte(kvm, ptep, _PAGE_DIRTY, 0,
gpa, shift);
KVM: PPC: Book3S HV: Handle page fault for a nested guest Consider a normal (L1) guest running under the main hypervisor (L0), and then a nested guest (L2) running under the L1 guest which is acting as a nested hypervisor. L0 has page tables to map the address space for L1 providing the translation from L1 real address -> L0 real address; L1 | | (L1 -> L0) | ----> L0 There are also page tables in L1 used to map the address space for L2 providing the translation from L2 real address -> L1 read address. Since the hardware can only walk a single level of page table, we need to maintain in L0 a "shadow_pgtable" for L2 which provides the translation from L2 real address -> L0 real address. Which looks like; L2 L2 | | | (L2 -> L1) | | | ----> L1 | (L2 -> L0) | | | (L1 -> L0) | | | ----> L0 --------> L0 When a page fault occurs while running a nested (L2) guest we need to insert a pte into this "shadow_pgtable" for the L2 -> L0 mapping. To do this we need to: 1. Walk the pgtable in L1 memory to find the L2 -> L1 mapping, and provide a page fault to L1 if this mapping doesn't exist. 2. Use our L1 -> L0 pgtable to convert this L1 address to an L0 address, or try to insert a pte for that mapping if it doesn't exist. 3. Now we have a L2 -> L0 mapping, insert this into our shadow_pgtable Once this mapping exists we can take rc faults when hardware is unable to automatically set the reference and change bits in the pte. On these we need to: 1. Check the rc bits on the L2 -> L1 pte match, and otherwise reflect the fault down to L1. 2. Set the rc bits in the L1 -> L0 pte which corresponds to the same host page. 3. Set the rc bits in the L2 -> L0 pte. As we reuse a large number of functions in book3s_64_mmu_radix.c for this we also needed to refactor a number of these functions to take an lpid parameter so that the correct lpid is used for tlb invalidations. The functionality however has remained the same. Reviewed-by: David Gibson <david@gibson.dropbear.id.au> Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org> Signed-off-by: Michael Ellerman <mpe@ellerman.id.au>
2018-10-08 13:31:07 +08:00
kvmppc_radix_tlbie_page(kvm, gpa, shift, kvm->arch.lpid);
KVM: PPC: Book3S HV: Keep rc bits in shadow pgtable in sync with host The rc bits contained in ptes are used to track whether a page has been accessed and whether it is dirty. The accessed bit is used to age a page and the dirty bit to track whether a page is dirty or not. Now that we support nested guests there are three ptes which track the state of the same page: - The partition-scoped page table in the L1 guest, mapping L2->L1 address - The partition-scoped page table in the host for the L1 guest, mapping L1->L0 address - The shadow partition-scoped page table for the nested guest in the host, mapping L2->L0 address The idea is to attempt to keep the rc state of these three ptes in sync, both when setting and when clearing rc bits. When setting the bits we achieve consistency by: - Initially setting the bits in the shadow page table as the 'and' of the other two. - When updating in software the rc bits in the shadow page table we ensure the state is consistent with the other two locations first, and update these before reflecting the change into the shadow page table. i.e. only set the bits in the L2->L0 pte if also set in both the L2->L1 and the L1->L0 pte. When clearing the bits we achieve consistency by: - The rc bits in the shadow page table are only cleared when discarding a pte, and we don't need to record this as if either bit is set then it must also be set in the pte mapping L1->L0. - When L1 clears an rc bit in the L2->L1 mapping it __should__ issue a tlbie instruction - This means we will discard the pte from the shadow page table meaning the mapping will have to be setup again. - When setup the pte again in the shadow page table we will ensure consistency with the L2->L1 pte. - When the host clears an rc bit in the L1->L0 mapping we need to also clear the bit in any ptes in the shadow page table which map the same gfn so we will be notified if a nested guest accesses the page. This case is what this patch specifically concerns. - We can search the nest_rmap list for that given gfn and clear the same bit from all corresponding ptes in shadow page tables. - If a nested guest causes either of the rc bits to be set by software in future then we will update the L1->L0 pte and maintain consistency. With the process outlined above we aim to maintain consistency of the 3 pte locations where we track rc for a given guest page. Signed-off-by: Suraj Jitindar Singh <sjitindarsingh@gmail.com> Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2018-12-21 11:28:43 +08:00
/* Also clear bit in ptes in shadow pgtable for nested guests */
rmapp = &memslot->arch.rmap[gfn - memslot->base_gfn];
kvmhv_update_nest_rmap_rc_list(kvm, rmapp, _PAGE_DIRTY, 0,
old & PTE_RPN_MASK,
1UL << shift);
spin_unlock(&kvm->mmu_lock);
}
return ret;
}
long kvmppc_hv_get_dirty_log_radix(struct kvm *kvm,
struct kvm_memory_slot *memslot, unsigned long *map)
{
unsigned long i, j;
int npages;
for (i = 0; i < memslot->npages; i = j) {
npages = kvm_radix_test_clear_dirty(kvm, memslot, i);
/*
* Note that if npages > 0 then i must be a multiple of npages,
* since huge pages are only used to back the guest at guest
* real addresses that are a multiple of their size.
