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Solving the Expression Problem in C++, á la LMS



We give a C++ solution to the Expression Problem that takes a components-for-cases approach. Our solution is a C++ transliteration of how Lightweight Modular Staging solves the Expression Problem. It, furthermore, gives a C++ encoding to object algebras and object algebra interfaces. We use our latter encoding by tying its recursive knot as in Datatypes à la Carte.
Solving the Expression Problem in C++, `a la
(Author’s Version) ?
Seyed H. Haeri (Hossein)1and Paul Keir2
1Universit´e catholique de Louvain, Louvain-la-Neuve, Belgium
2University of the West of Scotland, UK
Abstract. We give a C++ solution to the Expression Problem that
takes a components-for-cases approach. Our solution is a C++ translit-
eration of how Lightweight Modular Staging solves the Expression Prob-
lem. It, furthermore, gives a C++ encoding to object algebras and object
algebra interfaces. We use our latter encoding by tying its recursive knot
as in Datatypes `a la Carte.
1 Introduction
The Expression Problem (EP) [6,31,37] is a recurrent problem in Programming
Languages (PLs), for which a wide range of solutions have been proposed. Con-
sider those of Torgersen [35], Odersky and Zenger [20], Swierstra [34], Oliveira
and Cook [23], Bahr and Hvitved [2], Wang and Oliveira [38], Haeri and Schupp
[16], and Haeri and Keir [12], to name a few. EP is recurrent because it is repeat-
edly faced over embedding DSLs – a task commonly taken in the PL community.
Embedding a DSL is often practised in phases, each having its own Algebraic
Datatype (ADT) and functions defined on it. For example, take the base and
extension to be the type checking and the type erasure phases, respectively. One
wants to avoid recompiling, manipulating, and duplicating one’s type checker if
type erasure adds more ADT cases or defines new functions on them.
Haeri [11] phrases EP as the challenge of implementing an ADT – defined by
its cases and the functions on it – that:
E1. is extensible in both dimensions: Both new cases and functions can be added.
E2. provides weak static type safety: Applying a function fon a statically3con-
structed ADT term tshould fail to compile when fdoes not cover all the
cases in t.
?This work is partially funded by the LightKone European H2020 Pro ject under Grant
Agreement No. 732505 and partially by the Belgian National Fund for Scientific
Research (F.R.S.-FNRS).
3If the guarantee was for dynamically constructed terms too, we would have called it
strong static type safety.
E3. upon extension, forces no manipulation or duplication to the existing code.
E4. accommodates the extension with separate compilation: Compiling the ex-
tension imposes no requirement for repeating compilation or type checking
of existing ADTs and functions on them. Compilation and type checking of
the extension should not be deferred to the link or run time.
On the other hand, Rompf and Odersky [32] coin Lightweight Modular Stag-
ing (LMS) for Polymorphic Embedding [17] of DSLs in Scala. They employ a
fruitful combination of the Scala features detailed in [21] that, as a side-product,
offers a very simple yet effective solution to the EP. We call that side-product
the “Scala LMS-EPS.” In this paper, we offer a new C++ solution that is greatly
inspired by the Scala LMS-EPS. We call our own solution the “C++ LMS-EPS.”
Amongst the EP solutions, LMS is distinctive for its ease of extension: both in
adding new ADT cases and functions defined on them. We chose to implement
LMS in C++ to show the independence of LMS from Scala’s combination of
following features: traits, abstract type members, and super calls. Instead, the
C++ LMS-EPS makes use of the following C++ features: curiously recurring
template pattern (§3.2), abbreviated function templates of C++204(§3.3),
user-defined deduction guides and variadic templates (§5.3), and, most notably,
std::variant (§3.1). (For the unfamiliar reader, an introduction to those C++
features comes in Appendix A.) Unlike Scala, C++ is a mainstream language
which is well-known for its efficiency. Similar to Scala, C++ is a multi-paradigm
language with a high level of abstraction (from C++17 onward).
Given its presentation in C++, the C++ LMS-EPS machinery may look as
an EP solution that is too specific to C++. In order to correct that impression,
we recall that it is typical for EP solutions to be presented with tactful uses of
a single language. Take Datatypes `a la Carte [34], CDTs [1], PCDTs [2], and
MRM [25] in Haskell, Polymorphic Variants [9] in OCaml, and LMS [32] and
MVCs [22] in Scala. The C++ LMS-EPS is amongst the few EP solutions which
are presented in a mainstream programming language.
Here is a list of our contributions:
The C++ LMS-EPS takes a components-for-cases (C4C) [11] approach (§2.1
and 3.1). It implements ADTs (§3.2) using an encoding of object algebra inter-
faces [26] that is akin to Swierstra’s sum of functors [34]. We tie the recursive
knot using F-Bounding [4]. To implement functions on ADTs (§3.3), the C++
LMS-EPS gets a simulation of Haskell’s Combinator Pattern [28,§16] (§5.3) to
first acquire an encoding of object algebras. Our latter encoding, however, does
not use self -references [27]. The C++ LMS-EPS outperforms its Scala predeces-
sor by ensuring strong static type safety (§4.2). The way to distinguish between
the C++ LMS-EPS and EP solutions that use Generalised Algebraic Datatypes
(GADTs) is in §6. Detailed discussion on the related work comes in §7.
4Although our codebase remains fully functional without that (using ordinary type
parametrisation), we retain its usage here for enhanced readability.
