A Uniform Account of Realizability in Abstract Argumentation

A Uniform Account of Realizability in Abstract Argumentation

Thomas Linsbichler

and J¨org P ¨uhrer

and Hannes Strass

Abstract. We introduce a general framework for analyzing real- izability in abstract dialectical frameworks (ADFs) and various of its subclasses. In particular, the framework applies to Dung argu- mentation frameworks, SETAFs by Nielsen and Parsons, and bipolar ADFs. We present a uniform characterization method for the ad- missible, complete, preferred and model/stable semantics. We em- ploy this method to devise an algorithm that decides realizability for the mentioned formalisms and semantics; moreover the algorithm al- lows for constructing a desired knowledge base whenever one exists.

The algorithm is built in a modular way and thus easily extensible to new formalisms and semantics. We have implemented our approach in answer set programming, and used the implementation to obtain several novel results on the relative expressiveness of the abovemen- tioned formalisms.

1 Introduction

The abstract argumentation frameworks (AFs) introduced by Dung [9] have garnered increasing attention in the recent past. In his seminal paper, Dung showed how an abstract notion of argu- ment (seen as an atomic entity) and the notion of individual attacks between arguments together could reconstruct several established KR formalisms in argumentative terms. Despite the generality of those and many more results in the field that was sparked by that pa- per, researchers also noticed that the restriction toindividual attacks is often overly limiting, and devised extensions and generalizations of Dung’s frameworks: directions included generalizing individual attacks tocollective attacks[23], leading to so-called SETAFs; oth- ers started offering asupportrelation between arguments [8], prefer- ences among arguments [1, 22], or attacks on attacks into arbitrary depth [2]. This is only the tip of an iceberg, for a more comprehensive overview we refer to the work of Brewka, Polberg, and Woltran [5].

One of the most recent and most comprehensive generalizations of AFs has been presented by Brewka and Woltran [6] (and later continued by Brewka et al. [4]) in the form ofabstract dialectical frameworks (ADFs). These ADFs offer any type of link between arguments: individual attacks (as in AFs), collective attacks (as in SETAFs), and individual and collective support, to name only a few.

This generality is achieved through so-calledacceptance conditions associated to each statement. Roughly, the meaning of relationships between arguments is not fixed in ADFs, but is specified by the user for each argument in the form of Boolean functions (accept- ance functions) on the argument’s parents. However, this generality comes with a price: Strass and Wallner [29] found that the complex- ity of the associated reasoning problems of ADFs is in general higher

1Institute of Information Systems, TU Wien, Vienna, Austria

2Computer Science Institute, Leipzig University, Leipzig, Germany

than in AFs (one level up in the polynomial hierarchy). Fortunately, the subclass ofbipolar ADFs(defined by Brewka and Woltran [6]) is as complex as AFs (for all considered semantics) while still offering a wide range of modeling capacities [29]. However, there has only been little concerted effort so far to exactly analyze and compare the expressiveness of the abovementioned languages.

This paper is about exactly analyzing means of expression for ar- gumentation formalisms. Instead of motivating expressiveness in nat- ural language and showing examples that some formalisms seem to be able to express but others do not, we tackle the problem in a formal way. We use a precise mathematical definition of expressiveness: a set of interpretations isrealizableby a formalism under a semantics if and only if there exists a knowledge base of the formalism whose semantics is exactly the given set of interpretations. Studying real- izability in AFs has been started by Dunne et al. [11, 10], who ana- lyzed realizability for extension-based semantics, that is, interpreta- tions represented by sets where arguments are either accepted (in the extension set) or not accepted (not in the extension set). While their initial work disregarded arguments that are never accepted, there have been continuations where the existence of such “invisible” ar- guments is ruled out [3, 20]. Dyrkolbotn [12] began to analyze realiz- ability for labeling-based semantics of AFs, that is, three-valued se- mantics where arguments can be accepted (mapped to true), rejected (mapped to false) or neither (mapped to unknown). Strass [28] star- ted to analyze the relative expressiveness of two-valued semantics for ADFs (relative with respect to related formalisms). Most re- cently, P¨uhrer [26] presented precise characterizations of realizab- ility for ADFs under several three-valued semantics, namely admiss- ible, grounded, complete, and preferred. The term “precise charac- terizations” means that he gave necessary and sufficient conditions for an interpretation set to be ADF-realizable under a semantics.

The present paper continues this line of work by lifting it to a much more general setting. We combine the works of Dunne et al. [10], P¨uhrer [26], and Strass [28] into a unifying framework, and at the same time extend them to formalisms and semantics not considered in the respective papers: we treat several formalisms, namely AFs, SETAFs, and (B)ADFs, while the previous works all used differ- ent approaches and techniques. This is possible because all of these formalisms can be seen as subclasses of ADFs that are obtained by suitably restricting the acceptance conditions.

Another important feature of our framework is that we uniformly use three-valued interpretations as the underlying model theory. In particular, this means that arguments cannot be “invisible” any more since the underlying vocabulary of arguments is always implicit in each interpretation. Technically, we always assume a fixed underly- ing vocabulary and consider our results parametric in that vocabu- lary. In contrast, for example, Dyrkolbotn [12] presents a construc-

tion for realizability that introduces new arguments into the realizing knowledge base; we do not allow that. While sometimes the intro- duction of new arguments can make sense, for example if new in- formation becomes available about a domain or a debate, it is not sensible in general, as these new arguments would be purely tech- nical with an unclear dialectical meaning. Moreover, it would lead to a different notion of realizability, where most of the realizability problems would be significantly easier, if not trivial.

The paper proceeds as follows. We begin with recalling and intro- ducing the basis and basics of our work – the formalisms we analyze and the methodology with which we analyze them. Next we intro- duce our general framework for realizability; the major novelty is our consistent use of so-called characterization functions, firstly in- troduced by P¨uhrer [26], which we adapt to further semantics. The main workhorse of our approach will be a parametric propagate-and- guess algorithm for deciding whether a given interpretation set is realizable in a formalism under a semantics. We then analyze the rel- ative expressiveness of the considered formalisms, presenting several new results that we obtained using an implementation of our frame- work. We conclude with a discussion.

2 Preliminaries

We make use of standard mathematical concepts like functions and partially ordered sets. For a functionf:X→Y we denote theup- date offwith a pair(x, y)∈X×Y byf|^xy:X→Y withz7→y ifz=x, andz7→f(z)otherwise. For a functionf:X→Y and y∈Y, its preimage isf⁻¹(y) ={x∈X|f(x) =y}. Apartially ordered setis a pair(S,v)withva partial order onS. A partially ordered set(S,v)is acomplete latticeif and only if everyS⁰⊆S has both a greatest lower bound (glb)d

S⁰∈S and a least upper bound (lub)F

S⁰∈S. A partially ordered set(S,v)is acomplete meet-semilatticeiff every non-empty subsetS⁰⊆S has a greatest lower boundd

S⁰∈S(themeet) and every ascending chainC⊆S has a least upper boundF

C∈S.

