Why the JVM Spec defines checkcast for interface types

I'm working on the specification of pointer analysis for Java using Datalog. Basically, a pointer analysis computes for each variable in a program the set of objects it may point to at run-time.

For this purpose I need to express parts of the JVM Spec in Datalog as well. As a simple example, the following Datalog rules define when a class is a subclass of another class.

/**
 * JVM Spec:
 * - A class A is a subclass of a class C if A is a direct 
 *   subclass of C
 */
Subclass(?c, ?a) <-
  DirectSubclass[?a] = ?c.

/**
 * JVM Spec:
 * - A class A is a subclass of a class C if there is a direct
 *   subclass B of C and class A is a subclass of B
 */
Subclass(?c, ?a) <-
  Subclass(?b, ?a),
  DirectSubclass[?b] = ?c.

As you can see, this is remarkably close to the original specification (quoted in comments). You can clearly see the relationship between the spec and the code, even if you are not familiar with Datalog.

Recently, I was working on the specification of the checkcast instruction. This instruction performs the run-time check if an object can be cast to some type. The JVM Spec for checkcast first defines some variables:

The following rules are used to determine whether an objectref that is not null can be cast to the resolved type: if S is the class of the object referred to by objectref and T is the resolved class, array, or interface type, checkcast determines whether objectref can be cast to type T as follows:

So, this basically says that we're checking the cast (T) S.

The first rule for this cast is straightforward:

If S is an ordinary (nonarray) class, then:

• If T is a class type, then S must be the same class as T, or a subclass of T.
• If T is an interface type, then S must implement interface T.

Well, if you're somewhat familiar with Java, or object-oriented programming, then this part is obvious. Again, the specification in Datalog is easy:

CheckCast(?s, ?s) <-
  ClassType(?s).

CheckCast(?s, ?t) <-
  Subclass(?t, ?s).

CheckCast(?s, ?t) <-
  ClassType(?s),
  Superinterface(?t, ?s).

However, the next alternative in the specification is confusing:

If S is an interface type, then:

• If T is a class type, then T must be Object.
• If T is an interface type, then T must be the same interface as S or a superinterface of S.

The specification is crystal clear, but how can S ever be an interface type? S is the type of the object that is being cast, and how can an object ever have a run-time type that is an interface? Of course, the static type of an expression can be an interface, but we're talking about the run-time here!

I searched the web, which only resulted in a few hits. There was one question on a Sun forum years ago, where the one answer didn't make a lot of sense.

It turns out that this is indeed an `impossible' case. The reason why this item is in the specification, is because checkcast is recursively defined for arrays:

If S is a class representing the array type SC[], that is, an array of components of type SC, then:

• ...
• If T is an array type TC[], that is, an array of components of type TC, then one of the following must be true:
  • ...
  • TC and SC are reference types, and type SC can be cast to TC by recursive application of these rules.

So, if you have an object of type List[] that is cast to an Collection[], then the rules for checkcast get recursively invoked for the types S = List and T = Collection. Notice that List is an interface, but an object can have type List[] at run-time. If have not verified this with the JVM Spec maintainers, but as far as I can see, this is the only reason why the rule for interface types is there.

Just to show a little bit more of my specifications, here is the rule for the array case I just quoted from the JVM Spec:

CheckCast(?s, ?t) <-
  ComponentType[?s] = ?sc,
  ComponentType[?t] = ?tc,
  ReferenceType(?sc),
  ReferenceType(?tc),
  CheckCast(?sc, ?tc).

Isn't it beautiful how this exactly corresponds to the formal specification?

Unfortunately, even formal specifications can have errors, so I also specified a large testsuite that checks the specifications with concrete code. Here are some of the tests for CheckCast.

test Casting to self
  using database tests/hello/Empty.jar
  assert
    CheckCast("java.lang.Integer", "java.lang.Integer")

test Casting to superclasses
  using database tests/hello/Empty.jar
  assert
    CheckCast("java.lang.Integer", "java.lang.Number")
    CheckCast("java.lang.Integer", "java.lang.Object")

test Cast ArrayList to various superinterfaces
  using database tests/hello/Arrays.jar
  assert
    CheckCast("java.util.ArrayList", "java.util.List")
    CheckCast("java.util.ArrayList", "java.util.Collection")
    CheckCast("java.util.ArrayList", "java.io.Serializable")

test Cast class[] to implemented interface[]
  using database tests/hello/Arrays.jar
  assert
    CheckCast("java.util.ArrayList[]", "java.util.List[]")
    CheckCast("java.lang.Integer[]", "java.io.Serializable[]")

test Cast interface[] to superinterface[]
  using database tests/hello/Arrays.jar
  assert
    CheckCast("java.util.List[]", "java.util.Collection[]")

The tests are specified in a little domain-specific language for unit-testing Datalog that I implemented, initially for IRIS and later for LogicBlox. This tool is similar to parse-unit, a tool I wrote earlier for testing parsers in Stratego/XT. The concise syntax of a test encourages you to write a lot of tests. Domain-specific languages rock for this purpose!

New Conference on Software Language Engineering

This year is special. There is a new and exciting conference: the International Conference on Software Language Engineering (SLE). The deadline for submission of papers is July 14th, which is coming up soon! Before I start raving about the topics covered by this conference, here is the disclaimer: I'm on the program committee of this conference, and as such I believe it's my duty to advertise the conference.

Anyway, if done right, this conference has the potential to become a major and prestigious conference. The conference fills a clear gap: the topics of software language engineering do not exactly fit in major programming language conferences like OOPSLA, PLDI, POPL, and ECOOP. Nor do they fit exactly in the area of compiler construction (CC). CC does typically not accept more engineering or methodology-oriented papers. For OOPSLA and ECOOP the work more or less has to be in the context of object-oriented programming, for POPL it immediately has to be a principle (whatever that is), and for PLDI there are usually just a few slots available for papers that don't do something with memory management, garbage collection, program analysis, or concurrency. Personally, I've been pretty successful at getting papers in the area of software language engineering accepted at OOPSLA, but a full conference devoted to this topic is much better!

Another reason why I think that this conference has a lot of potential is that if I look at the list of topics of interest in the call for papers, then I can only think of one summary: everything that's fun! I'm convinced I'm not the only one who thinks these topics are fun. When talking to colleagues, I notice again and again most of us just love languages. The engineering of those languages is an issue for almost all computer scientists and many programmers in industry, and this conference will be the most obvious target for papers about this!

Also, the formalisms and used for the specification and implementation of (domain-specific) languages are still very much an open research topic. Standardization of languages is still far from perfect, as discussed by many posts on this blog. Also, new language implementation techniques are being proposed all the time, and extensible compilers for developing language extensions are more popular than ever. Not to mention the increasing interest in using domain-specific languages to help solve the software development problems we're facing.

Earlier in this post I wrote that this conference has major potential if done right. There are few risks. First, the conference has been started by two relatively small communities: ATEM and LDTA. I think the conference should attract a much larger community than the union of those two communities. I hope lots of people outside of the ATEM and LDTA communities will consider to submit a paper. Second, this year the conference is co-located with MODELS. Many programming language people are slightly allergic to model-driven engineering. I hope they will realize that this conference is not specifically a model-driven conference. Finally, the whole setup of the conference should be international and varied. I'm sorry to say that at this point I'm not entirely happy with the choice of keynote speakers. This nothing personal: I respect both keynote speakers, but the particular combination of the two speakers is a bit unfortunate. First, they are both Dutch. Second, neither of them is extremely well-known in the communities of OOPSLA, PLDI, or ECOOP. I hope that this will not affect the potential of this interesting conference.

Now go work on your submission!

Ph.D. Thesis: Exercises in Free Syntax

It has been awfully quiet here, I'm sorry about that. There are a few reasons for that. The first one is that I assembled my PhD thesis from my publications. This took quite some time and energy, but the result is great! My dissertation Exercises Free Syntax is available online. If you are interested in having dead tree version, just let me know!

I will defend my thesis tomorrow, January 21 (see the Dutch announcement). It's weird to realize that tomorrow is the accumulation of 4 years of working intensely!

For the library I created an English abstract. To give you an idea what the thesis is about, let me quote it here:

In modern software development the use of multiple software languages to constitute a single application is ubiquitous. Despite the omnipresent use of combinations of languages, the principles and techniques for using languages together are ad-hoc, unfriendly to programmers, and result in a poor level of integration. We work towards a principled and generic solution to language extension by studying the applicability of modular syntax definition, scannerless parsing, generalized parsing algorithms, and program transformations.