2 Background
2.1 Formal Notation
In this paper, we use parts of the γΦC0calculus developed for solving the
Expression Compatibility Problem [15]. γΦC0was developed after observing
that sharing ADT cases amongst ADTs is not limited to ADTs only extend-
ing one another. For example, consider the ADTs α1,α2, and α3defined as:
α1::= Num(Z)|Add(α1), α2::= Num(Z)|Add(α2)|Mul (α2), and α3::=
Num(Z)|Add(α3)|Sub(α3). Both α2and α3extend α1. But, neither of them is
an extension to the other. In order to share the implementation effort required
for encoding α2and α3, then γΦC0promotes the ADT cases to components (in
their Component-Based Software Engineering [33,§17],[29,§10] sense).
In γΦC0, ADT cases are independent of ADTs but still parameterised by
them. In the γΦC0notation, one would write α1=Num Add,α2=Num
Add Mul, and α3=Num Add Sub. In the γ ΦC0ADT definitions, what
comes to the r.h.s. of the “=” is called the case list of the ADT on the l.h.s. of
the “=”. The connection between γΦC0and C4C becomes more clear in §3.1.
Hereafter, we refer to α1as NA (for Numbers and Addition) and to α2as NAM
(for Numbers, Addition, and Multiplication).
2.2 The Scala LMS-EPS
Suppose one is interested in encoding NA and in evaluating its expressions. One
possible Scala implementation is:
1trait NA {
2trait Exp //Exp ::=
3case class Num(n: Int) extends Exp // Num(n) |
4case class Add(l: Exp, r: Exp) extends Exp // Add(Exp, Exp)
5def eval: Exp => Int = {
6case Num(n) => n
7case Add(l, r) => eval(l) + eval(r) } }
Scala uses inheritance for definition of ADT cases. In lines 3and 4above,
for example, Num and Add inherit from their ADT type, i.e., Exp. Implementing
NAM without manipulation or duplication of NA can now be done as:
1trait NAM extends NA { //Exp ::= ... |
2case class Mul(l: Exp, r: Exp) extends Exp // Mul(Exp, Exp)
3override def eval: Exp => Int = {
4case Mul(l, r) => eval(l) *eval(r)
5case e => super.eval(e)} }
Line 2above adds the new case (Mul). Line 4above handles its evaluation.
And, line 5above makes a super call to employ the evaluation already defined
at NA. Note that NAM inherits Num and Add because it extends NA.
Addition of a function on NA whilst addressing E3 and E4 is similar. For
example, here is how to provide pretty printing:
1trait NAPr extends NA {
2def to_str: Exp => String = {
3case Num(n) => n.toString
4case Add(l, r) => to_str(l) + "+"+ to_str(r)} }
3 The C++ Version
C++ offers no built-in support for ADTs. Neither does it support mixin-composi-
tion for a super call to be possible. The C++ LMS-EPS mitigates those by
exercising a coding discipline that is explained in §3.1 to §3.3. Term creation
and application of functions on that comes in §3.4.
3.1 Cases
An EP solution takes a C4C approach when each ADT case is implemented
using a standalone component that is ADT-parameterised. In the C++ LMS-
EPS, the ADT-parametrisation translates into type-parametrisation by ADT.
For example, here are the C++ counterparts of Num and Add in §2.2:
1template<typename ADT> struct Num {//Num α:Zα
2Num(int n): n_(n) {}
3int n_;
Above comes a C4C equivalent of Num in §2.2. Verbosity aside, an important
difference to notice is that Num in §2.2 is a case for the ADT Exp of §2.2,
exclusively. On the contrary, the above Num is a case for the encoding of every
ADT αsuch that Num cases (α). The Add below is similar.
1template<typename ADT> struct Add {//Add α:α×αα
2using CsVar = typename ADT::cases;
3Add(const CsVar& l, const CsVar& r):
4l_(std::make_shared<CsVar>(l)), r_(std::make_shared<CsVar>(r)) {}
5const std::shared_ptr<CsVar> l_, r_;
Terms created using Add, however, are recursive w.r.t. their ADT. That is
reflected in line 5with the l_ and r_ data members of Add being shared pointers
to the case list of ADT, albeit packed in a std::variant. (See line 2in NATemp
below.) Line 2is a type alias that will become more clear in §3.2. The need for
storing l_ and r_ in std::shared_ptrs is discussed in §5.1.
We follow the terminology of C++ IDPaM5[12] in calling Num and Add of
this section and similar C4C encodings of ADT cases the case components.
3.2 ADTs
Defining ADTs in the C++ LMS-EPS is less straightforward:
1template<typename ADT> struct NATemp
2{using cases = std::variant<Num<ADT>, Add<ADT>>;};
3struct NA: NATemp<NA> {};
In Swierstra’s terminology [34], lines 1and 2define a recursive knot that
line 3ties. In the terminology of Oliveira et al. [26], NATemp is an object algebra
interface. That is because NATemp declares a set of algebraic signatures (namely,
that of Num and Add) but does not define (implement) them. In other words,
those signatures do not pertain to a fixed ADT.
5Integration of a Decentralised Pattern Matching
What matters to the C++ LMS-EPS is that NATemp underpins every ADT,
for which instances of Num<ADT> or Add<ADT> are valid terms. (Using γΦC0, one
denotes that by α. α CNum App.) Given NATemp, in line 3, we introduce NA
as an instance of such ADTs. That introduction is done in a specific way for F-
Bounding [4] commonly referred to in C++ as the Curiously Recurring Template
Pattern (CRTP) [36,§21.2]. See §5.2 for why employing CRTP is required here.
The nested type name cases at line 2above is what we used in the definition
of CsVar at line 2of Add in §3.1.