Three-Valued Interpretations LetAbe a fixed finite set of state- ments. Aninterpretationis a mappingv:A→ {t,f,u}that assigns one of the truth values true (t), false (f) or unknown (u) to each statement. An interpretation istwo-valuedifv(A)⊆ {t,f}, that is, the truth valueuis not assigned. Two-valued interpretationsvcan be extended to assign truth valuesv(ϕ)∈ {t,f}to propositional for- mulasϕas usual.

The three truth values are partially ordered according to their in- formation content: we haveu<itandu<ifand no other pair in

<i, which intuitively means that the classical truth values contain more information than the truth value unknown. As usual, we denote by≤ithe partial order associated to the strict partial order<i. The pair({t,f,u},≤i)forms a complete meet-semilattice with the in- formation meet operationui. This meet can intuitively be interpreted asconsensusand assignstuit=t,fuif =f, and returnsuoth- erwise.

The information ordering≤iextends in a straightforward way to interpretationsv1, v2overAin thatv1≤iv2iffv1(a)≤iv2(a)for alla∈A. We say for two interpretationsv1, v2thatv2extendsv1

iffv1≤iv2. The setVof all interpretations overAforms a com- plete meet-semilattice with respect to the information ordering≤i. The consensus meet operation ui of this semilattice is given by (v1uiv2)(a) =v1(a)uiv2(a)for alla∈A. The least element of (V,≤i)is the valuationvu:A→ {u}mapping all statements to un- known – the least informative interpretation. ByV2we denote the set of two-valued interpretations; they are the≤i-maximal elements of

the meet-semilattice(V,≤i). We denote by[v]2the set of all two- valued interpretations that extendv. The elements of[v]2 form an

≤i-antichain with greatest lower boundv=d

i[v]2.

Abstract Argumentation Formalisms An abstract dialectical framework (ADF)is a tuple D= (A, L, C) whereA is a set of statements (representing positions one can take or not take in a debate), L⊆A×A is a set of links (representing dependencies between the positions),C={Ca}_a∈A is a collection of functions Ca: 2^par(a)→ {t,f}, one for each statement a∈A. The func- tion Ca is theacceptance condition of a and expresses whether a can be accepted, given the acceptance status of its parents par(a) ={b∈S|(b, a)∈L}. We usually represent eachCaby a propositional formulaϕaoverpar(a). For the acceptance condition Ca, we takeCa(M∩par(a)) =tto hold iffMis a model ofϕa.

Brewka and Woltran [6] introduced a useful subclass of ADFs: an ADFD= (A, L, C)isbipolariff all links inLare supporting or attacking (or both). A link(b, a)∈Lissupporting inDiff for all M⊆par(a), we have thatCa(M) =timpliesCa(M∪ {b}) =t.

Symmetrically, a link (b, a)∈L is attacking in D iff for all M⊆par(a), we have thatCa(M∪ {b}) =timpliesCa(M) =t.

Intuitively, a link(b, a)∈Lis supporting iff it can never be the case that there is some state of affairs where we acceptaand rejectb, but after additionally also acceptingbdo not acceptaany more. Sym- metrically, a link(b, a)∈Lis attacking iff it can never be the case that we rejectaandb, but after acceptingbalso accepta. If a link (b, a)is both supporting and attacking thenbhas no actual influence ona. (But the link does not violate bipolarity.) We write BADFs asD= (A, L⁺∪L⁻, C)and mean thatL⁺contains all supporting links andL⁻all attacking links; see also Example 1 below.³

The semantics of ADFs can be defined using an operatorΓDover three-valued interpretations [6, 4]. For an ADFDand a three-valued interpretationv, the interpretationΓD(v)is given by

a7→d

i{w(ϕa)|w∈[v]2}

That is, for each statementa, the operator returns the consensus truth value for its acceptance formulaϕa, where the consensus takes into account all possible two-valued interpretationswthat extend the in- put valuationv. If thisvis two-valued, we get[v]2={v}and thus ΓD(v)(a) =v(ϕa).

The standard semantics of ADFs are now defined as follows. For ADFD, an interpretationv:A→ {t,f,u}is

• admissibleiffv≤iΓD(v);

• completeiffΓD(v) =v;

• preferrediff it is≤i-maximal admissible;

• atwo-valued modeliff it is two-valued andΓD(v) =v.

We denote the sets of interpretations that are admissible, complete, preferred, and two-valued models byadm(D), com(D), prf(D) andmod(D), respectively. These definitions are proper generaliz- ations of Dung’s notions for AFs: For an AF(A, R), whereR ⊆ A ×A is the attack relation, the ADF associated to (A, R) is D_(A,R)= (A, R, C)withC={ϕa}a∈A andϕa=V

b:(b,a)∈R¬b for alla∈A. AFs inherit their semantics from the definitions for ADFs [4, Theorems 2 and 4]. In particular, an interpretation isstable for an AF(A, R)if and only if it is a two-valued model ofD(A,R).

3Other than a part of the name, there is no relationship of bipolar ADFs with the bipolar framework of Cayrol and Lagasquie-Schiex [8]; Brewka and Woltran gave a more detailed comparison of the two formalisms [6].

Example 1. Consider the bipolar ADFD= (A, L⁺∪L⁻, C)over vocabularyA={a, b, c}with

ϕa=b∧c, ϕb=¬a, ϕc=a∨ ¬b

whence it follows that L⁺={(b, a),(c, a),(a, c)} and L⁻={(a, b),(b, c)}. (The types of links can be read off the polarities of the statements in the acceptance formulas [28, The- orem 1]; statements occurring only positively are supporting, those that occur only negatively are attacking.) Intuitively, the acceptance conditionϕais a group support:acan only be accepted ifbothb andcare accepted. Forb, we have an individual attack just like in standard AFs:bis attacked bya, and therefore only be accepted ifa is not accepted. The acceptance condition ofcconsists of a support byathat overpowers an attack by b; in other words, to be able to acceptc, the support froma must be present or the attack fromb must be absent, and if both are present then the support is stronger.

(We could have specified that the attack is stronger than the support by writingϕc=a∧ ¬b.) Regarding the semantics ofD, we find that mod(D) =prf(D) ={v1} with v1={a7→f, b7→t, c7→f}.

Furthermore, we have adm(D) =com(D) =prf(D)∪ {v2} wherev2={a7→u, b7→u, c7→u}. Intuitively, setting all state- ments to u is always admissible; in this case it is also complete because no statement is unconditionally accepted or rejected. The non-trivial interpretation v1 is a model of the BADF because intuitively:ais rejected since it misses the support ofc;bis accepted because the attack fromadoes not materialize;cis rejected because it misses support fromaand at the same time is attacked byb.