We describe MetaBorg, a method for providing concrete syntax for domain abstractions to application programmers. Since object-oriented languages are designed for extensibility and reuse, the language constructs are often sufficient for expressing domain abstractions at the semantic level. However, they do not provide the right abstractions at the syntactic level. The MetaBorg method consists of embedding domain-specific languages in a general purpose host language and assimilating the embedded domain code into the surrounding host code. Instead of extending the implementation of the host language, the assimilation phase implements domain abstractions in terms of existing APIs leaving the host language undisturbed.

We present a solution to injection vulnerabilities. Software written in one language often needs to construct sentences in another language, such as SQL queries, XML output, or shell command invocations. This is almost always done using unhygienic string manipulation. A client can then supply specially crafted input that causes the constructed sentence to be interpreted in an unintended way, leading to an injection attack. We describe a more natural style of programming that yields code that is impervious to injections by construction. Our approach embeds the grammars of the guest languages into that of the host language and automatically generates code that maps the embedded language to constructs in the host language that reconstruct the embedded sentences, adding escaping functions where appropriate.

We study AspectJ as a typical example of a language conglomerate, i.e. a language composed of a number of separate languages with different syntactic styles. We show that the combination of the lexical syntax leads to considerable complexity in the lexical states to be processed. We show how scannerless parsing elegantly addresses this. We present the design of a modular, extensible, and formal definition of the lexical and context-free aspects of the AspectJ syntax. We introduce grammar mixins, which allows the declarative definition of keyword policies and combination of extensions.

We introduce separate compilation of grammars to enable deployment of languages as plugins to a compiler. Current extensible compilers focus on source-level extensibility, which requires users to compile the compiler with a specific configuration of extensions. A compound parser needs to be generated for every combination. We introduce an algorithm for parse table composition to support separate compilation of grammars to parse table components. Parse table components can be composed (linked) efficiently at runtime, i.e. just before parsing. For realistic language combination scenarios involving grammars for real languages, our parse table composition algorithm is an order of magnitude faster than computation of the parse table for the combined grammars, making online language composition feasible.

Also, they asked me for a Dutch, non-technical summary for news websites. For my Dutch readers:

We presenteren een verzameling van methoden en technieken om programmeertalen te combineren. Onze methoden maken het bijvoorbeeld mogelijk om in een programmeertaal die ontworpen is voor algemene doeleinden een subtaal te gebruiken die beter aansluit bij het domain van een bepaald onderdeel van een applicatie. Hierdoor kan een programmeur op een duidelijkere en compactere wijze een aspect van de software implementeren.

Op basis van dezelfde technieken presenteren we een methode die programmeurs beschermt tegen fouten die de oorzaak zijn van het meest voorkomende beveiligingsprobleem, een zogenaamde injectie aanval. Door op een iets andere wijze te programmeren, heeft de programmeur de garantie dat de software niet gevoelig is voor dergelijke aanvallen. In tegenstelling tot eerder voorgestelde oplossingen geeft onze methode absolute garanties, is eenvoudiger voor de programmeur, en kan gebruikt worden voor alle gevallen waarin injectie aanvallen kunnen voorkomen (bijvoorbeeld niet specifiek voor de taal SQL).

Tot slot maken onze technieken het mogelijk om de syntaxis van sommige programmeertalen duidelijker en formeler te definieren. Sommige moderne programmeertalen zijn eigenlijk een samensmelting van verschillende subtalen (zogenaamde taalagglomeraten). Van dergelijke talen was het tot nu toe onduidelijk hoe de syntaxis precies geformuleerd kon worden, wat voor standaardisering en compatibiliteit noodzakelijk is.

LDTA'07 slides on Grammar Engineering Tools

The slides of our presentation of the LDTA'07 paper Grammar Engineering Support for Precedence Rule Recovery and Compatibility Checking are now available online. The slides are a spectacular demonstration of latex masochism, so please take a look ;) . There are few bonus slides after the conclusion that I wasn't able to show during the 30-minutes version of the talk.

Migration of the YACC grammar for PHP to SDF

Last summer, Eric Bouwers started working on infrastructure for PHP program transformation and analysis, sponsored by the Google Summer of Code. He did an excellent job, thanks to his expertise in PHP and his thorough knowledge of Stratego/XT. To enjoy all the language engineering support in Stratego/XT, Eric developed a PHP grammar in SDF, the grammar formalism that is usually applied in Stratego/XT projects. Unfortunately it proved to be very difficult to get the grammar of PHP right.

PHP precedence problems

PHP features many operators, and the precedence of the operators is somewhat unusual and challenging for a grammar formalism. For example, PHP allows the weak binding assignment operator as an argument of the binary, strong binding && operator:

  if ($is_upload && $file = fopen($fname, 'w')) { 
    ...
  }

The same holds for the unary, strong binding ! operator:

  if(!$foo = getenv('BAR')){
    ...
  }

A similar precedence rule for the include operator allows an include to occur as the argument of the strong binding @ operator:

  @include_once 'Var/Renderer/' . $mode . '.php'

Precedence rule recovery

The most serious problem was to find out what the exact precedence rules of PHP operators are. The syntax of PHP is defined by a YACC grammar, which has a notion of precedence declarations that is heavily used by the PHP grammar. Unfortunately, for more complex grammars it is far from clear what the exact effect of the precedence declarations are. The precedence declarations are only used for conflict resolution in the parse table generator, so if there is no conflict, then the precedence declarations do not actually have any effect on a particular combination of operators. That's why we developed support for recovering precedence rules from YACC grammars, which I already wrote about in a previous blog. Based on these tools, we now have a very precise specification of the precedence rules of PHP.

The next step in the process of getting a perfect PHP grammar was to actually use this specification to develop very precise precedence declarations for the SDF grammar of PHP. However, the precedence rule specification involves about 1650 rules, so migrating these precedence rules to SDF precedence declarations by hand is not really an option. Fortunately, all the ingredients are actually there to generate SDF priority declarations from the precedence rules that we recover from the YACC grammar.

Argument-specific priorities

Thanks to two new features of SDF, these precedence rules can be translated directly to SDF. The first feature is argument-specific priorities. In the past, SDF only allowed priority declarations between productions. For example, the SDF priority

  E "*" E -> E > E "+" E -> E

defines that the production for the + operator cannot be applied to produce any of the E arguments of the production for the * operator, hence the production for the addition operator cannot be applied on the left-hand side or right-hand side of the multiplication operator. This priority implies that the multiplication operator binds stronger than the addition operator. This single SDF priority corresponds to the following two precedence rules in the grammar formalism independent notation we are using in the Stratego/XT Grammar Engineering Tools:

  <E -> <E -> E + E> * E>
  <E -> E * <E -> E + E>>

For many languages precedence rules are different for arguments of the same production. That's why we us the more specific representation of a precedence rules in our grammar engineering tools. Fortunately, SDF now supports argument-specific priorities as well. These argument-specific priorities are just plain numbers that indicate to which arguments of a production the priority applies. For example, the following SDF priority forbids the assignment operator only at the left-most and the right-most E of the conditional operator:

  E "?" E ":" E -> E <0,4> > E "=" E

This corresponds to the following precedence rules:

  <E -> <E -> E = E> ? E : E>
  <E -> E ? E : <E -> E = E>>

Non-transitive priorities

The second new SDF feature that is required for expressing the PHP precedence rules is non-transitive priorities. Before the introduction of this feature, all SDF priorities where transitively closed. For example, if there are two separate priorities

  "!" E -> E   > E "+" E -> E
  E "+" E -> E > V "=" E -> E

then by the transitive closure of priorities this would imply the priority

  "!" E -> E >  V "=" E -> E

This transitive closure feature is useful in most cases, but for some languages (such as PHP) the precedence rules are in fact not transitively closed, which makes the definition of these rules in SDF slightly problematic. For this reason, SDF now also features non-transitive priorities, using a dot before the > of the priority:

  "!" E -> E .> E "+" E -> E

Non-transitive priorities will not be included in the transitive closure, which gives you very precise control over the precedence rules.

Precedence rule migration

Thanks to the position-specific, non-transitive priorities of SDF, the precedence rules that we recover from the YACC grammar for PHP can now be mapped directly to SDF priority declarations. The two precedence rules mentioned earlier:

  <E -> <E -> E + E> * E>
  <E -> E * <E -> E + E>>

now translate directly to SDF priorities:

  E * E -> E <0> .> E + E -> E
  E * E -> E <2> .> E + E -> E

The migration of the recovered YACC precedence rules results in about 1650 of these SDF priorities, but thanks to the fully automatic migration this huge number of priorities is not really a problem. The resulting PHP syntax definition immediately solved all the known issues with the PHP syntax definition, which shows that this migration was most reliable and successful.