3.3 Functions
Just like that for ADTs, defining functions on ADTs takes two steps in the C++
First, for a function fon an ADT A, one implements an auxiliary func-
tion that takes a continuation as an argument. Suppose that one chooses the
name a_plus_f_matches for the auxiliary function. (See below for the in-
tention behind the naming of the auxiliary function.) Using the continuation,
a_plus_f_matches implements the raw pattern matching for fon every exten-
sion to A. Once called with fsubstituted for the continuation, a_plus_f_match-
es returns the pattern matching of f, now exclusively materialised for A.
Second, one implements fitself, which, by passing itself to
a_plus_f_matches, acquires the right pattern matching; and, then, visits f’s
parameter using the acquired pattern matching.
As an example for the above two steps, we implement below an evaluator for
NA expressions:
1template<typename ADT> auto na_plus_ev_matches(auto eval)
2{//na_plus_ev_matches<αCNum Add>
3return match {
4[] (const Num<ADT>& n) {return n.n_;}, //λNum(n). n
5//λAdd(l, r).eval (l) + eval (r)
6[eval](const Add<ADT>& a) {return eval(*a.l_) + eval(*a.r_);}
Above is the first step: na_plus_ev_matches is the auxiliary function for
evaluation. eval in line 1is the continuation. na_plus_ev_matches produces
the raw pattern matching for every ADT that extends NA. It does so by passing
match statements for Num<ADT> and Add<ADT> to the match combinator. In
line 4, a λ-abstraction is used for matching Num<ADT> instances. Line 6, on the
other hand, use a λ-abstraction to match Add<ADT> instances. The difference
is that the latter λ-abstraction is recursive and captures the variable eval (by
mentioning it between square brackets in line 6). Furthermore, rather than using
na_plus_ev_matches, it uses the continuation eval for recursion.
In short, the match combinator bundles a set of match statements together.
Such a match statement can be any callable C++ object. In this paper, we only
use λ-abstractions for our match statements. More on match in §5.3.
1int na_eval(const NA::cases& expr) {
2auto pm = na_plus_ev_matches<NA>(na_eval);
3return std::visit(pm, expr);
Above is the second step for provision of evaluation for NA expressions. In
line 2, it acquires the right pattern matching for NA by passing itself as the
continuation to na_plus_ev_matches. Then, in line 3, it visits the expression
to be evaluated using the acquired pattern matching.
We would like to end this subsection by emphasising on the following: In the
terminology of Oliveira et al. [26], na_plus_ev_matches is an object algebra.
In the latter work, compositionality of object algebras comes at the price of a
generalisation of self -references [27]. (In short, inside the body of an instance of
a given class, a self -references is a pointer/reference to the very instance itself.
Such a pointer/reference needs to also deal with virtual construction [8].) No-
tably, however, we achieve that (Section 4.1) without resorting to self -references.
3.4 Tests
Using the following two pieces of syntactic sugar for literals and addition
auto operator"" _n (unsigned long long n) {return Num<NA>(n);}
auto operator + (const NA::cases& l, const NA::cases& r) {return Add<NA>(l, r);}
na_eval(5_n + 5_n + 4_n) returns 14, as expected.
4 Addressing the EP Concerns
We now show how our technology is an EP solution.
4.1 E1 (Bidimensional Extensibility)
Extensibility in the dimension of ADTs is simple. Provided the Mul case com-
ponent below
1template<typename ADT> struct Mul{ //Mul α:: α×αα
2using CsVar = typename ADT::cases;
3Mul(const CsVar& l, const CsVar& r):
4l_(std::make_shared<CsVar>(l)), r_(std::make_shared<CsVar>(r)) {}
5const std::shared_ptr<CsVar> l_, r_;
encoding NAM using the C++ LMS-EPS can be done just like that for NA:
1template<typename ADT> struct NAMTemp
2{using cases = std::variant<Num<ADT>, Add<ADT>, Mul<ADT>>;};
3struct NAM: NAMTemp<NAM> {};
But, one can also extend NA to get NAM:
1template<typename ADT> struct NAMTemp
2{using cases = ext_variant_by_t<NATemp<ADT>, Mul<ADT>>;};
In the absence of a built-in extends for traits, that is the C++ LMS-EPS
counterpart for extending an ADT to another. See §5.4 for the definition of
Extensibility in the dimension of functions is not particularly difficult. For
example, here is how one does pretty printing for NA:
1template<typename ADT> auto na_plus_to_str_matches(auto to_string) {
2return match {
3[] (const Num<ADT>& n) {return std::to_string(n.n_);},
4[to_string](const Add<ADT>& a) {return to_string(*a.l_) + " + " +
8std::string na_to_string(const NA::cases& expr) {
9auto pm = na_plus_to_str_matches<NA>(na_to_string);
10 return std::visit(pm, expr);
11 }
na_plus_to_str_matches is the auxiliary function with to_string being
the continuation. na_to_string is the pretty printing for NA.
1template<typename ADT> auto nam_plus_to_str_matches(auto to_string) {
2return match {
4[to_string](const Mul<ADT>& m) {return to_string(*m.l_) + " *" +
On the other hand, the above auxiliary function called
nam_plus_to_str_matches reuses the match statements already developed by
na_plus_to_str_matches (line 3). It does so by including the latter function in
the list of match statements it includes in its match combinator. Note that the
former function, moreover, passes its own continuation (i.e., to_string) as an
argument to the latter function. Such a reuse is the C++ LMS-EPS counterpart
of the super call in line 5of NAM in §2.2.
The similarity becomes more clear when one observes that both the Scala
LMS-EPS and the C++ LMS-EPS scope the match statements and have mech-
anisms for reusing the existing ones. In the Scala LMS-EPS, the match state-
ments are scoped in a method of the base trait. That method, then, can be
overriden at the extension and reused via a super call. On the other hand, in
the C++ LMS-EPS, the match statements are scoped in the auxiliary functions.