A SETAF is a pairS= (A, X)whereX ⊆(2^A\ {∅})×Ais the (set) attack relation. We define three-valued counterparts of the se- mantics introduced by Nielsen and Parsons [23], following the same conventions as in three-valued semantics of AFs [7] and argument- ation formalisms in general. Given a statementa∈Aand an inter- pretationvwe say thataisacceptablewith respect tovif and only if∀(B, a)∈X∃a⁰∈B:v(a⁰) =fandaisunacceptablewith re- spect tovif and only if∃(B, a)∈X∀a⁰∈B:v(a⁰) =t.

For an interpretationv:A→ {t,f,u}it holds that

• v∈adm(S)iff for alla∈A,ais acceptable wrt.vifv(a) =t andais unacceptable wrt.vifv(a) =f;

• v∈com(S)iff for alla∈A,ais acceptable wrt.viffv(a) =t andais unacceptable wrt.viffv(a) =f;

• v∈prf(S)iffvis≤i-maximal admissible; and

• v∈mod(S)iffv∈adm(F)and@a∈A:v(a) =u.

For a SETAFS= (A, X)the corresponding ADFDS has accept- ance formulaϕa=V

(B,a)∈X

W

a⁰∈B¬a⁰for each statementa∈A.

Proposition 1. For any SETAF S= (A, X) it holds that σ(S) =σ(DS), whereσ∈ {adm,com,prf,mod}.

Proof. Given interpretation v and statement a, it holds that ΓD_S(v)(a) =t iff ∀w∈[v]2:w(a) =t iff ∀(B, a)∈X

∃a⁰∈B:v(a⁰) =f iffais acceptable wrt.vandΓD_S(v)(a) =f iff ∀w∈[v]2:w(a) =f iff ∃(B, a)∈X ∀a⁰∈B:v(a⁰) =t iff a is unacceptable wrt. v. Hence σ(S) =σ(DS) for

σ∈ {adm,com,prf,mod}. 2

Realizability A set V ⊆ V of interpretations is realizable in a formalismFunder a semanticsσif and only if there exists a know- ledge basekb∈ F having exactlyσ(kb) =V. P¨uhrer [26] charac- terized realizability for ADFs under various three-valued semantics.

We will reuse the central notions for capturing the complete se- mantics in this work.

Definition 1 (P ¨uhrer [26]). Let V be a set of interpretations. A functionf:V2→ V2 is acom-characterizationofV iff: for each v∈ Vwe havev∈V iff for eacha∈A:

• v(a)6=uimpliesf(v2)(a) =v(a)for allv2∈[v]2and

• v(a) =u implies f(v2⁰)(a) =t and f(v⁰⁰2)(a) =f for some

v₂⁰, v₂⁰⁰∈[v]2. N

Intuitively, a com-characterization f assigns the Boolean value f(v)(a)to a statementathat the acceptance condition ofawould have undervin an ADF that hasV as its complete semantics. From a function of this kind we can build a corresponding ADF by the following construction. Forf:V2→ V2, we defineDf as the ADF where the acceptance formula for each statementais given by

ϕ^fa= _

w∈V₂, f(w)(a)=t

φw with φw= ^

w(a⁰)=t

a⁰∧ ^

w(a⁰)=f

¬a⁰

Observe that for anyv∈ V2we havev(φw) =tiffv=wby defini- tion. Intuitively, the acceptance conditionϕ^f_ais constructed such that vis a model ofϕ^faif and only if we findf(v)(a) =t.

Proposition 2 (P ¨uhrer [26]). Let V ⊆ V be a set of interpreta- tions. (1) For each ADFD with com(D) =V, there is a com- characterization fD for V; (2) for each com-characterization f:V2→ V2forV we havecom(Df) =V.

The result shows thatV can be realized under complete semantics if and only if there is acom-characterization forV.

3 A General Framework for Realizability

The underlying idea of our framework is that all abstract argument- ation formalisms introduced in the previous section can be viewed as subclasses of ADFs. This is clear for ADFs themselves and for BADFs by definition; for (SET)AFs it is fairly easy to see. However, knowing that these formalisms can be recast as ADFs is not enough.

To employ this knowledge for realizability, we must precisely charac- terize the corresponding subclasses in terms of restricting the ADFs’

acceptance functions. Fortunately, this is also possible and paves the way for the framework we present in this section. Most importantly, we will make use of the fact that different formalisms and different semantics can be characterized modularly, that is, independently of each other.

Towards a uniform account of realizability for ADFs under differ- ent semantics, we start with a new characterization of realizability for ADFs under admissible semantics that is based on a notion similar in spirit tocom-characterizations.

Definition 2. LetV be a set of interpretations. A functionf:V2→ V2 is an adm-characterizationofV iff: for eachv ∈ V we have v∈V iff for everya∈A:

• v(a)6=uimpliesf(v2)(a) =v(a)for allv2∈[v]2. N Similar as for acom-characterization, anadm-characterizationfas- signs the valuef(v)(a)to a statementathat the acceptance condition ofawould evaluate to undervin an ADF that hasV as its admissible semantics. Note that the only difference to Definition 1 is dropping the second condition related to statements with truth valueu. While,

the two conditions in Definition 1 capture the relationΓD_f(v) =v, the remaining one in Definition 2 boils down tov≤i ΓD_f(v)that defines the admissible semantics.

Proposition 3. LetV ⊆ Vbe a set of interpretations. (1) For each ADFDsuch thatadm(D) =V, there is anadm-characterization fDforV; (2) for each adm-characterizationf:V2→ V2forV we haveadm(Df) =V.

Proof. (1) We define the function fD:V2→ V2 as fD(v2)(a) =v2(ϕa) for every v2∈ V2 and a∈A where ϕa

is the acceptance formula ofain D. We will show thatfD is an adm-characterization forV =adm(D). Letvbe an interpretation.

Consider the case v∈adm(D) and v(a)6=u for some a∈A and somev2 ∈[v]2. Fromv≤iΓD(v)we getv2(ϕa) =v(a). By definition of fD it follows that fD(v2)(a) =v(a). Now assume v6∈adm(D)and consequently v6≤iΓD(v). There must be some a∈A such thatv(a)6=uandv(a)6= ΓD(v)(a). Hence, there is some v2∈[v]2 with v2(ϕa)6=v(a) and fD(v2)(a)6=v(a) by definition offD. Thus,fDis anadm-characterization forV.