* Since we have at most one PTE covering any given guest
* real address, if npages > 1 we can skip to i + npages.
*/
j = i + 1;
KVM: PPC: Book3S HV: Unify dirty page map between HPT and radix Currently, the HPT code in HV KVM maintains a dirty bit per guest page in the rmap array, whether or not dirty page tracking has been enabled for the memory slot. In contrast, the radix code maintains a dirty bit per guest page in memslot->dirty_bitmap, and only does so when dirty page tracking has been enabled. This changes the HPT code to maintain the dirty bits in the memslot dirty_bitmap like radix does. This results in slightly less code overall, and will mean that we do not lose the dirty bits when transitioning between HPT and radix mode in future. There is one minor change to behaviour as a result. With HPT, when dirty tracking was enabled for a memslot, we would previously clear all the dirty bits at that point (both in the HPT entries and in the rmap arrays), meaning that a KVM_GET_DIRTY_LOG ioctl immediately following would show no pages as dirty (assuming no vcpus have run in the meantime). With this change, the dirty bits on HPT entries are not cleared at the point where dirty tracking is enabled, so KVM_GET_DIRTY_LOG would show as dirty any guest pages that are resident in the HPT and dirty. This is consistent with what happens on radix. This also fixes a bug in the mark_pages_dirty() function for radix (in the sense that the function no longer exists). In the case where a large page of 64 normal pages or more is marked dirty, the addressing of the dirty bitmap was incorrect and could write past the end of the bitmap. Fortunately this case was never hit in practice because a 2MB large page is only 32 x 64kB pages, and we don't support backing the guest with 1GB huge pages at this point. Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2017-10-26 13:39:19 +08:00
if (npages) {
set_dirty_bits(map, i, npages);
j = i + npages;
KVM: PPC: Book3S HV: Unify dirty page map between HPT and radix Currently, the HPT code in HV KVM maintains a dirty bit per guest page in the rmap array, whether or not dirty page tracking has been enabled for the memory slot. In contrast, the radix code maintains a dirty bit per guest page in memslot->dirty_bitmap, and only does so when dirty page tracking has been enabled. This changes the HPT code to maintain the dirty bits in the memslot dirty_bitmap like radix does. This results in slightly less code overall, and will mean that we do not lose the dirty bits when transitioning between HPT and radix mode in future. There is one minor change to behaviour as a result. With HPT, when dirty tracking was enabled for a memslot, we would previously clear all the dirty bits at that point (both in the HPT entries and in the rmap arrays), meaning that a KVM_GET_DIRTY_LOG ioctl immediately following would show no pages as dirty (assuming no vcpus have run in the meantime). With this change, the dirty bits on HPT entries are not cleared at the point where dirty tracking is enabled, so KVM_GET_DIRTY_LOG would show as dirty any guest pages that are resident in the HPT and dirty. This is consistent with what happens on radix. This also fixes a bug in the mark_pages_dirty() function for radix (in the sense that the function no longer exists). In the case where a large page of 64 normal pages or more is marked dirty, the addressing of the dirty bitmap was incorrect and could write past the end of the bitmap. Fortunately this case was never hit in practice because a 2MB large page is only 32 x 64kB pages, and we don't support backing the guest with 1GB huge pages at this point. Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
2017-10-26 13:39:19 +08:00
}
}
return 0;
}
void kvmppc_radix_flush_memslot(struct kvm *kvm,
const struct kvm_memory_slot *memslot)
{
unsigned long n;
pte_t *ptep;
unsigned long gpa;
unsigned int shift;
gpa = memslot->base_gfn << PAGE_SHIFT;
spin_lock(&kvm->mmu_lock);
for (n = memslot->npages; n; --n) {
ptep = __find_linux_pte(kvm->arch.pgtable, gpa, NULL, &shift);
if (ptep && pte_present(*ptep))
kvmppc_unmap_pte(kvm, ptep, gpa, shift, memslot,
kvm->arch.lpid);
gpa += PAGE_SIZE;
}
/*
* Increase the mmu notifier sequence number to prevent any page
* fault that read the memslot earlier from writing a PTE.