That auxiliary function, then, can be mentioned in the match of the extension’s
auxiliary function (just like the new match statements), enabling its reuse.
4.2 E2 (Static Type Safety)
Suppose that in the pretty printing for NAM, one mistakenly employs
na_plus_to_str_matches instead of nam_plus_to_str_matches. (Note that
the latter name starts with nam whilst the former only starts with na.) That
situation is like when the programmer attempts pretty printing for a NAM ex-
pression without having provided the pertaining match statement of Mul. Here
is the erroneous code:
1std::string nam_to_string(const NAM::cases& expr) {//WRONG!
2auto pm = na_plus_to_str_matches<NAM>(nam_to_string);
3return std::visit(pm, expr);
As expected, the above code fails to compile. As an example, GCC 7.1
produces three error messages. In summary, those error messages state that
na_plus_to_str_matches only has match statements for Num and Add (but not
Mul). Note that the code fails to compile even without passing a concrete ar-
gument into nam_to_string. That demonstrates our strong static type safety.
The C++ LMS-EPS can guarantee that because the compiler chooses the right
match statement using overload resolution, i.e., at compile-time. C.f. §5.3 for
4.3 E3 (No Manipulation/Duplication)
Notice how nothing in the evidence for our support for E1 and E2 requires ma-
nipulation, duplication, or recompilation of the existing codebase. Our support
for E3 follows.
4.4 E4 (Separate Compilation)
Our support for E4, in fact, follows just like E3. It turns out, however, that C++
templates enjoy two-phase translation [36,§14.3.1]: Their parts that depend on
the type parameters are type checked (and compiled) only when they are instan-
tiated, i.e., when concrete types are substituted for all their type parameters. As
a result, type checking (and compilation) will be redone for every instantiation.
That type-checking peculiarity might cause confusion w.r.t. our support for E4.
In order to dispel that confusion, we need to recall that Add, for instance, is
a class template rather than a class. In other words, Add is not a type (because
it is of kind ∗→∗) but Add<NA> is. The interesting implication here is that
Add<NA> and Add<NAM> are in no way associated to one another. Consequently,
introduction of NAM in presence of NA, causes no repetition in type checking
(or compilation) of Add<NA>. (Add<NAM>, nonetheless, needs to be compiled in
presence of Add<NA>.) The same argument holds for every other case component
already instantiated with the existing ADTs.
More generally, consider a base ADT Φb=γand its extension Φe= (γ)
(γ0). Let #(γ) = nand #(γ0) = n0, where #(.) is the number of components in
the component combination. Suppose a C++ LMS-EPS codebase that contains
case components for γ1, . . . , γnand γ0
1, . . . , γ0
n0. Defining Φbin such a codebase
incurs compilation of ncase components. Defining Φeon top incurs compilation
of n+n0case components. Nevertheless, that does not disqualify our EP solution
because defining the latter component combination does not incur recompilation
of the former component combination. Note that individual components differ
from their combination. And, E4 requires the combinations not to be recompiled.
Here is an example in terms of DSL embedding. Suppose availability of a
type checking phase in a codebase built using the C++ LMS-EPS. Adding a
type erasure phase to that codebase, does not incur recompilation of the type
checking phase. Such an addition will, however, incur recompilation of the case
components common between the two phases. Albeit, those case components will
be recompiled for the type erasure phase. That addition leaves the compilation
of the same case components for the type checking phase intact. Hence, our
support for E4.
A different understanding from separate compilation is also possible, in which:
an EP solution is expected to, upon extension, already be done with the type
checking and compilation of the “core part” of the new ADT. Consider extending
NA to NAM , for instance. With that understanding, Num and Add are consid-
ered the “core part” of NAM . As such, the argument is that the type checking
and compilation of that “core part” should not be repeated upon the extension.
However, before instantiating Num and Add for NAM, both Num<NAM> and Add<
NAM> are neither type checked nor compiled. That understanding, hence, refuses
to take our work for an EP solution. We find that understanding wrong because
the core of NAM is NA, i.e., the Num Add combination, as opposed to both
Num and Add but individually. Two quotations back our mindset up:
The definition Zenger and Odersky [20] give for separate compilation is as
follows: “Compiling datatype extensions or adding new processors should not
encompass re-type-checking the original datatype or existing processors [func-
tions].” The datatypes here are NA and NAM. Observe how compiling NAM does
not encompass repetition in the type checking and compilation of NA.
Wang and Oliveira [38] say an EP solution should support: “software evo-
lution in both dimensions in a modular way, without modifying the code that
has been written previously.” Then, they add: “Safety checks or compilation
steps must not be deferred until link or runtime.” Notice how neither definition
of new case components or ADTs, nor addition of case components to existing
ADTs to obtain new ADTs, implies modification of the previously written code.
Compilation or type checking of the extension is not deferred to link or runtime
For more elaboration on the take of Wang and Oliveira on (bidimensional)
modularity, one may ask: If NA’s client becomes a client of NAM , will the
client’s code remain intact under E3 and E4? Let us first disregard code that is
exclusively written for NA for it is not meant for reuse by NAM :
void na_client_f(const NA&) {...}
If on the contrary, the code only counts on the availability of Num and Add:
1template <
2typename ADT, typename = std::enable_if_t<adt_contains_v<ADT, Num, Add>>
3>void na_plus_client_f(const ADT& x) {...}
Then, it can expectedly be reused upon transition from NA to NAM . (We drop
the definition of adt_contains_v due to space restrictions.)