(2) Observe that for every two-valued interpretationv2and every a∈Awe havef(v2)(a) =v2(ϕ^fa).(⊆): Letv∈adm(Df)be an interpretation and a∈A a statement such thatv(a)6=u. Let v2

be a two-valued interpretation withv2∈[v]2. Sincev≤iΓD_f(v) we havev(a) =v2(ϕ^fa). Therefore, by our observation it must also hold thatf(v2)(a) =v(a). Thus, by Definition 2,v∈V.(⊇): Con- sider an interpretation v such that v6∈adm(Df). We show that v6∈V. Fromv6∈adm(Df)we getv6≤iΓD_f(v). There must be some a∈A such that v(a)6=uand v(a)6= ΓD_f(v)(a). Hence, there is some v2∈[v]2 with v2(ϕ^f_a)6=v(a) and consequently f(v2)(a)6=v(a). Thus, by Definition 2 we havev6∈V. 2 When listing sets of interpretations in examples, for the sake of readability we represent three-valued interpretations by sequences of truth values, tacitly assuming that the underlying vocabu- lary is given and has an associated total ordering. For example, for the vocabulary A={a, b, c} we represent the interpretation {a7→t, b7→f, c7→u}by the sequencetf u.

Example 2. Consider the sets V1={uuu,tf f,f tu} and V2={tf f,f tu}of interpretations overA={a, b, c}. The mapping f ={ttt7→f tt,ttf 7→tf t,tf t7→ttt,tf f 7→tf f,f tt7→f tf, f tf 7→ f tt,f f t 7→ ttf,f f f 7→ f tf}is anadm-characterization forV1. Thus, the ADFDf hasV1as its admissible interpretations.

Indeed, the realizing ADF has the following acceptance conditions:

ϕ^f_a≡(a∧b∧ ¬c)∨(a∧ ¬b)∨(¬a∧ ¬b∧c) ϕ^f_b ≡(a∧c)∨(¬a∧b)∨(¬a∧ ¬b∧ ¬c) ϕ^fc ≡(a∧b)∨(¬a∧b∧ ¬c)∨(¬b∧c)

ForV2 noadm-characterization exists becauseuuu6∈V2, but the implication of Definition 2 trivially holds fora,b, andc.

We have seen that the constructionDf for realizing under com- plete semantics can also be used for realizing a setV of interpret- ations under admissible semantics. The only difference is that we here require f to be an adm-characterization instead of a com- characterization forV. Note that admissible semantics can be char- acterized by properties that are easier to check than existence of an adm-characterization (see the work of P¨uhrer [26]). However, us- ing the same type of characterizations for different semantics allows for a unified approach for checking realizability and constructing a realizing ADF in case one exists.

For realizing under the model semantics, we can likewise present an adjusted version ofcom-characterizations.

Definition 3. Let V ⊆ V be a set of interpretations. A function f:V2→ V2 is amod-characterizationofV if and only if: (1)f is defined onV (that is,V ⊆ V2) and (2) for eachv∈ V2, we have

v∈V ifff(v) =v. N

As we can show, there is a one-to-one correspondence between mod-characterizations and ADF realizations.

Proposition 4. LetV ⊆ V be a set of interpretations. (1) For each ADFDsuch thatmod(D) =V, there is amod-characterizationfD

forV; (2) vice versa, for eachmod-characterization f:V2→ V2

forV we findmod(Df) =V.

A related result was given by Strass [28, Proposition 10]. The char- acterization we presented here fits into the general framework of this paper and is directly usable for our realizability algorithm. The next result summarizes how ADF realizability can be captured by differ- ent types of characterizations for the semantics we considered so far.

Theorem 5. Let V ⊆ V be a set of interpretations and consider σ∈ {adm, com, mod}. There is an ADFDsuch thatσ(D) =V if and only if there is aσ-characterization forV.

The preferred semantics of an ADFDis closely related to its ad- missible semantics as, by definition, the preferred interpretations of D are its ≤i-maximal admissible interpretations. As a con- sequence we can also describe preferred realizability in terms of adm-characterizations. We use the lattice-theoretic standard notation max≤_iV to denote the≤i-maximal elements of a given setV. Corollary 6. LetV ⊆ Vbe a set of interpretations. There is an ADF Dwithprf(D) =V iff there is anadm-characterization for some V⁰⊆ VwithV ⊆V⁰andmax≤_iV⁰=V.

Finally, we give a result on the complexity of deciding realizability for the mentioned formalisms and semantics. We assume here that the representation of an interpretation-setV over vocabularyAhas sizeΘ(3^|A|), that is, the size grows asymptotically in the order of 3^|A|. A possible encoding could be a bit string of length3^|A|where the presence (or absence) of eachv∈V is encoded by a 1 (or 0) at a particular position in the string. There might be specificV with smaller possible representations, but we have no grounds to presume a representation that is exponentially better in the general case.

Proposition 7. Let F ∈ {AF,SETAF,BADF,ADF}be a formal- ism andσ∈ {adm,com,prf,mod}be a semantics. The decision problem “Given a vocabularyAand a setV ⊆ Vof interpretations overA, is there akb∈ Fsuch thatσ(kb) =V?” can be decided in nondeterministic time that is polynomial in the size ofV.

Proof. Roughly, we guess a functionf:V2 → V2 and verify that it is aσ-characterization. Such a functionf can be represented in sizeO(2^|A|· |A|), that is, at most polynomial in the input of size O(3^|A|): the fact thatn·2ⁿ∈o(3ⁿ)⊆O(3ⁿ)follows from

n→∞lim n·2ⁿ

3ⁿ = lim

n→∞

n

3 2

n

= lim∗ n→∞

1

ln³₂· ³₂n = 0 where the starred equality holds by L’Hˆopital’s rule.

To verify that the guessedf is indeed aσ-characterization, we check (some of) the properties of Definition 1. Forσ=com, this can be done in polynomial time as follows: for eachv∈ Vanda∈A, we

look at the set[v]2⊆ V2 (which is at most polynomial in the input) and check for the respective witness interpretations (if v(a) =u) or their absence (ifv(a)6=u). Forσ=adm, there are even less conditions to check. For σ=mod, we compute the setF of fix- points off (by going through Vonce and checkingf(v) =vfor eachv∈ V) and verify thatF =V. Forσ=prf, we guess theV⁰ (withV ⊆V⁰⊆ V) from Corollary 6 alongsidefand verify thatf is anadm-characterization forV⁰and thatmax≤_iV⁰=V. 2

3.1 Deciding Realizability: Algorithm 1

Our main algorithm for deciding realizability is a propagate-and- guess algorithm in the spirit of the DPLL algorithm for deciding propositional satisfiability [19]. It is generic with respect to (1) the formalismFand (2) the semanticsσfor which should be realized.

To this end, the propagation part of the algorithm is kept exchange- able and will vary depending on formalism and semantics. Roughly, in the propagation step the algorithm uses the desired setV of in- terpretations to derive certain necessary properties of the realizing knowledge base (line 2). This is the essential part of the algorithm:

the derivation rules (propagators) used there are based on charac- terizations of realizability with respect to formalism and semantics.