Future

There is a lot of interesting work left to be done. First, it would be interesting to develop a more formal grammar for PHP, similar to the grammars of the C, Java, and C# specifications. These specifications all encode the precedence rules of the operators in the production rules, by introducing non-terminals for all the precedence levels. It should not be too difficult to automatically determine such an encoding from the precedence rules we recover. This would result in a formal specification of the PHP syntax, which will benefit many other parser generators. One of the remarkable things we found out is that the unary - operator has the same precedence as the binary - (usually it binds stronger), which results in -1 * 3 being parsed as -(1 * 3). We have not been able to find an example where this strange precedence rule results in unexpected behaviour, but for the development of a solid parser is it essential that such precedence rules are defined precisely.

Second, it would be good to try to minimize the number of generated SDF priorities by determining a priority declaration that you can actually oversee as a human. This would involve finding out where the transitive closure feature of SDF priorities can be used to remove redundant priority declarations.

Third, it would great to integrate the precedence rule migration in a tool that completely migrates a YACC/FLEX grammar to SDF. For this, we need tools to parse and understand a FLEX specification and extend the existing support for precedence rule migration to other YACC productions.

Clearly, there is lots of interesting (and useful!) grammar engineering work to do in this direction!

x86-64 support for Stratego/XT!

Today is the day that Stratego/XT supports 64-bit processors! Stratego/XT supports x86-64 from release 0.17M3pre16744 (or later), the sdf2-bundle from release 2.4pre212034 (or later). The releases are available from our new Nix buildfarm at the TU Delft.

Some history

About 6 years ago, various people started to complain about the lack of 64-bit processor support. At the time, most complaints came from our very own Unix geek Armijn Hemel, mostly because of his passion for Sun and these strange UltraSparc machines. However, similar to the limited distribution of Unix geeks, 64-bit systems were rather uncommon at the time. The requests we got were about more obscure processors, like Sun's UltraSparc and Intel's IA-64 (Itanium).

The 64-bit issues were never solved because (1) we never had a decent 64-bit machine at our disposal, (2) users with 64-bit system were uncommon, and (3) most of the issues were actually not Stratego/XT issues, but problems in the ATerm library, which is not maintained by the Stratego/XT developers.

Some first steps

However, it is not possible anymore to ignore 64 bit systems: Intel and AMD both sell 64-bit processors for consumers these days. Several users of Stratego/XT already have x86-64 machines, and the only reason why they don't complain en masse is that there is always the option to compile in 32-bit mode (using gcc -m32).

At the TU Delft (the new Stratego/XT headquarters), we now have an amazing buildfarm with some real, dedicated hardware bought specifically for the purpose of building software. At the moment, all our build machines (except for the Mac Minis) have x86-64 processors, so the lack of 64-bit machines is no longer an excuse.

Also, the ATerm library now enjoys a few more contributors. Last summer, Eelco Dolstra from Utrecht University created the first complete 64-bit patch for the ATerm library (Meta-Environment issue 606), simply because his Nix package management system uses the ATerm library and portability of Nix is important. Also, Erik Scheffers from the Eindhoven University of Technology has done an excellent job on the development of ATerm branches that support GCC 4.x and 64-bit machines.

The final step .. uh, steps

As a result, it was now feasible to fully support x86-64 systems. The only thing left for me to do was to use all the right patches and branches and enable an x86-64 build in our buildfarm. At least, that's what I thought ... Well, if you know computer scientists, then you'll also know that they are always far too optimistic.

In the end, it took me a day or four to get everything working. This is rather stressful work, I must say. Debugging code that is affected by a mixture of 32-bit assumptions and aliasing bugs introduced by GCC optimizations is definitely not much fun. You can stare at C code for as long as you like, but if the actual code being executed is completely different, then this won't help much. In the end, this little project resulted in quite a few new issues:

• STR-701 is a bug that was raised by casting a pointer to an integer in the address strategy of libstratego-lib, which returns an integer representation of the address of an ATerm. The Stratego Library has had this strategy for a long time, and indeed the most natural representation of an address is an integer datatype. Unfortunately, ATerm integers are fixed size, 32-bit integers, hence it cannot be used to represent a pointer of 64 bits. The new representation is a string, which is acceptable for most of the applications of address.

• Meta-Environment issue 720 is related to GCC optimizations based on strict alias analysis. In this case, the optimization seems to be applied only in the x86-64 backend of GCC, while the underlying problem is in fact architecture independent.

  The code that raises this bug applies efficient memory allocation by allocating blocks of objects rather than individual ones. The available objects are encoded efficiently in a linked list, with only a next field. This next field is used for the actual data of the object, as well as the link to next available object. The objects are character classes, having the name CC_Class, which is a typedef for an array of longs. Roughly, the invalid code for adding a node to the linked list looks like this:

  struct CC_Node {
    struct CC_Node *next;
  };
  
  static struct CC_Node *free_nodes = NULL;
  
  void add_node(CC_Class* c) {
    struct CC_Node *node = (struct CC_Node *) c;
    node->next = free_nodes;
    free_nodes = node;
  }
  

  The problem with this efficient linked list is that the same memory location is accessed through pointers of different types, in this case a pointer to a CC_Node struct and a pointer to a CC_Class. Hence, the code creates aliases of different types, which is invalid in C (see for example this nice introduction to strict aliasing). In this case, C compilers are allowed to assume that the two variables do not alias, which enables a whole bunch of optimizations that are invalid if they do in fact alias.

  The solution for this is to use a C union, which explicitly informs the compiler that a certain memory location is accessed through two different types. Using a union, the above code translates to:

  union CC_Node {
    CC_Class *cc;
    CC_Class **next;
  };
  
  static union CC_Node free_node = {NULL};
  
  void add_node(CC_Class* c) {
    node.cc = c;
    *(node.next) = free_node.cc;
    free_node.cc = node.cc;
  }
  

  Sidenote: I'm not really a C union expert, and I'm not 100% sure whether in this case a union is necessary for a CC_Class* and CC_Class** or CC_Class and CC_Class*. The union I've chosen solves the bug, but I should figure out what the exact solution should be. Feedback is welcome.

• Meta-Environment issue 718 is related to the previous bug. The problem here is that the same memory locations are accessed through a generic datatype (ATerm) as well as pointers to more specific structs, which again leads to strict aliasing problems. This time, the issue has been solved in a more ad-hoc way by declaring a variable as volatile. This solves the issue for now, but a more fundamental solution (probably a union) is necessary here as well.

• STR-705 adds some checks for the size of various types to the Stratego/XT build system, called Auto/XT. These checks are necessary for the header files of the ATerm library, which determine the characteristics of the platform based on the size of longs, integers, and void pointers, which are defined as macros (a feature that is under discussion in Meta-Environment issue 606: it does not play very well with cross compilation and compilation in 32-bit mode on a 64-bit platform). The ATerm library we are using at the moment is the branch 64-bit-fixes, which has been developed by Eelco Dolstra and Erik Scheffers.

  The new macro XT_C_TYPE_CHARACTERISTICS checks the sizes and defines the macros that are required by these headers. The macro XT_SETUP invokes the XT_C_TYPE_CHARACTERISTICS macro, so all packages based on Stratego/XT will automatically support the 64-bit branch of the ATerm library.

• STR-703 is related to the previous issues. In packages based on the GNU Autotools and Auto/XT, the C code is compiled by the Automake-based build system, not by the Stratego compiler itself (which only produces the C code). In this case, the XT_C_TYPE_CHARACTERISTICS takes care of the required defines. However, the Stratego compiler can also be used as a standalone compiler, where strc invokes the C compiler itself. In this case, strc needs to pass the definitions of macros to the C compiler.

• STR-704 drops the use of autoheader in stratego-libraries. Autoheader replaces the command-line definition of macros to a generated config.h. This generated file used to be installed as stratego-config.h, but this header file is no longer necessary: there is no configuration option in this file that is still necessary as part of the Stratego/XT installation. The mechanism of config.h installation is rather fragile (some macro definitions have to be removed), so if it is not necessary anymore, then why not drop it ...

  The relation to x86-64 support is that several C files in the stratego-libraries package did not correctly include the generated config.h before aterm2.h. This breaks on x86-64 systems because aterm2.h requires the aforementioned macro definitions.

Short version

The net result of this operation is that we now support x86-64 systems. And this time we will keep supporting 64-bit processors, whatever it takes.

It would be fun to check now if UltraSparc and IA-64 machines work out of the box, but I don't have access to any of these. If you have one, I would love to know if it works.