5 Technicality
5.1 Why std::shared ptr?
Although not precisely what the C++ specification states, it is not uncommon
for the current C++ compilers to require the types participating in the formation
of a std::variant to be default-constructable. That requirement is, however,
not fulfilled by our case components. As a matter of fact, ADT cases, in general,
are unlikely to fulfil that requirement.
But, as shown in line 2of NATemp, the C++ LMS-EPS needs the case com-
ponents to participate in a std::variant. Wrapping the case components in a
default-constructable type seems inevitable. We choose to wrap them inside a
std::shared_ptr because, then, we win sub-expression sharing as well.
5.2 Why CRTP?
The reader might have noticed that, in the C++ LMS-EPS, defining ADTs is
also possible without CRTP. For example, one might try the following for NA:
1struct OtherNA { using cases = std::variant<Num<OtherNA>, Add<OtherNA>>; };
Then, extending OtherNA to an encoding for NAM will, however, not be
possible as we extended NATemp to NAMTemp in §4.1. In addition to employing
a different extension metafunction than ext_variant_by_t in §5.4, we would
need some extra work in the case components. For example, here is how to enrich
1template<typename ADT> struct Add
2{/*... like before ... */template<typename A> using case_component = Add<A>;};
Then, we can still extend NATemp to get NAM:
1struct NAM { using cases = ext_variant_by_t<NATemp<NAM>, Mul<NAM>>; };
If one wishes to, it is even possible to completely abolish NATemp – and, in
fact, all the CRTP:
1struct NAM { using cases = ext_to_by_t<NA, NAM, Mul<NAM>>; };
where ext_to_by_t is defined in §5.4.
5.3 The match Combinator
The definition of our match combinator is as follows6:
1template<typename... Ts> struct match: Ts...
2{using Ts::operator()...;};
3template<typename... Ts> match(Ts...) -> match<Ts...>;
As one can see above, match is, in fact, a type parameterised struct. In
lines 1and 2above, match derives from all its type arguments. At line 2, it also
makes all the operator()s of its type arguments accessible via itself. Accord-
ingly, match is callable in all ways its type arguments are.
Line 3uses a C++ feature called user-defined deduction guides. Recall that
C++ only offers automatic type deduction for template functions. Without
line 3, thus, match is only a struct, missing the automatic type deduction. The
programmer would have then needed to list all the type arguments explicitly to
instantiate match. That would have been cumbersome and error-prone – espe-
cially, because those types can rapidly become human-unreadable. Line 3helps
6This is a paraphrase of the overloaded combinator taken from the std::visit’s
online specification at the C++ Reference:
the compiler to deduce type arguments for the struct (i.e., the match to the
right of “->”) in the same way it would have done that for the function (i.e., the
match to the left of “->”).
One may wonder why we need all those Ts::operator ()s. The reason is
that, according to the C++ specification, the first argument of std::visit
needs to be a callable. The compiler tries the second std::visit argument
against all the call pieces of syntax that the first argument provides. The mech-
anism is that of C++’s overload resolution. In this paper, we use match only for
combining λ-abstractions. But, all other sorts of callable are equally acceptable
to match.
Finally, we choose to call match a combinator because, to us, its usage is
akin to Haskell’s Combinator Pattern [28,§16].
5.4 Definitions of ext variant by t and ext to by t
Implementation of ext_variant_by_t is done using routine template metapro-
1template<typename,typename...> struct evb_helper;
2template<typename... OCs, typename... NCs>
3struct evb_helper<std::variant<OCs...>, NCs...>
4{using type = std::variant<OCs..., NCs...>;};
5template<typename ADT, typename... Cs> struct ext_variant_by
6{using type = typename evb_helper<typename ADT::cases, Cs...>::type;};
7template<typename ADT, typename... Cs>
8using ext_variant_by_t = typename ext_variant_by<ADT, Cs...>::type;
ext_variant_by_t (line 8) extends an ADT by the cases Cs.... To that
end, ext_variant_by_t is a syntactic shorthand for the type nested type of
ext_variant_by.ext_variant_by (line 5) works by delegating its duty to
evb_helper after acquiring the case list of ADT (line 6). Given a std::variant
of old cases (OCs...) and a series of new cases (NCs...), the metafunction
evb_helper type-evaluates to a std::variant of old and new cases (line 4).
Implementing ext_to_by_t is not particularly more complicated. So, we
drop explanation and only provide the code:
1template<typename,typename>struct materialise_for_helper;
2template<typename ADT, typename... Cs>
3struct materialise_for_helper<ADT, std::variant<Cs...>>
4{using type = std::variant<typename Cs::template case_component<ADT>...>;};
6template<typename ADT1, typename ADT2> struct materialise_for {
7using type = typename materialise_for_helper<ADT2, typename ADT1::cases>::type;
10 template<typename ADT1, typename ADT2, typename... Cs> struct ext_to_by {
11 using type = typename evb_helper<typename materialise_for<ADT1, ADT2>::type,
12 Cs...>::type;
13 };
15 template<typename ADT1, typename ADT2, typename... Cs>
16 using ext_to_by_t = typename ext_to_by<ADT1, ADT2, Cs...>::type;
6 C4C versus GADTs
When embedding DSLs, it is often convenient to piggyback on the host lan-
guage’s type system. In such a practice, GADTs are a powerful means to guar-
antee the absence of certain type errors. For example, here is a Scala translit-
eration7of the running example Kennedy and Russo [18] give for GADTs in
object-oriented languages:
1sealed abstract class Exp[T]
2case class Lit(i: Int) extends Exp[Int]
3case class Plus(e1: Exp[Int], e2: Exp[Int]) extends Exp[Int]
4case class Equals(e1: Exp[Int], e2: Exp[Int]) extends Exp[Boolean]
5case class Cond(e1: Exp[Boolean], e2: Exp[Int], e3: Exp[Int]) extends Exp[Int]
6/*... more case classes ... */
7def eval[T](exp: Exp[T]): T = exp match {...}
Notice first that Exp is type parameterised, where Tis an arbitrary Scala type.