(Propagators will be explained in detail in the next two subsections.) Once propagation of properties has reached a fixed point (line 7), the algorithm checks whether the derived information is sufficient to construct a knowledge base. If so, the knowledge base can be con- structed and returned (line 9). Otherwise (no more information can be obtained through propagation and there is not enough information to construct a knowledge base yet), the algorithm guesses another as- signment for the characterization (line 11) and calls itself recursively.

The main data structure that Algorithm 1 operates on is a set of triples(v, a,x)consisting of a two-valued interpretationv∈ V2, an atoma∈Aand a truth valuex∈ {t,f}. This data structure is inten- ded to represent theσ-characterizations introduced in Definitions 1 to 3. There, aσ-characterization is a functionf:V2→ V2from two- valued interpretations to two-valued interpretations. However, as the algorithm builds theσ-characterization step by step and there might not even be aσ-characterization in the end (becauseV is not real- izable), we use a setF of triples(v, a,x)to be able to represent both partial and incoherent states of affairs. Theσ-characterization candidate induced byFis partial if we have that for somevanda, neither(v, a,t)∈Fnor(v, a,f)∈F; likewise, the candidate is in- coherent if for somevanda, both(v, a,t)∈F and(v, a,f)∈F.

IfF is neither partial nor incoherent, it gives rise to a unique σ- characterization that can be used to construct the knowledge base realizing the desired set of interpretations. The correspondence to the characterization-function is then such that f(v)(a) =x iff (v, a,x)∈F.

In our presentation of the algorithm we focused on its main fea- tures, therefore the guessing step (line 11) is completely “blind”. It is possible to use techniques known from constraint satisfaction prob- lems, such as shaving (removing guessing possibilities that directly lead to inconsistency). Finally, we remark that the algorithm can be extended to enumerate all possible realizations of a given interpreta- tion set – by keeping all choice points in the guessing step and thus exhaustively exploring the whole search space.

In the case where the constructed relationF becomes functional at some point, the algorithm returns a realizing knowledge base kb^F_σ(F). For ADFs, this just means that we denote by f the σ- characterization represented byFand setkb^ADF_σ (F) =Df. For the remaining formalisms we will introduce the respective constructions

Algorithm 1realize(F, σ, V, F) Input: • a formalismF

• a semanticsσforF

• a setV of interpretationsv:A→ {t,f,u}

• a relationF⊆ V2×A× {t,f}, initially empty Output: akb∈ Fwithσ(kb) =V or “no” if none exists

1: repeat

2: setF∆:= S

p∈P^F_σ

p(V, F)\F

3: setF :=F∪F∆

4: if∃v∈ V2,∃a∈A:{(v, a,t),(v, a,f)} ⊆F then 5: return“no”

6: end if 7: until F∆=∅

8: if∀v∈ V2,∀a∈A,∃x∈ {t,f}: (v, a, x)∈Fthen 9: returnkb^Fσ(F)

10: end if

11: choosev∈ V2, a∈Awith(v, a,t)∈/F,(v, a,f)∈/F 12: ifrealize(F, σ, V, F∪ {(v, a,t)})6=“no”then 13: returnrealize(F, σ, V, F∪ {(v, a,t)}) 14: else

15: returnrealize(F, σ, V, F∪ {(v, a,f)}) 16: end if

in later subsections.

The algorithm is parametric in two dimensions, namely with re- spect to the formalismFand with respect to the semanticsσ. These two aspects come into the algorithm via so-calledpropagators. A propagator is a formalism-specific or semantics-specific set of deriv- ation rules. Given a setV of desired interpretations and a partialσ- characterizationF, a propagatorpderives new triples(v, a,x)that must necessarily be part of any totalσ-characterizationfforV such thatfextendsF. In what follows, we present semantics propagators for admissible, complete and two-valued model (in (SET)AF terms stable) semantics, and formalism propagators for BADFs, AFs, and SETAFs.

3.2 Semantics Propagators

The semantics propagators are defined in Figure 1. They are dir- ectly derived from the properties ofσ-characterizations presented in Definitions 1 to 3. While the definitions provide exact conditions to check whether a given function is aσ-characterization, the propagat- ors allow us to derive definite values of partial characterizations that are necessary to fulfill the conditions for being aσ-characterization.

For admissible semantics, the condition for a functionf to be an adm-characterization of a desired set of interpretationsV (cf. Defin- ition 2) can be split into a condition for desired interpretationsv∈V and two conditions for undesired interpretationsv /∈V. Propagator p^∈_admderives new triples by considering interpretationsv∈V. Here, for all two-valued interpretationsv2 that extendv, the valuef(v2) has to be in accordance withvonv’s Boolean part, that is, the al- gorithm adds(v2, a, v(a))wheneverv(a)6=u. On the other hand, p^∈_adm^/ derives new triples forv /∈V in order to ensure that there is a two-valued interpretationv2extendingvwheref(v2)differs fromv on a Boolean value ofv. Note that whilep^∈_admimmediately allows us to derive information aboutFfor each desired interpretationv∈V, propagatorp^∈_adm^/ is much weaker in the sense that it only derives a triple ofF if there is no other way to meet the conditions for an undesired interpretation. Special treatment is required for the inter- pretationvuthat maps all statements touand is admissible for every

p^∈_adm(V, F) ={(v2, a, v(a)) | v∈V, v₂∈[v]₂, v(a)6=u} p^∈,u_com(V, F) ={(v2, a,¬x)| v∈V, v₂∈[v]₂, v(a) =u, p^∈_adm^/ (V, F) ={(v₂, a,¬v(a)) | v∈ V \V, v₂∈[v]₂, x∈ {t,f},∀v₂⁰∈[v]₂:v₂6=v⁰₂→(v⁰₂, a,x)∈F}

v(a)6=u,∀b∈A\v⁻¹(u),∀v⁰₂∈[v]₂: p^6∈,tf_com(V, F) ={(v2, a,¬v(a)) | v∈ V \V, v₂∈[v]₂, v(a)6=u, (a, v2)6= (b, v₂⁰)→(v⁰₂, b, v(b))∈F} ∀b∈A\v⁻¹(u),∀v₂⁰∈[v]2: (a, v2)6= (b, v⁰₂)→(v⁰₂, b, v(b))∈F, p_adm(V, F) ={(v, a,t),(v, a,f) | v∈ V₂, a∈A, v_u6∈V} ∀b∈v⁻¹(u),∃v⁰⁰₂, v⁰⁰⁰₂ ∈[v]₂: (v⁰⁰₂, b,t),(v⁰⁰⁰₂ , b,f)∈F}

p^∈_mod(V, F) ={(v, a, v(a)) |v∈V, a∈A} p^6∈,u_com(V, F) ={(v2, a,¬x)| v∈ V \V, v2∈[v]2, v(a) =u, p^∈_mod^/ (V, F) ={(v, a,¬v(a)) |v∈ V2\V, a∈A, ∀b∈A\v⁻¹(u),∀v₂⁰∈[v]2: (v2, b, v(b))∈F,