Base access in the C# specification

In a previous post, I discussed a bug in the Java Language Specification on super field access of protected fields. If you haven't read this yet, I would suggest to give it a read before you continue with this post. Thanks to a discussion with Alex Buckley (the new maintainer of the Java Language specification), there is now a proposal to fix this bug in an elegant way. I'll report on the solution and the nice discussion on the relation to super field accesses in bytecode later.

However, first I would like to illustrate the risk of reuse. While writing on issues in the Java Language Specification, I figured that the C# specification probably has the same issue. After all, C# features the same details of protected access. Consider the following two C# classes:

  class A {
    protected int secret;
  }

  class B : A {
    public void f(A a) {
      a.secret = 5;
    }
  }

Due to the details of protected access, this example won't compile. The Mono C# compiler clearly explains the problem:

  A.cs(17,5): error CS1540: Cannot access protected 
  member  `A.secret' via a qualifier of type `A'. The 
  qualifier must be of type `B' or derived from it

Of course, C# also support access to fields of base classes (aka super classes). Indeed, checking the C# specification reveals that the definition of base access is exactly the same as super field access in Java. In Section 14.5.8 of the C# Language Specification (ECMA-334), the semantics of a base access expressions is defined in the following way:

"At compile-time, base-access expressions of the form base.I and base[E] are evaluated exactly as if they were written ((B)this).I and ((B)this)[E], where B is the base class of the class or struct in which the construct occurs."

Compare this definition to the Java Language Specification:

"Suppose that a field access expression super.name appears within class C, and the immediate super class of C is class S. Then super.name is treated exactly as if it had been the expression ((S)this).name."

The good thing about this reuse is that I can reuse the examples of my previous post as well. Consider the following two C# classes that compile without any problem:

  class A {
    protected int secret;
  }

  class B : A {
    public void f() {
      base.secret = 5;
    }
  }

Next, consider the derivative of this example, where class B has been modified to refer to the field secret using (A) this which is exactly the same as a reference through base, according to the specification.

  class B : A {
    public void f() {
      ((A) this).secret = 5;
    }
  }

Similar to Java, this class won't compile, due to the details of protected access in C#. Again, the Mono C# compiler explains the issue:

  A.cs(13,7): error CS1540: Cannot access protected 
  member `A.secret' via a qualifier of type `A'. The 
  qualifier must be of type `B' or derived from it

This example shows that for C# the two expressions base.secret and ((A) this).secret are not evaluated in the same way, so the previously reported problem in the Java Language Specification also applies to the C# specification.

Now I have to figure out how to report issues in the C# specification ...

Informal specifications are not so super

I'm trying to get back to the good habit of blogging about our work. I'm not very fond of dumping random links or remarks, so the most challenging part of blogging is to find a good topic to write a decent story about. This time, I fallback to a topic I've actually been working on about two years ago, but is still most relevant. At that time, I was actively developing a type checker for Java, as part of the Dryad project. This story is about a bug in the Java Language Specification, that for whatever bizarre reason has never been reported (afaik).

Super field access

Java supports access to fields of super classes using the super keyword, even if this field is hidden by a declaration of another field with the same name. For example, the following sample will print super.x = 1 and x = 2.

  class S {
    int x = 1;
  }

  class C extends S {
    int x = 2;

    void print() {
      System.out.println("super.x = " + super.x + " and x = " + x);
    }

    public static void main(String[] ps) {
      new C().print(); 
    }
  }

To allow access from inner classes to hidden fields of enclosing instances, Java also supports qualified super field accesses. In this case, the super keyword is prefixed with the name of a lexically enclosing class. This feature is related to the qualified this expression, which allows you to refer to an enclosing instance.

Current specification

We all have a reasonable, though informal, idea what the semantics of this language feature is. Of course for a real specification the semantics has to be defined more precisely. For example, two things that need to be define are what the type of such an expression is and if the field is accessible at all. The specification concisely defines the semantics of this language feature by forwarding the semantic rules to existing, more basic language features. For super.name, the JLS specifies:

"Suppose that a field access expression super.name appears within class C, and the immediate super class of C is class S. Then super.name is treated exactly as if it had been the expression ((S)this).name."

So, in the example I gave, super.x would be exactly equivalent to ((S)this).x. Obviously, the emphasis of exactly is on purpose. Why would they use this word? Does this suggest that there is also a notion of being treated almost exactly in the same way? ;) .

For qualified field access, the specification is almost the same, but this time using a qualified this instead of this.

"Suppose that a field access expression T.super.name appears within class C, and the immediate super class of the class denoted by T is a class whose fully qualified name is S. Then T.super.name is treated exactly as if it had been the expression ((S)T.this).name."

This specification looks very reasonable, considering that for a field access only the compile-time type of the subject expression is used to determine which field is to be used. By casting the subject expression (this) to the right type, the expected field of the super class is accessed.

Oops

Of course, it is always nice to have your type checker as compact as possible, so I was very happy with this specification. I could just forward everything related to super field accesses to the corresponding expression with a cast and a this expression. The Dryad typing rules looked something like this:

  attributes:
    |[ super.x ]| -> <attributes> |[ ((reft) this).x ]|
    where
      reft is the superclass of the current class

  attributes:
    |[ cname.super.x ]| ->  <attributes> |[ ((reft) cname.this).x ]|
    where
      reft is the superclass of the class cname

This implementation looks very attractive, but ... it didn't work. The reason for this is that super.name is in fact not exactly the same as ((S)this).name. The reason for this are the details of protected access, which I've previously written about on my blog. I'm not going to redo that, so let me just give an example (based on the example in the previous post) where this assumed equivalence is invalid. First, the following two classes are valid and can be compiled without any problems:

  package a;
  public class A {
    protected int secret;
  }

  package b;
  public class B extends a.A {
    void f() {
      super.secret = 5;
    }
  }

Next, let's now change the assignment to the expression ((a.A) this).secret, which is equivalent to super.secret according to the specification.

  package b;

  public class B extends a.A {
    void f() {
      ((a.A) this).secret = 5;
    }
  }

Unfortunately, this won't compile, due to the details of protected access:

b/B.java:5: secret has protected access in a.A
    ((a.A) this).secret = 5;
    ^
1 error

This example shows that the two expressions are not treated in the same way, so this looks like a problem in the Java Language Specification to me. Also, this shows that in semantics of languages likje Java the devil is really in the detail. What surprises me is that nobody has mentioned this before. Several major Java compilers have been implemented, right? Shouldn't the programmers responsible for these compilers have encountered this problem?

Java Virtual Machine

Another interesting thing is how the Java Virtual Machine specification deals with this. There is no special bytecode operator for accessing fields of super classes: all field assignments are performed using the putfield operator. Assuming that the source compiler would ignore the protected access problem, the two unequal examples I just gave would compile to exactly the same bytecode. So how can the JVM report an error about an illegal access for the expression ((a.A) this).secret? Well, it turns out that it doesn't.

We can show this by first making secret public, compile B, then make secret protected, and only recompile A. This works like a charm: doing this trick for the following example prints secret = 5.

  package a;
  public class A {
    protected int secret;
    public void print() {
      System.out.println("secret = " + secret);
    }
  }

  package b;
  public class B extends a.A {
    void f() {
      ((a.A) this).secret = 5;
    }
    public static void main(String[] ps) {
      B b = new B();
      b.f();
      b.print();
    }
  }

However, if this would be allowed by bytecode in general, then this would mean that the security vulnerability that was fixed with the details of protected access, would actually only give source level protection. Obviously, that would be no protection at all: you can safely assume that attackers are capable of writing bytecode. So let's try to make the example a bit more adventurous by passing the subject expression to the f method:

  package b;
  public class B extends a.A {
    void f(a.A a) {
      a.secret = 5;
    }
    public static void main(String[] ps) {
      B b = new B();
      b.f(b);
      b.print();
    }
  }

This time, the verifier reports an error:

  Exception in thread "main" java.lang.VerifyError:
    (class: b/B, method: f signature: (La/A;)V)
    Bad access to protected data

This error report is correct, so apparently the verifier does check for illegal protected access. In the first case, it was just a bit more liberal than the source language. The question is, how is this specified in the Java Virtual Machine specif cation? My first impression was that there might be some special handling of accesses to this. However, this would require the verifier to trace which local variables might have the value of this, which is rather unlikely. Then, Dick Eimers (who did lots of scary bytecode stuff for his master thesis) pointed me to a paper that exactly covers this subject: Checking Access to Protected Members in the Java Virtual Machine by Alessandro Coglio. Strange enough, this paper is not cited anywhere, while I think that the discussion of this issue is pretty good.