That is how Lit can derive from Exp[Int] whilst Equals derives from Exp[
Boolean]. Second, note that Plus takes two instances of Exp[Int]. Contrast
that with the familiar encodings of α=Plus(α, α)|. . . , for some ADT α. Unlike
the GADT one, the latter encoding cannot outlaw nonsensical expressions such
as Plus(Lit(5), Lit(true)). Third, note that eval is polymorphic in the
carrier type of Exp, i.e., T.
The similarity between the above case definitions and our case components
is that they are both type parameterised. Nevertheless, the former are parame-
terised by the type of the Scala expression they carry. Whereas, our case com-
ponents are parameterised by their ADT types. The impact is significant. Sup-
pose, for example, availability of a case component Boolwith the corresponding
operator ""_b syntactic sugar. In their current presentation, our case compo-
nents cannot statically outlaw type-erroneous expressions like 12_n + "true"
_b. On the other hand, the GADT Cond is exclusively available as an ADT case
of Exp and cannot be used for other ADTs.
Note that, so long as statically outlawing 12_n + "true"_b is the concern,
one can always add another layer in the Exp grammar so that the integral cases
and Boolean cases are no longer at the exact same ADT. That workaround,
however, will soon become unwieldy. That is because, it involves systematically
separating syntactic categories for every carrier type – resulting in the craft of
a new type system. GADTs employ the host language’s type system instead.
The bottom line is that GADTs and C4C encodings of ADTs are orthogonal.
One can always generalise our case components so they too are parameterised
by their carrier types and so they can guarantee similar type safety.
7 Related Work
The support of the Scala LMS-EPS for E2 can be easily broken using an in-
complete pattern matching. Yet, given that Scala pattern matching is dynamic,
7Posted online by James Iry on Wed, 22/10/2008 at
whether LMS really relaxes E2 is debatable. Note that the problem in the Scala
LMS-EPS is not an “Inheritance is not Subtyping” one [7]: The polymorphic
function of a deriving trait does specialise that of the base.
In comparison to the Scala LMS-EPS, we require one more step for defining
ADTs: the CRTP. Nevertheless, given that the C++ LMS-EPS is C4C, spec-
ifying the cases of an ADT is by only listing the right case components in a
std::variant. Defining functions on ADTs also requires one more step in the
C++ LMS-EPS: using the continuation. When extending a function for new
ADT cases, the C++ LMS-EPS, however, needs no explicit super call, as re-
quired by the Scala LMS-EPS.
A note on the Expression Compatibility Problem is appropriate here. As de-
tailed earlier [11,§4.2], the Scala LMS-EPS cannot outlaw incompatible exten-
sions. Neither can the current presentation of the C++ LMS-EPS. Nonetheless,
due to its C4C nature, that failure is not inherent in the C++ LMS-EPS. One
can easily constrain the ADT type parameter of the case components in a similar
fashion to the Scala IDPaM [16] to enforce compatibility upon extension.
The first C4C solution to the EP is the Scala IDPaM [16]. ADT creation in
the Scala IDPaM too requires F-Bounding. But, the type annotation required
when defining an ADT using their case components is heavier.
In the Scala IDPaM, the number of type annotations required for a function
taking an argument of an ADT with ncases is O(n). That is O(1) in the C++
LMS-EPS. The reason is that, in C++, with programming purely at the type
level, types can be computed from one another. In particular, an ADT’s case list
can be computed programmatically from the ADT itself. That is not possible in
Scala without implicits, which are not always an option. In the Scala IDPaM
too, implementation of functions on ADTs is nominal: For every function on a
given ADT α, all the corresponding match components—i.e., match statements
also delivered as components— of α’s cases need to be manually mixed in to form
the full function implementation. The situation is similar for the C++ LMS-EPS
in that all the match statements are required to be manually listed in the match
combinator. However, instead of using a continuation, in the Scala IDPaM, one
mixes in a base case as the last match component. Other than F-Bounding, the
major language feature required for the Scala IDPaM is stackability of traits.
In the C++ LMS-EPS, that is variadic templates. The distinctive difference
between the C++ LMS-EPS and the Scala IDPaM is that the latter work
relaxes E2 in the absence of a default [39]. On the contrary, the C++ LMS-
EPS guarantees strong static type safety.
The second C4C solution to the EP is the C++ IDPaM [12]. There are
two reasons to prefer the C++ IDPaM over the C++ LMS-EPS: Firstly, in
the C++ IDPaM, definition of a function fon ADTs amounts to provision of
simple (C++) function overloads, albeit plus a one-off macro instantiation for f.
(Those function overloads are called match components of the C++ IDPaM.)
Secondly, in the C++ IDPaM, function definition is structural: Suppose the
availability of all the corresponding match components of α’s case list and the
macro instantiation for f. Then, unlike the C++ LMS-EPS, to define fon α, the
programmer need not specify which match statements to include in the pattern
matching. The compiler deductively obtains the right pattern matching using
α’s structure, i.e. α’s case list.
There are two reasons to prefer the C++ LMS-EPS over the C++ IDPaM.
Firstly, implementing ADTs and functions on them is only possible in the C++
IDPaM using a metaprogramming facility shipped as a library. That library was
so rich in its concepts that it was natural to extend [13] for multiple dispatch.