∀c∈A\ {a}: (v, c, v(c))∈F} ∀b∈v⁻¹(u)\ {a}:∃v⁰⁰₂, v⁰⁰⁰₂ ∈[v]2: (v₂⁰⁰, b,t), p_mod(V, F) ={(v, a,t),(v, a,f) | v∈ V2, a∈A, V 6⊆ V2} (v⁰⁰⁰₂ , b,f)∈F,∀v⁰⁰⁰⁰₂ ∈[v]₂\ {v2}: (v₂⁰⁰⁰⁰, b,x)∈F}

Figure 1: Semantics propagators for the complete (P_com^ADF={p^∈,tf_com, p^∈,ucom, p^6∈,tfcom, p^6∈,ucom} with p^∈,tfcom(V, F) =p^∈_adm(V, F)), admissible (P_adm^ADF ={p^∈_adm, p^∈_adm^/ , p_adm}), and model semantics (P_mod^ADF={p^∈_mod, p^∈_mod^/ , p_mod}).

ADF. This is not captured byp^∈_admandp^∈_adm^/ as these deal only with interpretations that have Boolean mappings. Thus, propagatorp_adm serves to check whethervu∈V. If this is not the case, the propag- ator immediately makes the relationFincoherent and the algorithm correctly answers “no”.

For complete semantics and interpretations v∈V, propagator p^∈,tf_com derives triples just like in the admissible case. Propagator p^∈,ucom deals with statementsa∈Ahavingv(a) =ufor which there have to be at least two v2, v⁰₂∈[v]2 having f(v2)(a) =t and f(v2⁰)(a) =f. Hencep^∈,ucomderives triple(v2, a,¬x)if for all other v⁰₂∈[v]2 we find a triple (v⁰₂, a,x). For interpretationsv /∈V it must hold that there is some a∈A such that (i) v(a)6=u and f(v2)(a)6=v(a) for somev2∈[v]2 or (ii) v(a) =u but for all v2∈[v]2,f(v2)assigns the same Boolean truth valuextoa. Now if neither (i) nor (ii) can be fulfilled by any statementb∈A\ {a}

due to the current contents ofF, propagatorsp^6∈,tf_com andp^6∈,u_com derive triple(v2, a,¬v(a))forv(a)6=uif needed forato fulfill (i) and (v2, a,¬x)forv(a) =uif needed forato fulfill (ii), respectively.

Example 3. Consider the set V3={uuu,f uu,uuf,f tf}. First, we consider a run of realize(ADF,adm, V3,∅). In the first itera- tion, propagatorp^∈_adm ensures thatF∆in line 2 contains(f f f, a,f), (f tf, a,f), (f tf, c,f), and (f f f, c,f). Based on the latter three tuples and f uf∈/V3, propagator p^∈_adm^/ derives (f f f, a,t) in the second iteration which together with (f f f, a,f) causes the al- gorithm to return “no”. Consequently, V3 is not adm-realizable.

A run of realize(ADF,com, V3,∅) on the other hand returns com-characterization f for V3 that maps ttf to tf f, f tt to f f t, f tf and f f f to f tf and all other v2∈ V2 to f f f. Hence, ADF Df, given by the acceptance conditions ϕ^fa=a∧b∧ ¬c, ϕ^f_b = (¬a∧b∧ ¬c)∨(¬a∧ ¬b∧ ¬c), andϕ^f_c =¬a∧b∧c, has

V3as its complete semantics.

Finally, for two-valued model semantics, propagator p^∈_mod derives new triples by looking at interpretationsv∈V. For those, we must findf(v) =vin eachmod-characterizationf by definition. Thus the algorithm adds(v, a, v(a))for eacha∈Ato the partial charac- terizationF. Propagatorp^∈_mod^/ looks at interpretationsv∈ V2\V, for which it must hold thatf(v)6=v. Thus there must be a statement a∈Awithv(a)6=f(v)(a), which is exactly what this propagator derives whenever it is clear that there is only one statement can- didate left. This, in turn, is the case whenever all b∈Awith the opposite truth value ¬v(a) and allc∈A withc6=a cannot co- herently become the necessary witness any more. The propagator p_mod checks whetherV ⊆ V2, that is, the desired set of interpret-

Algorithm 2realizePrf(F, V) Input: • a formalismF

• a setV of interpretationsv:A→ {t,f,u}

Output: Return somekb∈ F withprf(kb) =V if one exists or

“no” otherwise.

1: ifmax≤_iV 6=V then 2: return“no”

3: end if

4: setV^<i :={v∈ V | ∃v⁰∈V :v <iv⁰} 5: setX:=∅

6: repeat

7: chooseV⁰⊆V^<iwithV⁰∈/X 8: setX :=X∪ {V⁰}

9: setV^adm:=V ∪V⁰

10: ifrealize(F,adm, V^adm,∅)6=“no”then 11: returnrealize(F,adm, V^adm,∅) 12: end if

13: until∀V⁰⊆V^<i:V⁰∈X 14: return“no”

ations consists entirely of two-valued interpretations. In that case this propagator makes the relationFincoherent, following a similar strategy asp_adm.

The Special Case of Preferred Semantics Realizing a given set of interpretationsV under preferred semantics requires special treat- ment. We do not have a σ-characterization function for σ=prf at hand to directly check realizability ofV but have to find some V⁰⊆ {v∈ V | ∃v⁰∈V :v <iv⁰} such that V ∪V⁰ is realizable under admissible semantics (cf. Corollary 6). Algorithm 2 imple- ments this idea by guessing such a V⁰ (line 7) and then using Algorithm 1 to try to realize V ∪V⁰ under admissible semantics (line 11). Ifrealizereturns a knowledge basekbrealizingV ∪V⁰ underadmwe can directly usekbas solution ofrealizePrf since it holds thatprf(kb) =V, given thatV is an≤i-antichain (line 2).