*/
kvm->mmu_notifier_seq++;
spin_unlock(&kvm->mmu_lock);
}
static void add_rmmu_ap_encoding(struct kvm_ppc_rmmu_info *info,
int psize, int *indexp)
{
if (!mmu_psize_defs[psize].shift)
return;
info->ap_encodings[*indexp] = mmu_psize_defs[psize].shift |
(mmu_psize_defs[psize].ap << 29);
++(*indexp);
}
int kvmhv_get_rmmu_info(struct kvm *kvm, struct kvm_ppc_rmmu_info *info)
{
int i;
if (!radix_enabled())
return -EINVAL;
memset(info, 0, sizeof(*info));
/* 4k page size */
info->geometries[0].page_shift = 12;
info->geometries[0].level_bits[0] = 9;
for (i = 1; i < 4; ++i)
info->geometries[0].level_bits[i] = p9_supported_radix_bits[i];
/* 64k page size */
info->geometries[1].page_shift = 16;
for (i = 0; i < 4; ++i)
info->geometries[1].level_bits[i] = p9_supported_radix_bits[i];
i = 0;
add_rmmu_ap_encoding(info, MMU_PAGE_4K, &i);
add_rmmu_ap_encoding(info, MMU_PAGE_64K, &i);
add_rmmu_ap_encoding(info, MMU_PAGE_2M, &i);
add_rmmu_ap_encoding(info, MMU_PAGE_1G, &i);
return 0;
}
int kvmppc_init_vm_radix(struct kvm *kvm)
{
kvm->arch.pgtable = pgd_alloc(kvm->mm);
if (!kvm->arch.pgtable)
return -ENOMEM;
return 0;
}
static void pte_ctor(void *addr)
{
memset(addr, 0, RADIX_PTE_TABLE_SIZE);
}
static void pmd_ctor(void *addr)
{
memset(addr, 0, RADIX_PMD_TABLE_SIZE);
}
struct debugfs_radix_state {
struct kvm *kvm;
struct mutex mutex;
unsigned long gpa;
int lpid;
int chars_left;
int buf_index;
char buf[128];
u8 hdr;
};
static int debugfs_radix_open(struct inode *inode, struct file *file)
{
struct kvm *kvm = inode->i_private;
struct debugfs_radix_state *p;
p = kzalloc(sizeof(*p), GFP_KERNEL);
if (!p)
return -ENOMEM;
kvm_get_kvm(kvm);
p->kvm = kvm;
mutex_init(&p->mutex);
file->private_data = p;
return nonseekable_open(inode, file);
}
static int debugfs_radix_release(struct inode *inode, struct file *file)
{
struct debugfs_radix_state *p = file->private_data;
kvm_put_kvm(p->kvm);
kfree(p);
return 0;
}
static ssize_t debugfs_radix_read(struct file *file, char __user *buf,
size_t len, loff_t *ppos)
{
struct debugfs_radix_state *p = file->private_data;
ssize_t ret, r;
unsigned long n;
struct kvm *kvm;
unsigned long gpa;
pgd_t *pgt;
struct kvm_nested_guest *nested;
pgd_t pgd, *pgdp;
pud_t pud, *pudp;
pmd_t pmd, *pmdp;
pte_t *ptep;
int shift;
unsigned long pte;
kvm = p->kvm;
if (!kvm_is_radix(kvm))
return 0;
ret = mutex_lock_interruptible(&p->mutex);
if (ret)
return ret;
if (p->chars_left) {
n = p->chars_left;
if (n > len)
n = len;
r = copy_to_user(buf, p->buf + p->buf_index, n);
n -= r;
p->chars_left -= n;
p->buf_index += n;
buf += n;
len -= n;
ret = n;
if (r) {
if (!n)
ret = -EFAULT;
goto out;
}
}
gpa = p->gpa;
nested = NULL;
pgt = NULL;
while (len != 0 && p->lpid >= 0) {
if (gpa >= RADIX_PGTABLE_RANGE) {
gpa = 0;