It turns out that the difference in accessibility between super field accesses and ordinary field accesses is handled implicitly thanks to the type inferencer used by Java Virtual Machine. The inferred type of the operand of the field access will be more specific than the type in the corresponding source code, which makes the access to the protected field valid in bytecode. I don't think that this implicit handling of the observed difference is a very good idea.

Our take on injection attacks

If you haven't been hiding under some really impressive rock for the last decade, then you probably know that injection attacks are a major issue in web applications. The problem of SQL injection is well-known, but you see similar issues everywhere: SQL, Shell, XML, HTML, LDAP search filters, XPath, XQuery, and a whole series of enterprisey query languages, such as HQL, JDOQL, EJBQL, OQL are all potential candidates for injection attacks. Just search for any of these languages together with the term injection and observe the horror. Recently, it has also become more popular to mix a program written in Java with scripts, usually something like JavaScript, Ruby or Groovy. If you include user input in the script, then this is yet another vector of attack.

Solutions?

Of course it is possible to just advice programmers to properly escape all user inputs, which prevents most of the injection attacks. However, that's like telling people to do their own memory management or to do the dishes every day (which is a particular problem I have). In other words: you won't get it right.

Most of the research on injection attacks has focused on finding injection problems in existing source code using static and/or runtime analysis. Usually, this results in tools that check for injection attacks for specific languages (e.g. SQL) in specific host languages (e.g. PHP). This is very important and useful work, since it can easily be applied to detect or prevent injection attacks in existing code bases. However, at some point we just fundamentally need to reconsider the way we program. Why just fight the symptoms if you can fix the problem?

So that's what we've done in our latest work called StringBorg. I'm not going to claim that all your injection problems will be over tomorrow, but at least I think that what we propose here gives us some perspective on solving theses issues once and for all in a few years. The solution we propose is to use syntax embeddings of the guest languages (SQL, LDAP, Shell, XPath, JavaScript) in the host language (PHP, Java) and let the system do all the proper escaping and positive checking of user input.

Examples

The paper I'll mention later explains all the technical details, and I cannot redo that in a better way in a blog, so I'll just give a bunch of examples that illustrate how it works.

SQL

The first example is an embedding of SQL in Java. This example illustrates how you can insert strings in SQL queries and compose SQL queries at runtime. The first code fragment is the classic, vulnerable, way of composing SQL queries using string concatenation.

  String s = "'; DROP TABLE Users; --";
  String e = "username = \'" + s + "\'";
  String q = "SELECT password FROM Users WHERE " + e;
  System.out.println(q);

Clearly, if the string s was provided by the user, then this would result in an injection attack: the final query is SELECT password FROM Users WHERE username = ''; DROP TABLE Users; --'. Bad luck, the Users table is gone! (or maybe you can thank your database administrator).

With StringBorg, you can introduce some kind of literal syntax for SQL. The SQL code is written between the quotation symbols <|...|>. SQL code or strings can be inserted in another SQL query using the syntax ${...}. The example would be written in StringBorg as:

  String s = "'; DROP TABLE Users; --";
  SQL e = <| username = ${s} |>;
  SQL q = <| SELECT password FROM Users WHERE ${e} |>;
  System.out.println(q.toString());

This will result in the correct query, SELECT password FROM Users WHERE username = '''; DROP TABLE Users; --', where the single quotes have been escaped by StringBorg according to the rules of the SQL standard (the exact escaping rules depend on the SQL dialect). Not only does the StringBorg solution solve the injection problem, it is also much prettier! This example also shows that it is not required to know the full SQL query at compile-time, for example the actual condition e could be different for two branches of an if statement, or could even be constructed in a while statement.

The nice thing about StringBorg is that the SQL support is not restricted to a specific language, in this case Java. For PHP, you can do exactly the same thing:

  $s = "'; DROP TABLE Users; --";
  $e = <| username = ${$s} |>;
  $q = <| SELECT password FROM Users WHERE ${$e} |>;
  echo $q->toString(), "\n";

LDAP

Using user input in LDAP search filters has very similar injection problems. First a basic example, where there is no problem with the user input:

  String name = "Babs Jensen";
  LDAP q = (| (cn=$(name)) |);
  System.out.println(q.toString());

The resulting LDAP filter will be (cn=Babs Jensen), which is what you would except. If the string has the value Babs (Jensen), then the parentheses need to be escaped. Indeed, StringBorg will produce the filter (cn=Babs \28Jensen\29). This input might have been an accident, but of course we can easily change this into a real injection attempt by using the string *. Again, StringBorg will properly escape this, resulting in the query (cn=\2a).

Shell

Programs that invoke shell command could be vulnerable to injection attacks as well (as the TWiki developers and users have learned the hard way). Similar to the other examples, StringBorg introduces a syntax to construct shell commands, and escape strings:

  Shell cmd = <| /bin/echo svn cat http://x -r <| s |> |>;
  System.out.println(cmd.toString());

If s has the values bravo, foo bar, * and ; echo pwn3d! respectively, then the resulting commands are:

  /bin/echo svn cat http://x -r bravo
  /bin/echo svn cat http://x -r foo\ bar
  /bin/echo svn cat http://x -r \*
  /bin/echo svn cat http://x -r \;\ echo\ pwn3d\!

JavaScript

Not only does StringBorg prevent injection attacks, it also makes composing SQL, XQuery, JavaScript, etc more attractive: you don't have to concatenate all these nasty strings anymore. For example, the following example taken from and article on the new Java scripting support is just plain ugly:

  jsEngine.eval(
    "function printNames1(namesList) {" +
    "  var x;" +
    "  var names = namesList.toArray();" +
    "  for(x in names) {" +
    "    println(names[x]);" +
    "  }" +
    "}" +

    "function addName(namesList, name) {" +
    "  namesList.add(name);" +
    "}"
  );

whereas this looks quite reasonable:

  jsEngine.eval(|[
    function printNames1(namesList) {
      var x;
      var names = namesList.toArray();
      for(x in names) {
        println(names[x]);
      }
    }

    function addName(namesList, name) {
      namesList.add(name);
    }
  ]| );

Of course, this would be easy to fix by introducing multi-line string literals in Java, but in addition to the nicer syntax, you get protection against injection attacks and compile-time syntactic checking of the code for free!

Generic, generic, generic

Now, if you are familiar with our work, then this solution won't really surprise you, since we have been working on syntax embeddings for some time now (although in different application areas, such as meta programming). However, this work is quite a fundamental step towards making these syntax embeddings easier to use by ordinary programmers. First, the system now supports ambiguities, which always was the weak point of our code generation work: if you don't support ambiguities, then the programmer needs to be familiar with the details of the grammar of the guest language, which you really don't want. Fortunately, this is now a technical detail that you now can forget about! Second, no meta-programming is required at all to add a new guest language (e.g. XPath) to the system. All you need to do is define the syntax of the language, define the syntax of the embedding, and optionally define escaping rules for strings and you're all set. Thus, compared to our previous work on MetaBorg (OOPSLA '04), there is no need for implementing the mapping from the syntax of the guest language to code in the host language.

This is a pretty amazing property: basically, this means that you can just use languages as libraries. You can just pick the languages you want to use in a source file and that's it! No difficult meta-programming stuff, no program transformation, no rewrite rules and strategies, no limitations. In fact, this even goes beyond libraries: libraries are always language specific (for example for Java or PHP), but the implementation of support for a guest language (e.g. SQL) is language independent. This means that if some person or company implements support for a guest language (e.g. SQL) then all host languages (Java, PHP, etc) are immediately supported.

Future?

The paper we wrote about this is titled "Preventing Injection Attacks with Syntax Embeddings. A Host and Guest Language Independent Approach" and is now available as technical report. Last week, we submitted this paper to the USENIX Security Symposium. We won't know if the paper is accepted until April 4, but I would be flabbergasted if it got rejected ;) . Our prototype implementation, called StringBorg, is available as well. I'm looking forward to your feedback and opinions. I'll add some examples to the webpage later this week, so make sure to come back!

Peter Ahé already has a general solution for embedding foreign languages on his wish list (as opposed to an XML specific solution), so could this actually be happening in the near future?

Some random thoughts on the complexity of syntax

For some reason I was invited to join the JSR-308 mailing list, which is a public mailing list discussing JSR-308. I've reported some problems with the Java grammar of the Java Language Specification in the past, so maybe I'm now on some list of people that might be interested. I'm not sure if I will contribute to the discussion, but at least lurking has been quite interesting. If you're not familiar with the proposal, JSR-308 has been started to allow annotations at more places in a Java program. The title of the JSR is "Annotations on Java Types", but the current discussion seems to interpret the goal of the JSR a bit more ambitiously, since there is a lot of talk going on about annotations of statements, and even expressions. I don't have a particularly strong opinion on this, but it's interesting to observe how the members of the list are approaching this change in the language.