3.3 Formalism Propagators

When constructing an ADF realizing a given setV of interpretations under a semanticsσ, the functionkb^ADFσ (F)makes use of theσ- characterization given byFin the following way:vis a model of the acceptance conditionϕaif and only if we find(v, a,t)∈F. Now as bipolar ADFs, SETAFs and AFs are all subclasses of ADFs by restricting the acceptance conditions of statements, these restrictions also carry over to theσ-characterizations. The propagators defined

p^SETAF(V, F) ={(vf, a,t)|a∈A} ∪ {(w, a,t)|(v, a,t)∈F, w∈ V2, w <tv} ∪ {(w, a,f)|(v, a,f)∈F, w∈ V2, v <tw}

p^AF(V, F) =p^SETAF(V, F)∪ {(v1ttv2, a,t)|(v1, a,t)∈F,(v2, a,t)∈F} L⁺=n (b, a)

(v, a,f)∈F, v(b) =f,(v|^b_t, a,t)∈Fo p^BADF(V, F) ={(v|^b_t, a,x)|(v, a,x)∈F,(w, a,¬x)∈F, w(b) =f,(w|^b_t, a,x)∈F} L⁻=n

(b, a)

(v, a,t)∈F, v(b) =f,(v|^b_t, a,f)∈Fo

Figure 2: Formalism propagators. For formalismF ∈ {AF,SETAF,BADF}and anyσ∈ {adm,com,prf,mod}, we set the respective propagator forFto P_σ^F=P_σ^ADF∪

p^F withp^Fas defined above.L⁺andL⁻define link polarities forkb^BADF_σ .

in Figure 2 use structural knowledge on the form of acceptance con- ditions of the respective formalisms to reduce the search space or to induce incoherence ofFwheneverV is not realizable.

Bipolar ADFs For bipolar ADFs, we use the fact that each of their links must have at least one polarity, that is, must be supporting or attacking. Therefore, if a link is not supporting, it must be attacking, and vice versa. For canonical realization, we obtain the polarities of links, that is, the setsL⁺andL⁻, as defined in Figure 2.

AFs To explain the AF propagators, we first need some more defin- itions. On the two classical truth values, we define the truth order- ingf<tt, whence the operations tt andut withfttt=tand futt=f result. These operations can be lifted pointwise to two- valued interpretations as usual, i.e.,(v1ttv2)(a) =v1(a)ttv2(a) and(v1utv2)(a) =v1(a)utv2(a). Again, the reflexive version of

<tis denoted by≤t. The pair(V2,≤t)of two-valued interpretations ordered by the truth ordering forms a complete lattice with glbutand lubtt. This complete lattice has the least elementvf :A→ {f}, the interpretation mapping all statements to false, and the greatest elementvt:A→ {t}mapping all statements to true, respectively.

Acceptance conditions of AF-based ADFs have the form of con- junctions of negative literals. In the complete lattice (V2,≤t), the model sets of AF acceptance conditions correspond to the lattice- theoretic concept of an ideal, a subset of V2 that is downward- closed with respect to≤tand upward-closed with respect tott. The propagator directly implements these closure properties: application ofp^AF ensures that when aσ-characterizationF that is neither in- coherent nor partial is found in line 8 of Algorithm 1, then there is, for eacha∈A, an interpretationvasuch that(va, a,t)∈Fand v≤tvafor each(v, a,t)∈F. Henceva is crucial for the accept- ance condition, or in AF terms the attacks, ofaand we can define kb^AFσ (F) = (A,{(b, a)|a, b∈A, va(b) =f}).

SETAFs The propagator for SETAFs,p^SETAF, is a weaker version of that of AFs, since we cannot presume upward-closure with respect to tt. In SETAF-based ADFs the acceptance formula is inconjunctive normal formcontaining only negative literals. By a transformation preserving logical equivalence we obtain an acceptance condition in disjunctive normal form, again with only negative literals; in other words, adisjunctionof AF acceptance formulas. Thus, the model set of a SETAF acceptance condition is not necessarily an ideal, but a union of ideals. For the canonical realization we can make use of the fact that, for eacha∈A, the setV_a^t={v∈ V2|(v, a,t)∈F} is downward-closed with respect to≤t, hence the set of models of W

v∈max_≤ tV^t

V

v(b)=f¬bis exactlyV_a^t. The clauses of its corres- ponding CNF-formula exactly coincide with the sets of arguments attackingainkb^SETAFσ (F).

3.4 Correctness

For a lack of space, we could not include a formal proof of soundness and completeness of Algorithm 1, but rather present arguments for termination and correctness.

Termination With each recursive call, the setFcan never decrease in size, as the only changes toFare adding the results of propagation in line 3 and adding the guesses in line 11. Also within the until-loop, the setF can never decrease in size; furthermore there is only an overall finite number of triples that can be added toF. Thus at some point we must haveF∆=∅and leave the until-loop. SinceFalways increases in size, at some point it must either become functional or incoherent, whence the algorithm terminates.

Soundness If the algorithm returnskb^Fσ(F)as a realizing know- ledge base, then according to the condition in line 8 the relationFin- duced a total functionf:V2→ V2. In particular, because the until- loop must have been run through at least once, there was at least one propagation step (line 2). Since the propagators are defined such that they enforce everything that must hold in aσ-characterization, we conclude that the induced functionfindeed is aσ-characterization forV. By construction, we consequently find thatσ(kb^F_σ(F)) =V. Completeness If the algorithm answers “no”, then the execution reached line 5. Thus, for the constructed set F, there must have been an interpretation v∈ V2 and a statement a∈A such that {(v, a,t),(v, a,f)} ⊆F, that is,Fis incoherent. SinceFis initially empty, the only way it could get incoherent is in the propagation step in line 2. (The guessing step cannot create incoherence, since exactly one truth value is guessed forvanda.) However, the propagators are defined such that they infer only assignments (triples) that are necessary for the givenF. Consequently, the given interpretation set V is such that either there is no realization within the ADF fragment corresponding to formalismF(that is, the formalism propagator de- rived the incoherence) or there is noσ-characterization forV with respect to general ADFs (that is, the semantics propagator derived the incoherence). In any case,V is notσ-realizable forF.

4 Implementation

As Algorithm 1 is based on propagation, guessing, and checking it is perfectly suited for an implementation using answer set pro- gramming (ASP) [24, 21] as this allows for exploiting conflict learn- ing strategies and heuristics of modern ASP solvers. Thus, we de- veloped ASP encodings in thegringolanguage [17] for our ap- proach. Similar as the algorithm, our declarative encodings are mod- ular, consisting of a main part responsible for constructing set F and separate encodings for the individual propagators. If one wants, e.g., to compute an AF realization under admissible semantics for a setV of interpretations, an input program encodingV is joined with the main encoding, the propagator encoding for admissible semantics as well as the propagator encoding for AFs. Every an- swer set of such a program encodes a respective characterization function. Our ASP encoding for preferred semantics is based on the admissible encoding and guesses further interpretations follow- ing the essential idea of Algorithm 2. For constructing a know- ledge base with the desired semantics, we also provide two ASP encodings that transform the output to an ADF in the syntax of

the DIAMOND tool [14], respectively an AF in ASPARTIX syn- tax [13, 15]. Both argumentation tools are based on ASP themselves.

The encodings for all the semantics and formalisms we covered in the paper can be downloaded from http://www.dbai.tuwien.

ac.at/research/project/adf/unreal/.