Behind the scenes, the library performs iterative pointer introspection to choose
the right match statements. In the C++ LMS-EPS, that pointer introspection
is done using the compiler’s built-in support for std::variant. That saves
the user from having to navigate the metaprogramming library upon mistakes
(or bugs). Furthermore, when it comes to orchestrating the pattern matching,
the compiler is likely to have more optimisation opportunities than the library.
Secondly, unlike their C++ IDPaM equivalents, case components of the C++
LMS-EPS do not inherit from their ADT. This entails weaker coupling between
case components and ADT definitions.
Instead of std::variant, one can use boost::variant8to craft a simi-
lar solution to the C++ LMS-EPS. Yet, the solution would have not been as
clean with its auxiliary functions as here. In essence, for a function f, one would
have needed to manually implement each match statement as a properly-typed
overload of F::operator (). Extending fto handle new ADT cases, neverthe-
less, would have been more akin to the Scala LMS-EPS. That is because, then,
providing the new match statements would have amounted to implementing
the corresponding FExtended::operator () overloads, for some FExtended
that derives from F. (Compare with §2.2.) Moreover, boost::variant requires
special settings for working with recursive types (such as ADTs) that damage
Using object algebras [10] to solve EP has become popular over recent years.
Oliveira and Cook [23] pioneered that. Oliveira et al. [26] address some awkward-
ness issues faced upon composition of object algebras. Rendel, Brachth¨auser and
Ostermann [30] add ideas from attribute grammars to get reusable tree traver-
sals. As also pointed out by Black [3], an often neglected factor about solutions
to EP is the complexity of term creation. That complexity increases from one
work to the next in the above literature. The symptom develops to the extent
that it takes Rendel, Brachth¨auser and Ostermann 12 non-trivial lines of code to
create a term representing “3+5”. Of course, those 12 lines are not for the latter
task exclusively and enable far more reuse. Yet, those 12 lines are inevitable for
term creation for “3 + 5”, making that so heavyweight. The latter work uses au-
tomatic code generation for term creation. So, the ADT user has a considerably
more involved job using the previous object algebras technologies for EP than
that of ours. Additionally, our object algebras themselves suffer from much less
syntactic noise. Defining functions on ADTs is slightly more involved in the C++
LMS-EPS than object algebras for the EP. For example, pretty-printing for NA
8 67 0/doc/html/variant.html
takes 12 (concise) Scala lines in the latter work, whereas that is 14 (syntactically
noisy) C++ lines in ours.
Garrigue [9] solves EP using global case definitions that, at their point of
definition, become available to every ADT defined afterwards. Per se, a function
that pattern matches on a group of these global cases can serve any ADT con-
taining the selected group. OCaml’s built-in support for Polymorphic Variants
[9] makes definition of both ADTs and functions on them easier. However, we
minimise the drawbacks [3] of ADT cases being global by promoting them to
Swierstra’s Datatypes `a la Carte [34] uses Haskell’s type classes to solve EP.
In his solution too, ADT cases are ADT-independent but ADT-parameterised.
He uses Haskell Functors to that end. Defining functions on ADTs amounts
to defining a type class, instances of which materialising match statements for
their corresponding ADT cases. Without syntactic sugaring, term creation can
become much more involved than that for ordinary ADTs of Haskell. Defining
the syntactic sugar takes many more steps than us, but, makes term creation
straightforward. Interestingly enough, using the Scala type classes [24] can lead
to simpler syntactic sugar definition but needs extra work for the lack of direct
support in Scala for type classes. In his machinery, Swierstra offers a match that
is used for monadically inspecting term structures.
Bahr and Hvitved extend Swierstra’s work by offering Compositional
Datatypes (CDTs) [1]. They aim at higher modularity and reusability. CDTs
support more recursion schemes, and, extend to mutually recursive data types
and GADTs. Besides, syntactic sugaring is much easier using CDTs because
smart constructors can be automatically deduced for terms.
Later on, they offer Parametric CDTs (PCDTs) [2] for automatic α-equivale-
nce and capture-avoiding variable bindings. PCDTs achieve that using Difun-
tors [19] (instead of functors) and a CDT encoding of Parametric Higher-Order
Abstract Syntax [5]. Case definitions take two phases: First an equivalent of
our case components need to be defined. Then, their case components need to
be materialised for each ADT, similar to but different from that of Haeri and
Schupp [14,11].
The distinctive difference between C4C and the works of Swierstra, Bahr,
and Hvitved is the former’s inspiration by CBSE. Components, in their CBSE
sense, ship with their ‘requires’ and ‘provides’ interfaces. Whereas, even though
the latter works too parametrise cases by ADTs, the interface that CDTs, for
instance, define do not go beyond algebraic signatures. Although we do not
present those for C++ LMS-EPS here, C4C goes well beyond that, enabling easy
solutions to the Expression Families Problem [23] and Expression Compatibility
Problem [16] as well as GADTs. The respective article is in submission.
8 Conclusion
In this paper we show how a new C4C encoding of ADTs in C++ can solve EP
in a way that is reminiscent to the Scala LMS-EPS. On its way, our solution
gives rise to simple encodings for object algebras and object algebra interfaces
and relates to Datatypes `a la Carte ADT encodings.
Given the simplicity of our encoding for object algebras and object algebra
interfaces in the absence of heavy notation for term creation, an interesting fu-
ture work is mimicking the earlier research on object algebra encodings for EP.
We need to investigate whether our technology still remains simple when we
take all the challenges those works take. Another possible future work is exten-
sion of our (single dispatch) mechanism for implementing functions on ADTs
to multiple dispatch. Of course, C++ LMS-EPS needs far more experimenta-
tion with real-size test cases to study its scalability. Finally, we are working on
a C++ LMS-EPS variation that, unlike our current presentation, structurally
implements functions on ADTs. The latter variation has thus far presented itself
as a promising vehicle for also delivering multiple dispatch.