Neal Gafter seems to represent the realistic camp in the discussion (not to call his opinion conservative). Neal was once the main person responsible for the Java Compiler at Sun, so you can safely assume that he knows what he's talking about. Together with Joshua Bloch, he is now mainly responsible for the position of Google in these Java matters. Last week, he sent another interesting message to the list: Can we agree on our goals?. As I mentioned, I don't have a very strong opinion on what the goal of the JSR should be, but Neal raised a point about syntax that reminded me again of some thoughts on syntax that have been lingering in my mind for some time now. Neal wrote:

"I think the right way to design a language with a general annotation facility is to support (or at least consider supporting) a way of annotating every semantically meaningful nonterminal. Doing that requires a language design with a very simple syntax. Java isn't syntactically simple, and I don't think there is anything we can do it make it so. If we wanted to take this approach with Java we'd have to come up with a syntactic solution for every construct that we want to be annotatable. Given the shape of the Java grammar, that solution would probably be a different special case for every thing we might want to annotate."

Whether you like it or not, this is a most valid concern. The interesting point about this annotation thing is that it is a language feature that applies in a completely different way to existing language constructs. Adding an expression, a statement, or some modifier to the language is not difficult to do, since this adds only an alternative to the existing structure of the language. Annotations, on the other hand, do not add just another alternative, but crosscut (sorry, I couldn't avoid the term) the language. If you are an annotation guy, then you want to have them everywhere, since you essentially want to add information to arbitrary language constructs. Now this is quite a problem if you have a lot of language constructs, not alternative language constructs, but distinct kinds of language constructs (of course known as nonterminals to grammar people). This would be trivial to do in language where there are not that many language constructs, such as Lisp and Scheme, and even model-based languages.

This makes you wonder what is a good language syntax. Should adding such a crosscutting language feature be easy? Conceptually, it is beyond any doubt attractive to have a limited number of language constructs, but on the other hand it is very convenient that Java has this natural syntax for things like modifiers, return types, formal parameters, formal type parameters, throws clauses, array initializers, multiple variable declarations at the same line, and so on. If you want to add annotations to all these different language constructs, then you basically have to break their abstraction, which suddenly makes them look unnatural, since it becomes clear that a syntactical construct that used to be easy to read, has some explicit semantic meaning. That is, the entire structure of the language is exposed in this way. It is no longer possible to read a program, abstracting over all the details of the language. Also, for several locations it is very unclear to the reader where an annotation refers to. For example, the current draft specification states that

"There is no need for new syntax for annotations on return types, because Java already permits an annotation to appear before a method return type. Currently, such annotations are interpreted as on the method declaration — for example, the @Deprecated annotation indicates that the method is deprecated. The person who defines the annotation decides whether an annotation that appears before the return value applies to the method declaration or to the return type.

Clearly, there is a problem in this case, since an annotation in the header of the method could refer to several things. The reason for this, is the syntactical conciseness of the language for method declarations: you don't have to identify every part explicitly, hence if you want to annotate some part only, then you have a problem. Moving that decision to the declaration side of the annotation is a not an attractive solution, for example there will be annotations that are applicable to both declarations and types.

This all brings us to the question how to determine if a syntax of a programming language is simple? Is that really just some subjective idea, or is it possible to determine this semi-automatically with more objective methods? I assume that the answer depends on the way the language is applied. For example, in program transformation it is rather inconvenient to have all kinds of optional clauses for a language construct. This reminds me of a post by Nick Sieger, who applied a visualization tool to some grammars. For some reason, this post was very popular and was discussed all over the web, including Lambda the Ultimate and LWN. However, most people seemed to agree that the visualizations did not tell much about the complexity of the languages. Indeed, the most visible aspects of the pictures are the encodings of the actual grammar that had to be applied to make the grammar non-ambiguous or to fit in the used grammar formalism. For example, the encoding of precedence rules for expressions makes the graph look pretty, but conceptually this is just a single expression. As a first guess, I would expect that some balance between the number of nodes and edges would be a better measurement: lots of edges to a single node, means that nonterminal is allowed at a lot of places, which is probably good for the orthogonality of the language (more people have been claiming this in the discussion about these visualizations).

But well, this makes you wonder if there has been any research on this. The only work I'm familiar with is SdfMetz, which is a metrics tool for SDF developed by Joost Visser and Tiago Alves. SDF grammars are usually closer to the intended design of a language than LR or LL grammars, so if you are interested in the complexity of the syntax of a language, then using SDF grammars sounds like a good idea. SdfMetz supports quite an interesting list of metrics. Some are rather obvious (count productions etc), but there are also some more complex metrics. I'm quite sure that (combinations of) these metrics can give some indication of the complexity of a language. Unfortunately, the work on SdfMetz was not mentioned at all in the discussion of these visualizations. Why is it that a quick and dirty blogpost is discussed all over the web and solid research does not get mentioned? Clearly, the SdfMetz researchers should just post a few fancy pictures for achieving instant fame ;) . Back to the question what is a good syntax, they have mostly focused on the facts until now (see their paper Metrication of SDF Grammars), and have not done much work on interpreting the metrics they have collected. It would be interesting if somebody would start doing this.

Joost Visser and Tiago Alves will be presenting SdfMetz at LDTA 2007, the Workshop on Language Descriptions, Tools and Application (program now available!). As I mentioned in my previous post, we will be presenting our work on precedence rule recovery and compatibility checking there as well. So, if you are in the neighbourhood (or maybe even visiting ETAPS), then make sure to drop by if you are interested!

Another thing that finally seems to get some well-deserved attention is ambiguity analysis. It strikes me that the people on the JSR-308 list approach this rather informally, by just guessing what might be ambiguous or not. It should be much easier to play with a language and determine how to introduce a new feature in a non-ambiguous way. Sylvian Schmitz will be presenting a Bison-based ambiguity detection tool at LDTA, so that should be interesting to learn about. The paper is already online, but I haven't read it yet. Maybe I'll report about it later.

Grammar Engineering, I'm loving it

I think that the most attractive thing to work on as a researcher are the problems that you actually encounter yourself. Obviously, it is an option to try to encounter these problems by studying or writing code that you are not actually directly interested in, but it is much more fun if you can work on issues in code that you just write out of your own interest.

This is what we've done in our latest paper, titled "Grammar Engineering Support for Precedence Rule Recovery and Compatibility Checking". Most of our work involves syntax definitions, for example to provide general support for the implementation of program transformations and also for our research on embedding and composing languages (for various applications). One of the problems we encounter is that the conversion of a grammar from one grammar formalism to another is rather unreliable. For example, if you need to convert a grammar from YACC to SDF, then you basically have no idea if the two grammars are compatible. For programs in general, this is understandable, since imperative source code are very difficult to compare. But, if you have a more or less declarative specification of a grammar, how is it possible that you cannot compare them at all?

As a first step towards supporting grammar compatibility checking, we have implemented a tool that compares the precedence rules of grammars. A very simple example of a precedence rule is that 1 + 2 * 3 should be parsed as 1 + (2 * 3). The precedence rules of a grammar might look like a trivial property at first sight, but actually it is rather complex to understand as a human what the precedence rules of a YACC or SDF grammar are. This tool has already been most successful for comparing existing C grammars written in YACC and SDF and deriving the exact precedence rules of PHP, which has quite a bizarre expression language.

The paper has been accepted for LDTA 2007, the Workshop on Language Descriptions, Tools and Applications, which is an excellent place for this subject. We will present our work at this workshop at the end of March. A draft version of the paper is available from the publication list at my homepage. The implementation is available as part of the Stratego/XT Grammar Engineering Tools. The website includes a bunch of examples. In the future, we hope to provide more tools to assist with the maintenance, testing, conversion, and analysis of grammars. In fact, Stratego/XT itself already contains some interesting tools for this, most prominently the grammar unit-testing tool parse-unit.

Now I think of it, it's probably a bad idea as a vegetarian to paraphrase a campaign by McDonald's ...

Java-front versus Jackpot, APT, Eclipse JDT

Recently, there have been a few interesting developments in standard support for open compilers and program transformation. For example, Sun released the annotation processing tools (APT) as part of the JDK5, which opens up Sun's Java compiler a bit. Also, there is Jackpot, a plugin for Netbeans for transformating Java code. The obvious question is how this relates to work that has been done in research on open compilers and program transformation.