5 Expressiveness Results

In this section we briefly present some results that we have obtained using our implementation. We first introduce some necessary nota- tion to describe the relative expressiveness of knowledge representa- tion formalisms [18, 28]. For formalismsF1andF2with semantics σ1 andσ2, we say thatF2underσ2is at least as expressive asF1

underσ1and writeF₁^σ1 ≤eF₂^σ2 if and only ifΣ^σ_F¹₁ ⊆Σ^σ_F²₂, where Σ^σF ={σ(kb)|kb∈ F }is thesignature ofF underσ. As usual, we defineF1<eF2if and only ifF1≤eF2andF26≤eF1.

We now start by considering the signatures of AFs, SETAFs and (B)ADFs for the unary vocabulary{a}:

Σ^adm_AF = Σ^adm_SETAF={{u},{u,t}}

Σ^com_AF = Σ^com_SETAF={{u},{t}}

Σ^prf_AF = Σ^prf_SETAF={{u},{t}}

Σ^mod_AF = Σ^mod_SETAF={∅,{t}}

Σ^adm_ADF = Σ^adm_BADF= Σ^adm_AF ∪ {{u,f},{u,t,f}}

Σ^com_ADF = Σ^com_BADF= Σ^com_AF ∪ {{f},{u,t,f}}

Σ^prf_ADF= Σ^prf_BADF= Σ^prf_AF ∪ {{f},{t,f}}

Σ^mod_ADF = Σ^mod_BADF= Σ^mod_AF ∪ {{f},{t,f}}

The following result shows that the expressiveness of the formalisms under consideration is in line with the amount of restrictions they impose on acceptance formulas.

Theorem 8. For anyσ∈ {adm,com,prf,mod}:

1. AF^σ<eSETAF^σ. 2. SETAF^σ<eBADF^σ. 3. BADF^σ<eADF^σ.

Proof. (1) AF^σ≤eSETAF^σ is clear (by modeling individual at- tacks via singletons). For SETAF^σ6≤eAF^σthe witnessing interpret- ation sets over vocabularyA={a, b, c}are{uuu,ttf,tf t,f tt} ∈ Σ^σ_SETAF \ Σ^σAF and {ttf,tf t,f tt} ∈ Σ^τ_SETAF \ Σ^τAF with σ∈ {adm,com} and τ∈ {prf,mod}. By each pair of argu- ments of A being t in at least one model, a realizing AF can- not feature any attack, immediately giving rise to the model ttt.

The respective realizing SETAF is given by the attack relation X={({a, b}, c),({a, c}, b),({b, c}, a)}.

(2) It is clear that SETAF^σ≤eBADF^σholds (SETAFs are bipolar since all parents are always attacking). For BADF^σ6≤eSETAF^σthe respective counterexamples can be read off the signatures above:

for σ∈ {adm,com} we find {u,t,f} ∈Σ^σBADF\Σ^σSETAF and for τ ∈ {prf,mod}we find{t,f} ∈Σ^τ_BADF\Σ^τ_SETAF. The realizing bi- polar ADF has acceptance conditionϕa=a.

(3) Forσ=mod the result is known [28, Theorem 14]; for the remaining semantics the model sets witnessing ADF^σ6≤eBADF^σ over vocabularyA={a, b}are

{uu,tu,tt,tf,f u} ∈Σ^admADF \Σ^admBADF

{uu,tu,tt,tf,f u} ∈Σ^comADF\Σ^comBADF

{tt,tf,f u} ∈Σ^prf_ADF\Σ^prf_BADF

A witnessing ADF is given byϕa=aandϕb=a↔b. 2 Theorem 8 is concerned with the relative expressiveness of the formalisms under consideration, given a certain semantics. Consider- ing different semantics we find that for all formalisms the signatures become incomparable:

Proposition 9. F₁^σ16≤eF₂^σ2 and F₂^σ26≤eF₁^σ1 for all form- alisms F1,F2∈ {AF,SETAF,BADF,ADF} and all semantics σ1, σ2∈ {adm,com,prf,mod}withσ16=σ2.

Proof. First, the result foradmandcomfollows by{u,t} ∈Σ^adm_AF , but {u,t}∈/Σ^com_ADF and {t} ∈Σ^com_AF , but {t}∈/Σ^adm_ADF. Moreover, taking into account that the set of preferred interpretations (resp. two- valued models) always forms a≤i-antichain while the set of admiss- ible (resp. complete) interpretations never does, the result follows for σ1∈ {adm,com}andσ2∈ {prf,mod}. Finally, since akb∈ F may not have any two-valued models and a preferred interpretation is not necessarily two-valued, the result forprf andmodfollows.2 Disregarding the possibility of realizing the empty set of interpret- ations under the two-valued model semantics, we obtain the follow- ing relation for ADFs.

Proposition 10. (Σ^modADF\ {∅})⊆Σ^prf_ADF.

Proof. Consider some V ∈Σ^modADF with V 6=∅. Clearly V ⊆ V2

and by Proposition 4 there is amod-characterizationf:V2→ V2

for V, that is, f(v) =v iff v∈V. Define f⁰:V2→ V2 such that f⁰(v) =f(v) =v for all v∈V and f⁰(v)(a) =¬v(a) for all v∈ V \V and a∈A. Now it holds that f⁰ is an adm- characterization of V⁰={v∈ V | ∀v2∈[v]2:v2∈V} ∪ {vu}.

Sincemax≤_iV⁰=V we get that the ADFDwith acceptance for- mulaϕ^fa⁰ for eacha∈Ahasprf(D) =V whenceV ∈Σ^prf_ADF. 2 In contrast, this relation does not hold for AFs, which was shown for extension-based semantics by Linsbichler et al. [20, Theorem 5] and immediately follows for the three-valued case.

6 Discussion

We presented a framework for realizability in which AFs, SETAFs, BADFs and general ADFs can be treated in a uniform way. The centerpiece of our approach is an algorithm for deciding realizab- ility of a given interpretation-set in a formalism under a semantics.

The algorithm makes use of so-called propagators, by which it can be adapted to the different formalisms and semantics. We also presented an implementation of our framework in answer set programming and several novel expressiveness results that we obtained using our im- plementation. In unpublished related work, our colleague Sylwia Pol- berg studied a wide range of abstract argumentation formalisms, in particular their relationship with ADFs [25]. This can be the basis for including further formalisms into our realizability framework: all that remains to do is figuring out suitable ADF fragments and developing propagators for them, just like we did exemplarily for Nielsen and Parsons’ SETAFs. For further future work, several semantics whose realizability is yet unstudied could be added to our framework, for example semantics based on conflict-freeness, like three-valued ver- sions of conflict-free, naive, and stage semantics [27, 16, 29].

Acknowledgements This research was supported by the German Research Foundation (DFG) under project BR 1817/7-1 and the Aus- trian Science Fund (FWF) under projects I1102, I2854 and P25518.