pgt = NULL;
if (nested) {
kvmhv_put_nested(nested);
nested = NULL;
}
p->lpid = kvmhv_nested_next_lpid(kvm, p->lpid);
p->hdr = 0;
if (p->lpid < 0)
break;
}
if (!pgt) {
if (p->lpid == 0) {
pgt = kvm->arch.pgtable;
} else {
nested = kvmhv_get_nested(kvm, p->lpid, false);
if (!nested) {
gpa = RADIX_PGTABLE_RANGE;
continue;
}
pgt = nested->shadow_pgtable;
}
}
n = 0;
if (!p->hdr) {
if (p->lpid > 0)
n = scnprintf(p->buf, sizeof(p->buf),
"\nNested LPID %d: ", p->lpid);
n += scnprintf(p->buf + n, sizeof(p->buf) - n,
"pgdir: %lx\n", (unsigned long)pgt);
p->hdr = 1;
goto copy;
}
pgdp = pgt + pgd_index(gpa);
pgd = READ_ONCE(*pgdp);
if (!(pgd_val(pgd) & _PAGE_PRESENT)) {
gpa = (gpa & PGDIR_MASK) + PGDIR_SIZE;
continue;
}
pudp = pud_offset(&pgd, gpa);
pud = READ_ONCE(*pudp);
if (!(pud_val(pud) & _PAGE_PRESENT)) {
gpa = (gpa & PUD_MASK) + PUD_SIZE;
continue;
}
if (pud_val(pud) & _PAGE_PTE) {
pte = pud_val(pud);
shift = PUD_SHIFT;
goto leaf;
}
pmdp = pmd_offset(&pud, gpa);
pmd = READ_ONCE(*pmdp);
if (!(pmd_val(pmd) & _PAGE_PRESENT)) {
gpa = (gpa & PMD_MASK) + PMD_SIZE;
continue;
}
if (pmd_val(pmd) & _PAGE_PTE) {
pte = pmd_val(pmd);
shift = PMD_SHIFT;
goto leaf;
}
ptep = pte_offset_kernel(&pmd, gpa);
pte = pte_val(READ_ONCE(*ptep));
if (!(pte & _PAGE_PRESENT)) {
gpa += PAGE_SIZE;
continue;
}
shift = PAGE_SHIFT;
leaf:
n = scnprintf(p->buf, sizeof(p->buf),
" %lx: %lx %d\n", gpa, pte, shift);
gpa += 1ul << shift;
copy:
p->chars_left = n;
if (n > len)
n = len;
r = copy_to_user(buf, p->buf, n);
n -= r;
p->chars_left -= n;
p->buf_index = n;
buf += n;
len -= n;
ret += n;
if (r) {
if (!ret)
ret = -EFAULT;
break;
}
}
p->gpa = gpa;
if (nested)
kvmhv_put_nested(nested);
out:
mutex_unlock(&p->mutex);
return ret;
}
static ssize_t debugfs_radix_write(struct file *file, const char __user *buf,
size_t len, loff_t *ppos)
{
return -EACCES;
}
static const struct file_operations debugfs_radix_fops = {
.owner = THIS_MODULE,
.open = debugfs_radix_open,
.release = debugfs_radix_release,
.read = debugfs_radix_read,
.write = debugfs_radix_write,
.llseek = generic_file_llseek,
};
void kvmhv_radix_debugfs_init(struct kvm *kvm)
{
kvm->arch.radix_dentry = debugfs_create_file("radix", 0400,
kvm->arch.debugfs_dir, kvm,
&debugfs_radix_fops);
}
int kvmppc_radix_init(void)
{
unsigned long size = sizeof(void *) << RADIX_PTE_INDEX_SIZE;
kvm_pte_cache = kmem_cache_create("kvm-pte", size, size, 0, pte_ctor);
if (!kvm_pte_cache)
return -ENOMEM;
size = sizeof(void *) << RADIX_PMD_INDEX_SIZE;
kvm_pmd_cache = kmem_cache_create("kvm-pmd", size, size, 0, pmd_ctor);
if (!kvm_pmd_cache) {
kmem_cache_destroy(kvm_pte_cache);
return -ENOMEM;
}
return 0;
}
void kvmppc_radix_exit(void)
{
kmem_cache_destroy(kvm_pte_cache);
kmem_cache_destroy(kvm_pmd_cache);
}