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A C++ Features Used
A C++ struct (or class) can be type parameterised. The struct Sbelow,
for example, takes two type parameters T1 and T2:
template<typename T1, typename T2> struct S {...};
Likewise, C++ functions can take type parameters:
template<typename T1, typename T2> void f(T1 t1, T2 t2) {...}
From C++20 onward, certain type parameters need not to be mentioned
explicitly. For example, the above function fcan be abbreviated as:
void f(auto t1, auto t2) {...}
A (template or non-template)struct can define nested type members. For
example, the struct Tbelow defines T::type to be int:
struct T {using type = int;};
Nested types can themselves be type parameterised, like Y::template type:
struct Y {template<typename>using type = int;};
C++17 added std::variant as a type-safe representation for unions. An in-
stance of std::variant, at any given time, holds a value of one of its alternative
types. That is, the static type of such an instance is that of the std::variant it
is defined with; whilst, the dynamic type is one and only one of those alternative
types. As such, a function that is to be applied on a std::variant needs to
be applicable to its alternative types. Technically, a visitor is required for the
alternative types. The function std::visit, takes a visitor in addition to a pack
of arguments to be visited.
auto twice = [](int n){return n*2;}
The variable twice above is bound to a λ-abstraction that, given an int,
returns its value times two. λ-abstractions can also capture unbound names. In
such a case, the captured name needs to be mentioned in the opening square
brackets before the list of parameters. For example, the λ-abstraction times
below captures the name m:
auto times = [m](int n){return n*m;}
A common approach to defining domain-specific languages (DSLs) is via a direct embedding into a host language. There are several well-known techniques to do such embeddings, including shallow and deep embeddings. However, such embeddings come with various trade-offs in existing programming languages. Owing to such trade-offs, many embedded DSLs end up using a mix of approaches in practice, requiring a substantial amount of code, as well as some advanced coding techniques. In this paper, we show that the recently proposed Compositional Programming paradigm and the CP language provide improved support for embedded DSLs. In CP we obtain a new form of embedding, which we call a compositional embedding, that has most of the advantages of both shallow and deep embeddings. On the one hand, compositional embeddings enable various forms of linguistic reuse that are characteristic of shallow embeddings, including the ability to reuse host-language optimizations in the DSL and add new DSL constructs easily. On the other hand, similarly to deep embeddings, compositional embeddings support definitions by pattern matching or dynamic dispatching (including dependent interpretations, transformations, and optimizations) over the abstract syntax of the DSL and have the ability to add new interpretations. We illustrate an instance of compositional embeddings with a DSL for document authoring called ExT. The DSL is highly flexible and extensible, allowing users to create various non-trivial extensions easily. For instance, ExT supports various extensions that enable the production of wiki-like documents, LaTeX documents, vector graphics or charts. The viability of compositional embeddings for ExT is evaluated with three applications.
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This paper presents a novel and simple solution to Wadler’s Expression Problem that works in conventional object-oriented languages. Unlike all existing solutions in Java-like languages, this new solution does not use any kind of generics: it relies only on subtyping. The key to the solution is the use of covariant type refinement of return types (or fields): a simple feature available in many object-oriented languages, but not as widely known or used as it should be. We believe that our results present valuable insights for researchers and programming language designers interested in extensibility. Furthermore our results have immediate applicability as practical design patterns for programmers interested in improving extensibility of their programs.
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The "Expression Problem" was brought to prominence by Wadler in 1998. It is widely regarded as illustrating that the two mainstream approaches to data abstraction---procedural abstraction and type abstraction---are complementary, with the strengths of one being the weaknesses of the other. Despite an extensive literature, the origin of the problem remains ill-understood. I show that the core problem is in fact the use of global constants, and demonstrate that an important aspect of the problem goes away when Java is replaced by a language like Grace, which eliminates them.
This paper presents Modular Reifiable Matching (MRM): a new approach to two level types using a fixpoint of list-of-functors representation. MRM allows the modular definition of datatypes and functions by pattern matching, using a style similar to the widely popular Datatypes 'a la Carte (DTC) approach. However, unlike DTC, MRM uses a fixpoint of list-of-functors approach to two-level types. This approach has advantages that help with various aspects of extensibility, modularity and reuse. Firstly, modular pattern matching definitions are collected using a list of matches that is fully reifiable. This allows for extensible pattern matching definitions to be easily reused/inherited, and particular matches to be overridden. Such flexibility is used, among other things, to implement extensible generic traversals. Secondly, the subtyping relation between lists of functors is quite simple, does not require backtracking, and is easy to model in languages like Haskell. MRM is implemented as a Haskell library, and its use and applicability are illustrated through various examples in the paper.
Oliveira and Cook (2012) and Oliveira et al. (2013) have recently introduced object algebras as a program structuring technique to improve the modularity and extensibility of programs. We analyze the relationship between object algebras and attribute grammars (AGs), a formalism to augment context-free grammars with attributes. We present an extension of the object algebra technique with which the full class of L-attributed grammars - an important class of AGs that corresponds to one-pass compilers - can be encoded in Scala. The encoding is modular (attributes can be defined and type-checked separately), scalable (the size of the encoding is linear in the size of the AG specification) and compositional (each AG artifact is represented as a semantic object of the host language). To evaluate these claims, we have formalized the encoding and re-implemented a onepass compiler for a subset of C with our technique. We also discuss how advanced features of modern AG systems, such as higher-order and parameterized attributes, reference attributes, and forwarding can be supported.