CRACK Windows 10 Pro X64 Build 14393 It-IT Sep 2016 {Generation2}
Elfreda Dasovich
Windows 10 installation on MAC book 2016 gets stuck on boot screen with black screen, blue windows logo and spinner. Windows 10 Installation was working perfectly fine until I enabled hyper v, now I can't start machine in safe mode because it needs function keys to boot machine in safe mode. on option seems to be reinstall the windows OS.
And it's the disabling of VT extensions in Mac EFI that's been causing Hyper-V installations to fail, according to most reports. I did a clean Boot Camp install (fresh partition) of Windows 10 Pro using a recent ISO downloaded from Microsoft. Once the Apple Boot Camp driver installation completed in Windows, but before joining a WiFi network or installing any MS Updates, I went to "Turn Windows features on or off", enabled Hyper-V, and allowed it to reboot.
It appears that you did not read the question right or did not get the context as both Parallels and Fusion are MacOS only virtualization software and this question pertains to Windows 10 running Bootcamp.
I have the same issue. I have a MBP 2016 13 inch (no touch bar). I have used Boot Camp to install Windows 10 Pro (10.0.14393) and then attempted to turn on Hyper V from programs and features, which installs and reboots once and during the second reboot it fails at the windows 10 boot screen. All Hyper V requirements are met:
That's a good question Loner T. I am doing a deep dive into Windows 10 UEFI support, changing UEFI settings for different windows boot options and different windows partition types to see if there's a way out of this.
After the introduction of ASLR, the advancement in mitigations against memory corruption has been impeded due to the implementation of sufficient mitigations across all platforms in the early 2000s. However, over time, offensive researchers have continuously presented bypasses and limitations of these initial generation mitigation techniques. Some of the primary limitations that persist include
These challenges have prompted the security industry to shift its focus towards mitigations that play a crucial role in preventing the exploitation of memory corruption, rather than solely detecting or preventing the corruption itself. This is particularly important in scenarios where the first generation techniques fail, whether due to the aforementioned issues or any other reasons.
In response to these requirements, the security industry has primarily concentrated on the concept of CFI (Control Flow Integrity). CFI serves as the fundamental principle behind all second generation techniques, which will be discussed below.
Attackers seek to inject code into a program to make use of its privileges or to extract data from its memory space. Before executable code is commonly made read-only, an attacker could arbitrarily change the code as it is run, targeting direct transfers or even do with no transfers at all. After W^X became widespread, an attacker wants to instead redirect execution to a separate, unprotected area containing the code to be run, making use of indirect transfers: one could overwrite the virtual table for a forward-edge attack or change the call stack for a backward-edge attack (return-oriented programming). CFI is designed to protect indirect transfers from going to unintended locations.
For the above high level implementation, Forward edge integrity will take care that function f1 will only call f2 or f2 is only called from f1. Whereas Backward edge integrity will take care that ret or function exit from f2 will only return to f1. For a successful CFI, both forward and backward edges would need to be protected to maintain control flow integrity and prevent attackers from diverting the program's execution in any direction.
In early stage CFI implementations like CCFIR and bin-CFI, Every instruction that is the target of a legitimate control-flow transfer is assigned a unique identifier (ID), and checks are inserted before control-flow instructions to ensure that only valid targets are allowed. Direct transfers have a fixed target and they do not require any enforcement checks. However, indirect transfers, like function calls and returns, and indirect jumps, take a dynamic target address as argument. As the target address could be controlled by an attacker due to a vulnerability, CFI checks to ensure that its ID matches the list of known and allowable target IDs of the instruction. The implementation can be understood better using below control flow graph:
In the year 2005, the initial surface of CFI research became publicly accessible. Subsequently, numerous additional researchers emerged over the course of several years, presenting evidence of various forms of control flow integrity. However, the primary release of CFI that was readily accessible occurred in 2014 for both Windows and Linux operating systems, with both the compilers of Windows and Linux incorporating some level of support for CFI. In 2014, Linux implemented its first CFI implementation, drawing upon the research conducted by the Google team as outlined in their paper titled "Enforcing Forward-Edge Control-Flow Integrity in GCC & LLVM". Following suit, Windows also introduced CFI support in November 2014, referring to it as Control Flow Guard. Given that the Linux implementation gained public exposure first, let us delve into the CFI implementation within the Linux system.