Olivier Lefevre sent me an email to ask how the tree for Java provided by Jackpot and javac compares to our support for parsing and transforming Java in Java-front. The answer is probably useful in general, so I'll quote it here. Feel free to share your opinion in the comments!

As you may know, starting with Java 6 the Sun JDK will ship with an API to the AST: see jackpot api

Yes, Jackpot and APT are great projects. However, there is not yet a full API to the AST in the standard JDK, afaik. The compiler will be more 'open' in two different ways.

First, the current annotation processing tool (APT) is going to be combined with javac, but APT only provides access to the global structure of a Java source file and does not include the statement and expression level. Also, this API does not allow modification of the Java representation. APT is read-only: you can only generate new code.

Second, there is Jackpot, which is a rule-based language for transforming Java code. For Jackpot, the representation of Java used by javac has been opened and cleaned up a bit to make it more usable in external tools. However, this representation is not standardized and Sun recommends not to use stuff from com.sun.*. Afaik, Jackpot will be shipped as part of NetBeans and not as part of the JDK.

How does this compare to Java-front?

That's a good question. The answer depends on the application.

If you just need an AST for Java, then the advantage of the com.sun.source.tree AST is that you are absolutely sure that the AST conforms to javac, since the implementation is exactly the same. Of course, the same holds for ecj and the AST of Java that is provided by Eclipse (org.eclipse.jdt.core.dom.*). However, the grammar provides by Java-front is very good, so I don't expect any parsing problems. It has been tested and used heavily in the last few years and the development of this grammar has even resulted in a number of fixes in the JLS.

An advantage of Java-front is that it is a bit more language independent. Obviously, the Eclipse and Javac ASTs are to be used in Java. If you want to implement a transformation of Java in a different language, then you have to write an exporter. Java-front outputs ASTs in a language independent exchange format (ATerms), which can also be converted to XML. Of course, Java-front is most useful if you combine it with a language that is designed for program transformation and operates on ATerms, such as Stratego. One of the biggest advantages of Stratego is that it is very easy to do traversals over the AST: no tiresome visitors.

If you need more information about Java than can be defined in a context-free grammar, then you need more than just a parser. For more complex transformations (which includes simple refactorings), you'll probably need an implementation of disambiguation (reclassification) and qualification of names. A simple statement like System.out.println is already highly ambiguous with an analysis: is System a variable? a class? a package? Is out an inner class? a field? Most likely, you'll need type information as well. Java and Eclipse have the major advantage that you can safely assume that their type checkers are pretty good. For Jackpot, I suppose that there is some way to get type information (since type information can be used in Jackpot), but I from a quick scan I cannot figure out how to do this from the public API. For Java-front, there is an extension (Dryad) that supports type-checking and disambiguation, but this work is not yet complete. Using an existing compiler is of course a safer alternative. For experiments, the stuff provided by Dryad should be ok (we use it in our course on program transformation).

A different application is the implementation of Java language extensions. Javac and ECJ do not support this. The Java representation is open, but not extensible. Java-front uses a modular syntax definition formalism (SDF) that allows you to extend the grammar of Java in an almost trivial way. The strength of this approach is illustrated by the embedding of the Java syntax in Stratego (GPCE '02) and Java (GPCE '05), the applications of the grammar in MetaBorg (OOPSLA '04), and the modular extension of the grammar for the definition of the AspectJ syntax (OOPSLA '06). Of course, these applications are not really interesting if you are just interested in a Java program transformation tool, but it illustrates the reusability of such a syntax definition (as opposed to the grammars used by ecj, javac and most other parser generators). You'll need tools for pretty-printing as well. Outside of Eclipse, pretty-printing the JDT Core DOM is troublesome and mostly useful for debugging the output only. Inside Eclipse, the support for pretty-printing and preserving the layout of a program is of course excellent (see the existing implementations of refactoring). Jackpot provides a pretty-printer as well, but I don't know if it can be used outside NetBeans. Java-front provides the tool pp-java, which has been heavily tested and can insert parentheses in exactly the right places.

I am interesting in implementing small refactorings.

If you want to implement solid refactorings that could eventually even be deployed, then I would suggest to use an existing framework for refactoring, since there is much more to do than just getting an AST. A few years ago, I implemented an extract method refactoring in JRefactory, which was quite a useful experience. I suppose it's a bit obsolete now, since the refactoring market is dominated by refactorings directly supported by IDEs. You could consider Eclipse or NetBeans.

If your objective is to play a bit with program transformations and maybe even be a bit more adventurous by using real program transformation languages, then it might be nice to use Java-front and Stratego. Using Stratego is a major advantage of the tiresome implementation of traversals in Java (and most other languages).

Hope this helps :) Feel free to ask more questions if anything is unclear :)

On the Details of Protected Access in Java

When I was working on my master thesis I spent most of my time in the Software Technology Lab of our department. This was really a gorgeous place to work, mostly due to a group of great fellow students. After finishing our master study, we all left this lab a few years ago and have a job now. One of the students I worked with first went to a research institute to work on maritime simulations and is now moving to a company that produces Real Software.

Of course, the problem is that universities do a terrible job at learning students how to write Real Software, so the first thing these companies do is give their employees some proper education. Real Software is written in Java, so the first step is to become a Sun Certified Programmer for the Java 2 Platform, aka SCJP.

This guy is now going through the SCJP materials, which covers Java at a surprising level of detail. Universities do at least one thing right: they stimulate you to think about the things you learn, instead of just accepting it all as-is. As a result, we have had some nice discussions about the details of Java, and he has only just started ;) .

Last week, we started talking about the details of protected access in Java. The study material seems to mention the detailed rules, but does not explain why the rules are the way they are. Strange enough, I could not find any good resource that does this, and I couldn't remember where I learned about this, so I decided to explain it myself.

Protected Access

The Java Language Specification has two subsections that define the rules for accessibility. The general rules are in subsection 6.6.1: Determining Accessibility and some of the more complicated rules for protected access are in subsection 6.6.2: Details on Protected Access. The rules of the first subsection are rather straightforward and don't need much explanation.

For protected members and constructors this subsection defines that the members are accessible if the access occurs from within the same package. This rule is clear, although many people seem to be surprised by this. However, the second case refers to subsection 6.6.2, where most of the confusion is about. This subsection defines the additional accessibility rules for protected access, to which most people are informally familiar as “a protected member is accessible from subclasses”. A simple example where this accessibility rule is applied:

package a;

public class A {
  protected int secret;
}

package b;

public class B extends a.A {
  void f() {
    secret = 5;
  }
}

However, the rules that define accessibility of protected members from subclasses are a bit more complex than you might expect. The problem is that the rule you know informally, is not that clear anymore if the access of the protected member is qualified (i.e. is applied to an object, not implicitly to this). Consider this simple example:

package a;

public class A {
  protected int secret;
}

package b;

public class B2 extends a.A {
  void f(a.A a) {
    a.secret = 5;
  }
}

In this example, the access to the protected instance field secret of A occurs from a subclass B of A, so according to our informal idea of protected access, this should be allowed. However, this example should make you feel a bit uncomfortable. Indeed, this is not allowed in Java. Let's take a look at what the compilers say:

$ javac b/B2.java
b/B2.java:5: secret has protected access in a.A
    a.secret = 5;

$ jikes b/B2.java
Found 1 semantic error compiling "b/B2.java":

     5.     a.secret = 5;
              ^----^

*** Semantic Error: The instance field "secret" in class "A" has protected access, but the qualifying expression is not of type "B2" or any of its enclosing types.

$ ecj b/B2.java
----------
1. ERROR in b/B2.java
 (at line 5)
        a.secret = 5;
        ^^^^^^^^
The field A.secret is not visible
----------

Side note: I was a bit surprised by the error report of ecj. This could be a bug: the protected field is visible but it is not accessible. The error report of jikes is by far the best.

Basically, if this access would be allowed, then you can access any protected field of any class, by just making a subclass of the class that declares the protected field. Hence, you could never really protect your protected fields if this kind of access would be allowed. Consider this example:

package a;

public class A {
  protected int secret;
}

package b;

public final class MySecurityHazard extends a.A {
}

package c;

public class C extends a.A {
  void f(b.MySecurityHazard b) {
    b.secret = 5;
  }
}

In this example the class MySecurityHazard has deliberately been declared to be final to avoid that the sensitive fields of this class can be accessed. However, according to our (now deprecated) informal knowledge of protected access, we can just create another subclass C of A that can be used to access the protected fields of MySecurityHazard.