In 2014, few google researchers have implemented the first practical CFI implementation. This particular implementation was employed internally in conjunction with Chromium and a select few other products, prior to its public release in August of that year. Consequently, this CFI implementation was integrated into both the GCC and LLVM compilers during the same aforementioned year. This initial work contain two different methods for CFI:
Both methods have been implemented in order to enforce the integrity of forward edges and primarily focus on safeguarding against arbitrary code execution resulting from memory corruption in the heap. The authors have observed that, although the process stack has already been adequately protected through the implementation of multiple mitigation measures up until 2014, the heap area of the process remains susceptible to memory corruptions, which can lead to arbitrary code execution without any significant mitigations in place. This has served as a motivation for the authors to incorporate compiler-based mechanisms, with the aim of further enhancing the protection and integrity of program control data, with particular emphasis on the integrity of control-transfer data stored in the heap.
The goal of VTV is to protect virtual calls in C++. The motivation behind this mitigation is to protect common types of code hijacking that happen in C++ based memory implementation due to heap exploitations.
To protect the above scenario, VTV verifies the validity, at each call site, of the vtable pointer being used for the virtual call, before allowing the call to execute. In particular, it verifies that the vtable pointer about to be used is correct for the call site, i.e., that it points either to the vtable for the static type of the object, or to a vtable for one of its descendant classes. The compiler passes to the verifier function the vtable pointer from the object and the set of valid vtable pointers for the call site. If the pointer from the object is in the valid set, then it gets returned and used. Otherwise, the verification function calls a failure function, which normally reports an error and aborts execution immediately.
VTV has two pieces: the main compiler part, and a runtime library (libvtv), both of which are part of GCC. In addition to inserting verification calls at each call site, the compiler collects class hierarchy and vtable information during compilation, and uses it to generate function calls into libvtv, which will (at runtime) build the complete sets of valid vtable pointers for each polymorphic class in the program.
To keep track of static types of objects and to find sets of vtable pointers, VTV creates a special set of variables called vtable-map variables (read only), one for each polymorphic class. At runtime, a vtable-map variable will point to the set of valid vtable pointers for its associated class. When VTV inserts a verification call, it passes in the appropriate vtable-map variable for the static type of the object, which points to the set to use for verification.
In the second image you will notice that before calling a1->getmass() (address [rax+8]), it is been verified if it is in the range of valid call site which is added during IR phase of compilation. The check can be explained in more details here:
Unlike VTV, IFCC mechanism protect integrity of all kinds of indirect calls. , it protects indirect calls by generating jump tables for indirect-call targets and adding code at indirect-call sites to transform function pointers, ensuring that they point to a jump-table entry. Any function pointer that does not point into the appropriate table is considered a CFI violation and will be forced into the right table by IFCC. The valid jump table entries are again created based on same function prototypes.
In the second case where IFCC is enabled, before calling add through operation (call rax), the address is first verified if matches the address in the jmp table ( mov rcx, offset add and cmp rax, rcx) or not. Since we only have 1 address in the jump table, the check is straight forward but in case of multiple entries, you will see ror (rotate) instruction to verify if the jump index is within range.
-fsanitize=cfi-cast-strict - If a class has a single non-virtual base and does not introduce or override virtual member functions or fields other than an implicitly defined virtual destructor, it will have the same layout and virtual function semantics as its base. By default, casts to such classes are checked as if they were made to the least derived such class.
-fsanitize=cfi-derived-cast and -fsanitize=cfi-unrelated-cast - These scheme checks that pointer casts are made to an object of the correct dynamic type; that is, the dynamic type of the object must be a derived class of the pointee type of the cast. The checks are currently only introduced where the class being casted to is a polymorphic class. First one is bad casts from a base class to a derived class and 2nd one is bad casts from a pointer of type void* or another unrelated type.
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