How can we define the protected access that we would like to have? Of course, qualified access to protected members could be forbidden completely (and maybe that would have been a good idea), but Java is a bit more flexible, without introducing security problems. The basic problem of the unwanted access is that you start a new inheritance branch and access protected fields from there. So, the qualified access to protected fields should be restricted to the same inheritance branch as the object to which it is applied. This is exactly what the details of protected access are about. They define that access from a class S is permitted only if the type of the qualifier is S or a subclass of S. Let me now finally quote the specification:

Let C be the class in which a protected member m is declared. Access is permitted only within the body of a subclass S of C. In addition, if Id denotes an instance field or instance method, then:

• If the access is by a qualified name Q.Id, where Q is an ExpressionName, then the access is permitted if and only if the type of the expression Q is S or a subclass of S.
• If the access is by a field access expression E.Id, where E is a Primary expression, or by a method invocation expression E.Id(. . .), where E is a Primary expression, then the access is permitted if and only if the type of E is S or a subclass of S.

See JLS3, Section 6.6.2.1: Access to a protected Member

Note that these more specific rules only apply to instance members, not to static ones. If you are using a protected static field or method, make sure that you understand what you're doing: anyone will be able to access the field or method by just creating another subclass of the class that declares the protected field or method. It would not make any sense to impose these additional rules on static members, since all subclasses actually share the static member! So, while protected access on instance members could be used for security reasons, protected access on static members is only useful for hiding the member by restricting its accessibility to a part of a program where it is relevant.

If you read the rules carefully, then they are not too unclear. The real source of confusion seems to be that there is no motivation why these more complex rules are necessary. I hope this blog has solved that.

Lifting Member Classes from Generic Classes

I've been working of Java generics and member classes during the last few weeks. In particular, I had to find out how the additional information on Java generics is exactly represented in bytecode attributes of generic classes and methods (aka generic signatures). I was surprised by the way member classes of generic classes are compiled and I'm worried about the consequences of this for future updates of the JVM specification. That's what this entry is about.

First, something about the relation between member classes and lambda lifting. The Java language supports member classes, but Java bytecode does not. Therefore, Java compilers have to lift member classes to top-level classes, a transformation that is comparable to lambda lifting (see for example the paper Lambda-Lifting in Quadratic Time).

Member classes are compiled to ordinary top-level classes where the constructor takes an extra argument for the instance of the enclosing class. For example, the constructor a member class Bar of class Foo will get an additional argument of type Foo for the enclosing instance of a Bar object. Constructors of local classes (classes declared in a method) also take additional arguments for the local variables that it uses from its enclosing method. This process of lifting classes (that have lexical scope) is very similar to lifting nested functions in lambda lifting: all local variables that are used in the nested class become explicit arguments to make the nested class scope insensitive. After that, the class can simply be lifted out of its original scope to the top-level. An essential property of the class (or function) after lifting is that the nested class (or function) no longer directly refers to variables of the original scope of the nested class or function.

In Java 5.0, parameterized types and methods (aka generics) have been introduced. In combination with member classes, this raises the question how type variables should be handled when lifting member classes. From the source code point of view, this is pretty obvious:

class Foo<A> {
  class Bar {
    A get() { ... }
  }
}

If the class Bar is lifted, then its constructor gets an additional parameter for the enclosing Foo instance. This Foo instance is parameterized using a type variable A, so the lifted class Bar should also be parameterized with a type: the type parameter of its enclosing instance. This lifting of type parameters is comparable to the lifting of parameters for normal variables. So, the result of source-level lifting the Bar class could be:

class Foo<A> {
}

class Bar<A> {
  private final Foo<A> _enclosing;

  public Bar(Foo<A> enclosing) {
    _enclosing = enclosing;
  }

  A get() { ... }
}

Indeed, the Eclipse implementation of the refactoring "Move Member Type to New File" adds the type parameter to the lifted class (thumbs up for the generics support in Eclipse!).

So, what happens in Java bytecode? Should the lifted class have a type parameter? Should the lifted class be a valid class, generically speaking? (of course, a JVM is currently not required to understand generics related information in bytecode).

Well, the lifted class does not have type parameter and, generically speaking, it is not a valid class. Let's take a look at the bytecode, represented in a structured way as an aterm, produced by a tool called class2aterm (I use ... to leave out some details and // to explain what the code means. The full aterm is available here)

$ class2aterm -i Foo\$Bar.class --parse-sig | pp-aterm
ClassFile(
  ...
  // field for the enclosing Foo instance.
  Field(
    AccessFlags([Final, Synthetic])
  , Name("this$0")
  , FieldDescriptor(ObjectType("Foo"))
  , Attributes([])
  )
  ...

  // constructor, taking a Foo argument.
  Method(
    AccessFlags([])
  , Name("<init>")
  , MethodDescriptor([ObjectType("Foo")], Void)
  , Attributes([])
  )
  ...

  // get method with a generic signature
  Method(
    AccessFlags([])
  , Name("get")
  , MethodDescriptor([], ObjectType("java.lang.Object"))
  , Attributes(
      [ MethodSignature(
          TypeParams([])
        , Params([])
        , Returns(TypeVar(Id("A")))
        , Throws([])
        )
      ]
    )
  )
  ...

  // attributes of the class Bar
  Attributes(
    [ SourceFile("Foo.java")
    , InnerClasses( ... )
    ]
  )
  ...
)

This disassembled class file reveals some interesting details about the way nested classes are lifted:

• The lifted class Bar is not parameterized: it has no ClassSignature attributed, which should be there if the class takes formal type parameters.
• The field for the enclosing class does not have a parameterized type. Its type is the raw type Foo!
• The constructor of Bar (the method name <init>) has no generic signature and takes a raw type Foo as an argument.
• The get method does have a generic signature, which describes that the method returns a type variable A.

Of course, all the information of the original source can be reconstructed by a tool that knows about member classes and generics. But, to a tool that only knows about generics, this code would be considered incorrect. Hence, if the virtual machine would support generics in the future (which is an option explicitly left open), then this code would be incorrect! The type variable mentioned in the generic signature of the get method is not in scope. Hence, the JVM would be required to have knowledge of inner classes as well as generics to be able to find out what type parameter this type variable refers to. Unless, of course, the bytecode format is changed, which will still make it impossible to run code compiled to the current bytecode format under the new JVM, which has always been a important requirement for Sun when working on extensions of the Java platform (language and virtual machine).

Furthermore, the type variable in the signature of the get method is not qualified. Every single name in Java bytecode is fully qualified, which is very useful for tools that need to work on bytecode: they don't have to name analysis to find out to what construct a name refers. Type variables are not qualified, which complicates the analysis that has to be performed by a tool that operates on bytecode. Not only can this type variable refer to type parameters of arbitrary enclosing classes, it could also refer to type parameters of enclosing generic methods (for local classes or member classes in local classes).

The fact that type variables in bytecode are not qualified is already quite annoying without considering member classes. In the Java language, it is allowed to redeclare type variables. For example:

class Foo<A> {
  <A> void foo(A x) {
  }
}

In this example the type parameter A of the foo method is a different type parameter then the A parameter of the class Foo. This basically means that a bytecode processing tool with knowledge of generics has to do name analysis, which is definitely not something that is desirable for a bytecode format. Introducing canonical, fully qualified names for type variables would solve this.

As you might know, I'm working on semantic analysis for Java in the context of the Stratego/XT project. My goal is to make it possible to define program transformations in Stratego at the semantic level: in program transformations consider the actual meaning of names, types of expressions, and so on, without requiring the programmer to redo the semantic analysis, which is quite complex for a 'real' language like Java. Obviously, I have decided to qualify type variables. For example, the parameter A of the method foo in class Foo in the last example is represented as:

Param(
  []
, TypeVar(
    MethodName(TypeName(PackageName([]), Id("Foo")), Id("foo"))
  , Id("A")
  )
, Id("x")
)

The MethodName is the qualifier of the type variable in this example. This qualifier makes it immediately clear that the type variable refers to the type parameter of the method foo.

I don't know if this would have been fixed (maybe I see this completely wrong), but still it's a pity that I wasn't able to give feedback on this before JSR14 was finished. At that time, I was still working on the syntactic part of my Java transformation project (which is now available as Java Front). I gave some feedback on the syntax of generics, annotations and enumerations (mostly typos and minor bugs), but that's about it. For reducing the number of possible problems, I think that it would be very useful if new language features, such as generics, are also implemented with alternative techniques and tools. For example, I was able to give some feedback on the syntax of Java, because I was implementing a parser by creating a declarative syntax definition in SDF, a modular syntax definition formalism that integrates lexical and context-free syntax. These unconventional approaches might in general result in valuable feedback on proposals for new language features.